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Efficient algorithms and data structures for VLSI CAD

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Efficient algorithms and data structures for VLSI CAD
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xii, 141 leaves : ill. ; 29 cm.

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Thesis (Ph. D.)--University of Florida, 1996.
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Includes bibliographical references (leaves 138-140).
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EFFICIENT ALGORITHMS AND DATA STRUCTURES FOR VLSI CAD











By

SEONGHUN CHO










A DISSERTATION PRESENTED TO THE GRADUATE SCHOOL OF THE UNIVERSITY OF FLORIDA IN PARTIAL FULFILLMENT OF THE REQUIREMENTS FOR THE DEGREE OF
DOCTOR OF PHILOSOPHY UNIVERSITY OF FLORIDA 1996














ACKNOWLEDGMENTS


My heartfelt appreciation goes to my advisor Professor Sartaj Sahni for giving me continued guidance in my thesis work. I thank him for the help, patience and support he provided throughout my stay in the University of Florida. Weekly meetings and discussions with him have spawned many ideas, for which I am thankful.

I would like to thank other members in my supervisory committee, Dr Li-Mmn Fu, Dr Theodore Johnson, Dr Sanguthevar Rajasekaran, and Dr Paul W. Chun for their interest and comments.

Thanks go to Venkat Thanvantri for his willingness to discuss the general subject of algorithms.

Thanks go to my wife Joungyim for her love, encouragement and patience. Finally, I would like to thank my parents for the love and support, without which I could not have pursued my doctoral studies. To them I dedicate this work.













TABLE OF CONTENTS




ACKNOWLEDGMENTS ............................. ii

LIST OF TABLES ................................. vii

LIST OF FIGURES ................................ x

ABSTRACT .... ....... .. ........... ............ xi

CHAPTERS

I INTRODUCTION ............................. 1
1.1 Background ..... .. ....... .. .. .......... 1
1.2 Dissertation Outline ......................... 2

2 MINIMUM AREA JOINING OF COMPACTED CELLS ........ 4
2.1 Introduction ..... ......... .. .. ...... .... 4
2.2 1-layer River Routing ......................... 9
2.3 Constraint Graph Representation .................. 16
2.4 Heuristics to Minimize Area ..................... 20
2.4.1 Heuristic I .. . . 20
2.4.2 Heuristic 2 .. .. ..... .. .. .. ... .... 21
2.4.3 Heuristic 3 ....... .. .. .. ... .... 23
2.5 Experimental Results . . . 25
2.6 Conclusion . . . 30

3 A NEW WEIGHT BALANCED BINARY SEARCH TREE ...... 32
3.1 Introduction ..... ......... .. .. .. ..... .... 32
3.2 Balanced Trees and Rotations .................... 34
3.3 fl-BBSTs ................................ 39
3.4 Search, Insert, and Delete in a O-BBST . 44
3.4.1 Search . . . 44
3.4.2 Insertion . . . 44
3.4.3 Deletion . . . 57
3.4.4 Enhancements . . . 63
3.4.5 Top Down Algorithms . . 64








3.5 Simple O-BBSTs . . . 66
3.6 BBSTs without Deletion . . 70
3.7 Experimental Results . . . 72
3.8 Conclusion . . . 98

4 WEIGHT BIASED LEFTIST TREES AND MODIFIED SKIP LISTS 105
4.1 Introduction . . . 105
4.2 Weight Biased Leftist Trees . . 106
4.3 Modified Skip Lists . . . 109
4.4 MSLs As Priority Queues . . 118
4.5 Experimental Results For Priority Queues . 122
4 6 Conclusion . . . 129

5 CONCLUSIONS . . . 135

A ABBREVIATIONS . . . 136


REFERENCES . . . . 138

BIOGRAPHICAL SKETCH . . . 141




























i %,
















LIST OF TABLES




2.1 Error rate (%) over optimal, I = 1 . . 27

2.2 Improvement (%) over Fang, 1 = 1 . . 27

2.3 Time taken, I = I . . . 28

2.4 Error rate (%) over optimal, I = 2 . . 29

2.5 Improvement over I = I cases . . 29

2.6 Improvement over Fang, I = 2 . . 30

2.7 Time taken, 1 = 2 . . . 31

3.1 The number of key comparisons on random inputs (version 1 code) 76

3.2 The number of key comparisons on ordered inputs (version 1 code) 77

3.3 Height of the trees on random inputs (version 1 code) . 77

3.4 Height of the trees on ordered inputs (version 1 code) . 77

3.5 The number of rotations on random inputs (version I code) 79

3.6 The number of rotations on ordered inputs (version I code) 80

3.7 Run time on random inputs using integer keys (version I code) 81

3.8 Run time on ordered inputs using integer keys (version 1 code) 82

3.9 Run time on random real inputs (version 1 code) . 83



V








3.10 Run time on ordered real inputs (version 1 code) . 84

3.11 The number of key comparisons on random inputs (version 1 code) 87

3.12 The number of key comparisons on ordered inputs (version 1 code) 88

3.13 Height of the trees on random inputs (version 1 code) . 89 3.14 Height of the trees on ordered inputs (version I code) . 89 3.15 The number of rotations on random inputs (version 1 code) 90 3.16 The number of rotations on ordered inputs (version I code) 91 3.17 Run time on random inputs using integer keys (version I code) 93 3.18 Run time on ordered inputs using integer keys (version I code) 94 3.19 Run time on random real inputs (version 1 code) . 95

3.20 Run time on ordered real inputs (version 1 code) . 96

3.21 The number of key comparisons on random inputs (version 2 code) 99

3.22 The number of key comparisons on ordered inputs (version 2 code) 100

3.23 Run time on random real inputs (version 2 code) . 101

3.24 Run time on ordered real inputs (version 2 code) . 103

4.1 The number of key comparisons . . 117

4.2 Num ber of levels . . . 118

4.3 Run tim e . . . . 119

4.4 The number of key comparisons . . 124

4.5 Height/level of the structures . . 126

4.6 Run time using integer keys . . 127

4.7 The number of key comparisons . . 130



VI








4.8 Height/level of the structures . . 131

4.9 Run time using integer keys . . 132

A.1 Abbreviations used in tables . . 137














































VII

















LIST OF FIGURES




2.1 Cell joining...................................... 5

2.2 I-layer river routing. .. .. .. .. ... .... ... ... ... .....8

2.3 Round robin and greedy layer assignments. .. .. .. ... ... ....11

2.4 Minimizing the number of tracks or layers. .. .. ... ... ... ..14

2.5 Constraint graph representation. .. .. .. ... ... ... ... ....18

2.6 Merge in constraint graph. .. .. .. .... ... ... ... ... ..19

2.7 Heuristic 1. .. .. .. .. ... ... .... ... ... ... ... ....20

2.8 Heuristic 2. .. .. .. .. ... ... ... .... ... ... ... ....22

2.9 Heuristic 3. .. .. .. .. ... ... ... .... ... ... ... ....24

3.1 LL and RL rotations. .. .. .. ... ... ... .... ... ... ..35

3.2 A tree in WVB(1/4) that is not L-balanced. .. .. ... ... .. ....42

3.3 !-balanced tree that is not a COST .. .. .. .. .... ... ... ..43

3.4 Lb rotation for insertion .. .. .. .. .... ... ... ... ... ..45

3.5 Substep (i) of insertion LR rotation .. .. .. .. .... ... ... ..46

3.6 Case Lb, for LR(ii) rotation .. .. .. ... .... ... ... ... ..49

3.7 Case LR for LR(ii) rotation .. .. .. ... ... .... ... ... ..51








3.8 LL rotation for deletion. .. .. .. .. ... ... ... .... ... ..58

3.9 LR rotation for deletion. .. .. ... ... ... ... ... ... ...60

3.10 Restructuring procedure. .. .. .. ... .... ... ... ... ....65

3.11 Simple restructuring procedure for insertion .. .. .. ... ... ....68

3.12 Simple restructuring procedure for deletion .. .. .. .. ... ... ..68

3.13 Simple restructuring procedure without a #i value .. .. .. ... ....70

3.14 Run time on real inputs (version 1 code) .. .. .. ... ... ... ..85

3.15 Run time on random real inputs (version 1 code) .. .. .. ... ...92

3.16 Run time on ordered real inputs (version 1 code) .. .. .. .... ..97

3.17 Run time on random real inputs (version 2 code). .. .. .. ... ..102

3.18 Run time on ordered real inputs (version 2 code). .. .. ... .. ..102

4.1 Example min-WBLTs .. .. .. .. ... ... ... ... ... ... ..108

4.2 min-WBLT Insert. .. .. .. .. .... ... ... ... ... ... ..109

4.3 min-WILT Delete-mmn.. .. .. .... ... ... ... ... ... ..110

4.4 Skip Lists. .. .. .. ... .... ... ... ... ... ... ... ...11

4.5 Modified Skip Lists .. .. .. .. ... .... ... ... ... ... ..113

4.6 MSL Search. .. .. .. ... ... ... ... ... ... .... ....113

4.7 MSL Insert .. .. .. ... ... ... .... ... ... ... ... ..114

4.8 MSL Delete. .. .. .. ... ... ... .... ... ... ... ....114

4.9 Run time. .. .. ... ... ... ... ... ... ... .... .....120

4.10 TMSL Insert. .. .. ... ... ... ... ... ... .... ... ..121

4.11 TMSL Delete-min.. .. .. ... ... .... ... ... ... ... ..121




ix








4.12 TMSL Delete-max ....... ............................ 122

4.13 Run time on randomly ....... .......................... 128

4.14 Run time on random2 ...... .......................... 129

4.15 Run time on randomly ...... .......................... 131

4.16 Run time on random2 ...... .......................... 133











































x
















Abstract of Dissertation Presented to the Graduate School
of the University of Florida in Partial Fulfillment of the Requirements for the Doctor of Philosophy EFFICIENT ALGORITHMS AND DATA STRUCTURES FOR VLSI CAD By


Seonghun Cho


May 1996



Chairman: Dr. Sartaj Sahni
Major Department: Computer and Information Science and Engineering


In this dissertation, we develop efficient algorithms and data structures for problems that arise in electronic computer aided design (ECAD).

We consider the problem of joining a row of compacted cells so as to minimize the area occupied by the cells and the interconnects. The cell joining process includes cell stretching and river routing. We propose several heuristics to join a row of cells in such a way that area is minimized. The proposed heuristics are compared experimentally with the previously proposed heuristic.

We develop a new class of weight balanced binary search trees called f-balanced binary search trees (0i-BBSTs). 3-BBSTs are designed to have reduced internal path


xi








length. As a result, they are expected to exhibit good search time characteristics. Individual search, insert, and delete operations in an n node 13-BBST take O(log n) time for 0 < j3 < V2- 1. Experimental results comparing the performance of #I3BSTs, W13(a) trees, AVL-trees, red/black trees, treaps, deterministic skip lists and skip lists are presented. Two simplified versions of /3-BBSTs are also developed.

We propose the weight biased leftist tree as an alternative to traditional leftist trees for the representation of mergeable priority queues. A modified version of skip lists that uses fixed size nodes is also proposed. Experimental results show our modified skip list structure is faster than the original skip list structure for the representation of dictionaries. Experimental results comparing weight biased leftist trees and competing priority queue structures as well as experimentat results for double ended priority queues are presented.



























X11















CHAPTER 1
INTRODUCTION


1.1 Background


In VLSI layout, we are concerned with transforming a circuit from its logical design to a physical implementation. The layout problem for VLSI circuits is generally decomposed into smaller problems such as partitioning, floorplanning, placement, routing and compaction.

The partitioning process decomposes a large circuit/module into a collection of smaller sub-ci rcui ts/ modules. In floorplanning, logical components of a circuit are assigned relative positions on a chip. The physical realization of each component (i.e., its area and aspect ratio) is also selected. The objectives of floorplanning include overall area minimization, minimization of power consumption. etc. The precise locations for the components of a design are then determined during the placement process to optimize the area and the timing. After the components are placed, the pins are connected during the routing process. During the process of compaction, the components and interconnections are moved so as to further optimize the layout in terms of area and delay.

The routing process is usually divided into three smaller sub-problems of global routing, detailed routing, and specialized routing. Global routing decomposes the




1




2


complex routing problem into small and manageable sub-problems. It assigns each net to a set of routing regions such as channel, switcbbox, and planar routing to minimize a combination of criteria such as area, circuit delay, etc. Steiner trees and spanning trees are the commonly used approaches for net connection in global routing. Specialized routing is used to connect power-ground nets or clock nets. Detailed routing has two types of routing, general and restricted and there are three types of detailed restricted routing, channel, switchbox, and planar.

In channel routing, all terminals of nets are located in two parallel rows across a routing region called channel. In switchbox routing, terminals of nets are located on the four sides of the routing region. Planar routing is a problem in which the interconnection topology of the nets is planar. That is, all connections can be realized on a single layer. Vias allow wires to change layers but the presence of vias reduces reliability and performance of a circuit. Single layer routing is not always possible.

River routing is a special case of planar routing in which all nets have exactly two terminals, one on each side of the channel, and the net sequence on each side of the channel is the same. River routing is used in PCB routing, particularly in dataflow architectures with multi-bit buses connecting a series of logic blocks, and symbolic IC design systems.

1.2 Dissertation Outline

This dissertation is divided into four chapters. In Chapter 2, we consider the problem of joining a row of compacted cells so as to minimize the area occupied by the cells and the interconnects. The cell joining process includes cell stretching and




3


river routing. We propose several heuristics to join a row of cells in such a way that area is minimized. The proposed heuristics are compared experimentally with the previously proposed one.

VLSI Physical Design Automation is essentially the study of algorithms and data structures related to the physical design process. Specific data structures can be used to improve the performance of algorithms. For example, maze routing algorithms and line-probe algorithms in global routing [30] use search structures, retiring algorithms [6, 291 use priority queue structures, and layout compaction algorithms [13] use radix priority search trees.

Chapters 3 and 4 are about search and priority queue structures. As specific data structures could be used to produce better performance of algorithms in VLSI design automation, new search and priority queue structures are proposed and thoroughly compared with other data structures.

Finally, in the last chapter, we present conclusions of this work.














CHAPTER
MINIMUM AREA JOINING OF COMPACTED CELLS


2.1 Introduction

When designing circuits with compacted symbolic sticks basic cells, the circuit is realized by a collection of compacted cells that tile a two-dimensional area. The intercell interconnects are such that each interconnect connects two terminals that are on adjacent boundaries of neighboring cells. So, for example, if cells A and B (Figure 2.1(a)) are neighboring cells of the circuit, then the right boundary of A is adjacent to the left boundary of B. The number of terminals on each of these boundaries will be the same and the Z'th terminal (from the bottom) on the right boundary of A is to be connected to the i'th terminal (from the bottom) on the left boundary of B.

Since the cells are available in compacted form, it is not possible to reduce the distance between any pair of terminals on any side of a cell. However, this distance can be increased by stretching the cell. In the example of Figure 2. 1 (a), we can stretch either cell vertically by defining a horizontal cut line at any position and pulling the two cell pieces apart by any desired amount (the cell can also be stretched horizontally by using a vertical cut line).






4




5





5 5 E- 5 35
4 15
A- 4 3 B 10
3 2
2


(a) Horizontal adjacent cells (b) Joining by stretching





12




(c) Joining by river routing (d) Combination cell joining

Figure 2. 1. Cell joining


The required interconnects between cells A and B of Figure 2.1(a) can be accornplished by stretching cells A and B so that the terminals of A and B line up as in Figure 2.1(b). The broken lines in Figure 2.1(a) indicate the cut lines used for stretching. The stretching enables us to join cells A and B using no routing tracks (by "join" we mean make the interconnects between cells A and B). This method of joining cells is also called pitch matching.

Another way to join cells A and B is to river route the interconnects as in Figure 2.1(c). This uses routing tracks in a channel between cells A and B but does riot increase cell height. The pitch matching and river routing approaches to cell joining have been studied in Boyer [51 and Weste [33]. Algorithms for single-layer




6



river routing can be found in several works [15, 19, 23. 24] and those for multilayer river routing can be found in Baratz [3]. Single-layer gridless river routing is studied in Tompa [32]. Two applications of river routing are hybrid circuit design and structured design (DSP).

Cell stretching (or pitch matching) increases the height of the layout while river routing increases its width. Both affect the layout area. The layout of Figure 2.1(b) has area 150. To compute the area of the layout of Figure 2.1(c), we assume tracks have unit separation. So, the layout width is 14 and height is 11. The layout has area 154. Cheng and Despain [S) have proposed using a combination of cell stretching and river routing so as to obtain layouts with smaller area than possible when only one of these joining methods is used. Figure 2.1(d) shows the result of joining cells A and B using both stretching and river routing. The area of this layout is 144. This is minimum for the instance of Figure 2.1(a).

Cheng and Despain [8] have proposed a heuristic for single layer joining of compacted cells. At each step of their heuristic either a row or column of compacted cells is joined. Following this, the row or column of joined cells is replaced by a composite cell that represents the result of joining. Notice that when a row (column) of cells is joined, cells may be stretched vertically (horizontally) and river routing is done in a vertical (horizontal) channel. To join a row of cells, Cheng and Despain [81 bound the maximum height to which a cell may be stretched. This bound is hmax + hb1v9/(4 h...)




7


where h,,,, is the height of the tallest compacted cell being joined and h,,,-, is the average height of the cells being joined.

Using this bound, cells are joined one-at-a-time using a penalty/reward scheme to determine if a pair of terminals is to be joined by stretching or by river routing.

Lim, Cheng, and Sahni [17] have considered the case when only two cells are to be joined. They develop fast polynomial time algorithms to obtain the minimum area join of two cells. In addition, they are able to obtain, in low order polynomial time, minimum area joins that rninirf ze the length of the longest wire or the total wire length. Lim [16] has proposed an 0(n(n/c)-') algorithm to find the minimum area join of c cells having a total of n terminals. This algorithm does an exhaustive search over all possible numbers of tracks in the c I routing channels between adjacent cells. A constraint graph is used to determine the minimum height layout for each assignment of number of tracks to routing channels. The time required per track assignment is O(n) and the worst case number of track assignments is 0((n/c)-1). The algorithm of Lim [16] is flawed as it handles channels with zero routing tracks by joining the adjacent cells using minimum height cell stretching and then considers the joined cells as one. This problem is easily fixed, however, by combining, in the constraint graph, pairs of vertices that represent corresponding terminals of the two cells (i.e., ZAli terminals of each cell) with zero routing tracks in between.

In this chapter, we consider the case when I > 1 routing layers are available to river route the inter cell connections. Note that while multiple layers do not affect layout area when cell stretching alone is used, a reduction in area is possible when




8




















(a) 1-layer (b) 2-layer

Figure 2.2. I-layer river routing


cell stretching is combined with river routing or when river routing alone is used. We assume that in each layer of each routing channel, the interconnects are to be accomplished using river routing. An alternative is to use HV routing when I = 2, HVH or VHV routing when I 3, and extensions of HVH and VHV routing for

1 > 3. However, for river routing instances, using routing layers in this way has no advantage over river routing in each layer (see Theorem 5, Section 2). When the number of layers available for river routing is increased, one may see a dramatic reduction in the number of routing tracks needed per layer. Figure 2.2 shows an instance that needs n tracks when routed in one layer but only one track/layer when routed in two layers.

We begin, in section 2, by stating the necessary and sufficient conditions for a river routing instance to be routable in I layers using at most t tracks per layer arid stating how to perform I-layer river routing when such a routing is possible. In




9



this section, we also show that IIV style routing has no advantage over river routing in each layer. In Section 2.3, we describe the constraint graph used to determine minimum height stretching of c cells. Heuristics for the minimum area joining of c cells are proposed in Section 2.4 and the results of experiments with these are provided in Section 2.5. Our conclusions appear in Section 2.6.



2.2 1-layer River Routing


Let (Ai, B,), 1 < I < mn be a set of terminal pairs such that the A,'s are on one side (say left or top) of a routing channel and the Bi's are on the other (right or bottom) side. Terminal Ai is to be connected to terminal B2, 1 < i < m. For this channel routing instance to be an instance of river routing, it must be the case that a, < a2 < ... < a,~ and b, < b2 < ... < bm where a, and bi, respectively, give the positions of terminals Ai and Bi, 1 K I < m. We may assume an underlying grid with each terminal being at a grid position. In the case of a horizontal (vertical) channel the ais and bis are grid column (row) numbers. Leiserson and Pinter [151 have obtained the following necessary and sufficient condition for a river routing instance to be routable in a single layer using at most t > 0 tracks. Theorem 1 [15] The river routing instance defined above is routable in a single layer using at most t > 0 tracks if and only if

(a) ai~t 6, > t (b) b~ ai > t

for every i < 771 t.




10


For the general case of I > 1 layers, we obtain the necessary and sufficient condition of Theorem 2.


Theorem 2 The river routing instance defined above is routable in I > 1 layers (each layer routing whole nets) using at most t > 0 tracks per layer if and only if

(a) ai+it bi > t (b) bj+lt a > t

for every i < m- lt.


Proof: First, we establish that (a) and (b) are necessary for I layer routing. Since the proofs for (a) and (b) are similar, we provide that for (a) only. Suppose that ai+lt bi < t for some i. Consider the It + 1 terminal pairs (Aj, Bi), i < j < i + It. When routing these on I layers, at least one layer has to be assigned > t + 1 terminal pairs. So, suppose that terminal pairs (A',B ), (A'2, B), ..., (A+ 1,B'+), ... are assigned to the same layer for river routin g. We may assume that a' < a'2 < ... < at+1 andb < b2 < < bt+1. Since a' < a1+< t and b' > b -b' ai+it-bi 1. 2 a+1l -- ait+t

From Theorem 1, it follows that the terminal pairs (A, B), 1 < j < t + 1, cannot be river routed on a single layer. Hence, (Aj, Bj), i < j < i + It, cannot be river routed on 1 layers. So, (Aj, Bj), 1 < J < m, cannot be river routed on I layers. As a result,

(a) is a necessary condition.

To show that (a) and (b) are sufficient conditions for routability, we present two algorithms (RoundRobin and Greedy) that assign the nets to layers in such a way that each layer is river routable when both (a) and (b) are satisfied. The correctness








procedure RoundRobin
{ Assign the m nets to I layers. }
begin
fori 1 to m do
assign net (Ai, B1) to layer (i mod 1) + 1;
end ;

procedure Greedy;
{ Assign the m nets to I layers. }
begin
fori:= 1 tomdo
assign net (Ai, B1) to layer q such that q is the smallest integer < I
for which the conditions of Theorem 1 are not violated on layer q
(if there is no such q, then fail)
end ;

Figure 2.3. Round robin and greedy layer assignments

of these algorithms is established in Theorems 3 and 4, respectively. 0



We later discovered that Baratz [3] has not only obtained the same condition but also proposed the same two algorithms for I-layer river routing. One assigns nets to layers in a round robin fashion and the other uses a greedy strategy. The corresponding procedures are given in Figure 2.3.


Theorem 3 The layer assignment produced by the RoundRobin procedure is river routable if

(a) ai+lt bi > t (b) bi+t ai > t for all i < m-It.




12



Proof: Let (A', B'), (A,, B'), ... = (Aj,Bj), (Aj+t, Bj+,), (Aj+2, Bj+21), ... be the nets assigned to layer (j mod 1) + 1, j < 1. So, a' = aj+(i-) and b' = bj+(i-I. Hence,
I
ai+t b = aj+(i+t_)l bj+(i-)l

aj+(i-1)I+ti bj+(i-1)l

> t (from (a))


Similarly, b, a > t. So, the layer assignment satisfies the conditions of Theorem 1 and is river routable using t tracks. 0


Theorem 4 If (a) ai+t bi t and (b) bi+tt ai t for all i < m lt, then procedure Greedy assigns nets to layers such that the assignment to each layer is routable using t tracks.


Proof: If procedure Greedy is able to assign each of the mn nets to a layer, then the layer assignments satisfy the conditions of Theorem 1 and so are routable using t tracks. Suppose the algorithm fails while trying to assign net (A,, B,) to a layer. At this time nets (Ai, Bi), 1 < i < r, have been assigned to layers so as to satisfy the conditions of Theorem 1 and the assignment of net (A,, Br) to each of these layers violates these conditions. Consider first those layers, L,, on which condition (a) is violated. For a layer s E La, suppose that the assigned nets are ..., (A-t, B-t),
..., (A1,Bj). Let (A, BJ) = (A,B,.). Since s E L,, we have a, bj_, < t. Now, if b,.- > bj'_t, then a. br-it < t which violates condition (a) of this theorem. So, b-t < b'>_. Since, b-., < b'. < ... < b6'2 < b'-_, I of the It 1 nets




13


(Ar -t+ I, Br..it+ ). ...., (Ar1, B,_1) have been assigned to layers E L,,. Consequently. the layers in L,, account for tiLal of these it 1 nets.

In a similar way, we can show that the remaining 1 ILail layers account for another W~ 1L.1) of these nets. This gives us a total of ti nets, whereas we had only ti 1. This contradiction implies that procedure Greedy cannot fail unless conditions

(a) and (b) are not satisfied. 11



Procedure RoundRobin is easily seen to have complexity of 0(m). A straightforward im-plementation of procedure Greedy will have complexity of 0(ml). However, by using priority search trees [18] the complexity can be reduced to O(m log 1). In practice, since I is quite small, it is unlikely that the priority search tree implementation will run faster than the straightforward implementation in which the 1 layers are checked in sequence. The actual routing for all I layers can be done in 0(mt) time using the computed layer assignment and the single layer routing algorithm of Leiserson and Pinter [15].

Using Theorem 2, we can develop a linear time algorithm to determine the minimum number of tracks needed to route an instance in I layers as well as to find the minimum number of layers needed for a t track routing. The algorithm for the former is given in Figure 2.4. This figure also shows the changes needed in case t is given and we wish to determine the minimum number of layers. The correctness of the algorithm follows from that of Theorem 2 and the fact that t (or 1) is increased only if the current t (1) is found to be infeasible. The complexity is 0(m) as neither




14


procedure MinimizeTracks; {or MinimizeLayers}
{ Determine the minimum number of tracks per layer (or minimum number
of layers) needed for multilayer river routing }
begin
t:0O; {or I := 1}

while (i < m It) do
if (ai+lt b, < t) or (bi+t ai < t)
thent:=t+ 1 {orl:=l+1}
else i :=i + 1;
end ;

Figure 2.4. Minimizing the number of tracks or layers

i nor t (1) can exceed m. So, neither clause of the if statement can be executed more than m 1 times.

Using the multilayer river routing results of Baratz [3], one can trivially extend all the results of Lim, Cheng and Sahni [17] to the case of multilayer joining of compacted cells. So, the multilayer minimum area join of two compacted cells with m nets can be obtained in O(m2) time. If we wish to minimize the maximum wire length while keeping area minimum, the asymptotic time complexity is still O(m2). The total wire length can be minimized while keeping area minimum in O(m2 log m) time.

In HV style routing, each routing layer is assigned a routing direction (either H or V). In an H (V) layer only horizontal (vertical) wire segments can be laid out. Horizontal segments on one layer connect to vertical segments, of the same net, on another layer by means of vias. In the case of river routing instances, one can see that there is no advantage to having more than two V-layers (i.e., two V-layers are sufficient to route all river routing instances).




15


Let RR(I, t) be the set of all river routing instances that can be routed in 1 layers, using t tracks per layer and using river routing in each layer. Let HV(I, t) be all river routing instances that can be routed using HV style routing, 1 layers, and t tracks per layer. Note that HV(I,t) includes instances routable with 0, 1, and 2 V-layers. Let HVV(1, t) be all river routing instances using HV style routing, 1 2 H-layers, and 2 V-layers. Theorems 5 and 6 below hold for both the knock-knee [25] and directional HV models. Theorem 7 holds only for the directional model. Theorem 5 HV(I,t) C RR(I,t) for every I > 1 and every t > 1. Proof: HV(l, t) RR(/, t) follows from a more general result obtained by Baratz [3]. Baratz [3] has shown that, for river routing instances, there is no advantage to using any routing scheme that wires a net on more than one layer. Since it is easy to construct river routing instances X such that X E RR(1, t) and X HV(l,t), it follows that HV(I, t) C RR(I, t).

We provide a simpler proof of HV(I, t) g RR(l, t). This proof will also establish our next result. We shall show that if X is a river routing instance such that X RR(/,t), then X 0 HV(/,t). Hence, HV(/,t) C RR(/,t).

Suppose that X RR(l, t). From Theorem 2, it follows that ai+t bi < t or b,+tj aj < t for some i. Suppose that a,+Ii bi < t (the proof is similar when bi+tt ai < t). So, ai+lt < bi t. Since X is a river routing instance, at least nets i + t,. ,i + It intersect a vertical cut line drawn at ai+lt. Hence, the density of X at aj+it is > i + It (i + t) + 1 = (I 1)t + 1. When HV style routing is used with




16


1, 1 > 1, layers, at most I 1 layers are available for horizontal routes. With t tracks per layer, densities of at most (I 1)t can be accommodated. So, X HV(/, t). 0 Theorem 6 HVV(l,t) C RR(l- 1,t) for everyI > 2 and everyt > 1. Proof: As in Theorem 5, suppose that X 0 RR(I- 1, t). Let i be such that ai+(-1)t bi < t. The net density at a1+(1-1)t is > (1 2)t + 1. In HVV routing, two layers are V-layers. So only I 2 layers are available for horizontal segments. This is not enough as the horizontal segment density is > (1 2)t + 1 at ai+(-1)t. Hence, X HVV(!, t).

One may easily construct river routing instances that are in RR(l 1, t) but not in HVV(l,t). 0


Theorem 7 RR(2, t) r HV(l, t) HVV(1, t) for every I > 1 and every t > 1. Proof: Consider the RR instance (a,,bl) = (1,2) and (a2,b2) = (2,3). This is in RR(2,t) for every t > 1 but is not in HV(/, t) HVV(/, t) for any 1. El



As remarked earlier, Theorem 7 holds only for the directional model. For the knock-knee model, one can show that RR(/, t) C HV(/ + 1, t) for every I > 1 and every t > 1.


2.3 Constraint Graph Representation

Lim [16] has proposed the use of a constraint graph to determine the terminal positions in a row of compacted cells. This is for the case when the number of tracks




17


in each routing channel is given and we wish to minimize the layout height. In the constraint graph, each cell is represented by a directed chain of vertices. Each cell terminal is represented by a vertex. The exception is when a compacted cell has terminals at the same y-position on both sides of the cell. In this case, the two terminals at the same y-position are represented by a single vertex. The vertex chain is linked in the direction of increasing y-position. The chain edges are labeled by the minimum allowable terminal separation. In addition, the constraint graph contains a source vertex that represents the bottom of the layout and a sink vertex that represents the layout top. The source vertex connects to the bottom of each chain and the top of each chain is connected to the sink vertex.

Figure 2.5(b) shows the chains (solid edges) for the four cell row of Figure 2.5(a). To complete the constraint graph, directed edges are added to introduce the channel routing constraints of Theorem 2. These are represented by the broken edges of Figure 2.5(b). Figure 2.5(b) is for the two layer case.

Lim [16] has shown that the constraint graph is acyclic provided the number of tracks in each routing channel is > 0. He has proposed handling channels with zero tracks by finding first the minimum area joining of the adjacent cells (only cell stretching is permitted now) and then combining these two cells into one. Le., the two cells are replaced by their minimum area join. This strategy can be shown to result in non-optimality of the algorithm proposed in Lim [16). To preserve optimality, it is necessary to merge the vertices that represent terminals that are the endpoints of




18



sink


2

C 9 M



C M f 2
b g b
L f

a d h k 3 2

a d h k

(a) Cell joining with 2 layers



source

(b) Constraint graph representation Figure 2.5. Constraint graph representation


nets that are to be routed using no tracks as in Figure 2.6. The resultant constraint graph is also acyclic.

It is easy to see that the number of vertices and edges in the constraint graph is O(n) where n is the total number of terminals. Furthermore, the graph can be constructed in O(n) time given the number of routing layers and the number of tracks in each channel. The constraint graph described by us is identical to that of Lim [16] except in the way channels with zero tracks are handled and in that our graph is defined for 1 > 1 routing layers while that of Lim [161 is only for I = 1.




19



sink


2

C 9 M


f 2

b


3 di 2
i V
a h k





source

Figure 2.6. Merge in constraint graph


The length of the longest path from the source vertex of the constraint graph to each of the remaining vertices can be computed in O(n) time by doing this in topological order [14, Section 6.5]. It is easy to see that if each terminal is placed at a vertical position given by the longest path length from the source, then all nets can be routed in the given number of tracks (as the conditions of Theorem 2 are satisfied in each routing channel). Furthermore, Lim [16] has shown that such a positioning of terminals results in a stretched layout of minimum height for the given channel widths. As a result, when channel widths are known, cells can be stretched to minimize area in O(n) time. The channel widths that result in minimum area can be determined in 0(n(n/c)c-1) time where c is the number of cells by trying out all




20


procedure Ileuristici
begin
for i: 1 to c -1 do
begin
determine the minimum area join of each pair of adjacent cells;
select the pair that has minimum area and replace it with its
minimum area join;
end;
end ;

Figure 2.7. Heuristic 1

possible channel widths [16]. Since this is feasible only for small c, we propose several heuristics in the next section.



2.4 Heuristics to Minimize Area


We formulate three greedy heuristics to obtain the minimum area join of a row of c compacted cells that have a total of n terminals.


2.4.1 Heuristic 1

The heuristic is described in Figure 2.7.

In each iteration of the for loop we examine every pair of adjacent cells. For each pair, the minimum area join is found using the algorithm of Lim, Cheng and Sahni [17] extended to the multilayer case as discussed in Section 2.2. The pair which has the minimum area join is replaced by a single cell that represents this join. So following each iteration of the for loop the number of cells decreases by one. When the for loop terminates, we are left with a Single cell that represents the join of all c cells. The time needed to determine the minimum area join of a pair of cells with ni nets between them is O(n?). The time to do this for all pairs of adjacent




21


cells is O(>ji n') 0 (n 2). So, the for ioop iteration with i = 1 takes 0(n2) time. On subsequent iterations, only the two pairs that include the cell introduced in the previous iteration need to have their minimum area join computed. Since each cell pair being considered includes at least one composite cell, the minimum area join is computed by considering the portion of the constraint graph that represents all the basic cells in the cell pair. Channel widths for channels within a composite cell are not changed while obtaining the minimum area join of the cell pair. However, as different channel widths for the channel between the two (composite) cells being joined are tried, the constraint graph is used to determine the minimum height of the combined cell. So, the time to combine two (composite) cells with ni terminals in the channel between them is O(nn1). Hence the time for the remaining c 2 iterations is O(n Xj7i n,) = 0(n'). The overall complexity of Heuristic 1 is therefore 0(n2). In case the terminals are uniformly distributed over the cells, ni = O(n/c) for all i. The time for the first iteration of the for loop is now O(n2/c) and that for each of the remaining iterations is 0(n2/c). The overall time is 0(n2).

2.4.2 Heuristic 2

In this heuristic, we begin by assigning each channel the number of tracks needed to route the channel with no cell stretching. This number can be determined in 0(n1) time for a channel with ni nets as described in Section 2.2. The time taken to do this for all c 1 channels is 0(n). The configuration obtained in this way is the maximum width layout. Starting from this configuration, we reduce the total number of tracks available across all c 1 channels by one on each iteration. For this, the




22


procedure Heuristic2
begin
for each channel determine the number of tracks, ti,
needed to route with no stretching, 1 < i < C;
t =E-1t
set up the constraint graph using ti tracks in channel i, 1 < i < c
compute layout area, A ;
for tracks :=t downto 1 do {reduce by 11
begin
for 1 i=1to c-i1 do
begin
reduce the number of tracks in channel i by 1
modify the constraint graph to reflect this ;
determine the length of the longest path in the graph
and from this the layout area, a1
end;
select 1 such that a1 = min{ a,
reduce the number of tracks in channel j' by I
A =min{f A, a,
end
end ;

Figure 2.8. Heuristic 2

effect of a one track reduction is computed for each channel. The minimum layout

height is determined by computing the length of the longest path in the constraint

graph of Section 2.3. The track reduction is done in the channel that results in the

smallest layout height (hence the minimum area for the given number of tracks). The

algorithm is stated more formally in Figure 2.8.

When the algorithm terminates, A is the area of the minimum area join found

by the heuristic. To reconstruct the layout, it is necessary to store the tracks per

channel each time A is updated in the statement A = min{f A, a, }.

For the time complexity, we see that the steps that precede the outer for loop

take O(n) time. Each iteration of the outer loop takes 0(nc) time. Hence this ioop




23


contributes a total of 0(net) to the time. Since t = O(n), the overall time complexity of Heuristic 2 is O(n 2 C).

2.4.3 Heuristic 3

Unlike Heuristic 2 which attempts to minimize the layout height for each value of t, the totaJ number of tracks, Heuristic 3 attempts to minimize the width (i.e., total number of tracks) for each choice of layout height. The heuristic begins with a layout height, ht, equal to the height of the tallest compacted cell. At each iteration, the next layout height to use is computed as described later. During each iteration, cells are combined in groups of at most k (k > 1 is a parameter to the heuristic). Each group of combined cells is replaced by its minimum area join subject to the constraint that the height of the join does not exceed ht. This joining of < k cells at a time continues until only one cell remains. Its area is computed and recorded. The minimum area obtained over all heights tried is then reported as the best. Heuristic

3 is given in Figure 2.9.

In our implementation of Heuristic 3, the minimum area join of k cells is found by considering the portion of the constraint graph for all the basic cells included in these k cells. So, for this purpose composite cells are not handled as single cells. Rather, as in Heuristic 1, the basic cells they are composed of are considered and channel widths previously assigned to the associated channels are not changed. Track assignment is done only for the k 1 channels between the k composite cells. We found this to give better results than when composite cells were regarded as atomic. For the case k = 2, the minimum area is determined by a binary search over the




24


procedure Heuristic3
begin
ht :=height of the tallest cell;
repeat I minimize width subject to height < ht}
repeat I do this by combining k cells at a time}
select k adjacent cells such that the minimum height cell is selected
and the height of the tallest selected cell is minimum
(if there are fewer than k cells, then select all of them)
obtain the minimum area layout for the selected cells under
the constraint that the layout height does not exceed ht
during the preceding step record the next value of ht
that is possible for a layout
until one cell remains;
compute the area of the remaining cell and record it
if it is less than the minimum area found so far;
if there is no next height then terminate;
hti: next height;
until false
end ;

Figure 2.9. Heuristic 3

number of tracks in the single channel. This takes O(n log n,) time where ni is the

number of nets in channel i. Thus the time needed for the inner repeat loop when

k = 2 is O(cn log n) (for uniform terminal distribution it is O(cn log(n/c)). During

the binary search, the heights corresponding to channel widths that require height

> ht are recorded. The minimum of these heights yields the next value of hi.

When k > 2, all track combinations for the k 1 channels are tried as in

Section 2.3. Again, each composite cell is broken up into its basic cells. As different

track combinations are tried, we record the minimum height > ht that results from

any track combination. This gives the next value of ht. The time for the inner

repeat loop is O((c/(k 1))n(n/k)k-1) (or O((c/(k 1))n(n/c)k1') when terminals

are uniformly distributed).




25


In all our experiments, the outer repeat loop was iterated fewer than (k 1)n times. To ensure that the number of iterations is O(kn), one may adopt the following scheme. When the number of iterations first reaches (k 1)n, compute a set of at most n new heights by beginning with the current constraint graph. This uses the current assignment of number of tracks in each channel. Heuristic 2 is next used to reduce the total number of available tracks by one and determine the height needed to complete the routing with the reduced number of tracks. This process gives us at most n new heights h, < h2 < ... < hp,. Heuristic 3 is now resumed with h, as the next height. Only two iterations are performed. Then Heuristic 3 is resumed with max{ h2, ht I as the next height. Again two iterations of the outer repeat loop are done. Next the heuristic is resumed with max I h3, ht } as the next height. This continues until we have gone through p resumptions of the heuristic. With this scheme to limit the number of iterations, the complexity of Heuristic 3 becomes O(cn 2 log n) when k = 2 and O((c/(k 1))kn 2(n/k)k1') = O(cn2(n/k)k1') when k > 2. For the case when the n terminals are uniformly distributed over the c cells, the complexity is O(cn'log(n/c)) when k = 2 and O((c/(k 1))kn(n/c)-1) = O(cn(n/c)k1') when k > 2. One may verify that since Heuristic 3 tries the maximum useful height (i.e., the height needed when no routing tracks are available), it generates optimal solutions when k = c.



2.5 Experimental Results

We programmed our three heuristics as well as the heuristic Fang [8) in C and ran tests on a single KSR processor. Optimal solutions for instances with up to nine




26


cells were obtained using the corrected version of the exhaustive search algorithm of Lim [161. Our test set consisted of instances that had a number of cells, c, equal to one of the numbers in the set 13, ..., 9, 10, 20, 50, 1001. For each value of c, there were twenty instances and the results were averaged over these instances. An instance with c cells had c I routing channels. The number, t, of terminals on either side of each routing channel was equal to c for 3 < c < 9 and was 10 for the other values of c. In addition, when c = 100, we also had instances with 20 terminals on either side. In our experiments, we considered only single layer and two layer routing.

Table 2.1 gives the average percentage by which the area of the single layer solutions generated by each of the heuristics exceeded the area of the single layer optimal solution. As is evident, each of the heuristics proposed in this chapter gave noticeably better solutions than did Fang. This table is only for the cases 3 < c < 9 as for c > 9 the optimal algorithm of Lim [16] required too much time to complete. When k > c, Heuristic 3 is guaranteed to generate an optimal solution. So, we did not run these cases.

In table 2.2, we have used the single layer solution produced by Fang as the benchmark against which the solutions obtained by our three heuristics are compared. This table gives the average percentage by which the area of the solutions produced by our heuristics is less than that of the solutions produced by Fang. Our solutions have area 9 to 18% less.

Table 2.3 compares the computing time requirements of the various algorithms for the case of one layer. The optimal algorithm is useful only for small values of c




27








Table 2.1. Error rate (%) over optimal, 1 = 1

cells___t Fang_ Hersi Heuisic He 31k =41
[~ ~ cel 21 Fag Huitc Huitc4__Huitc__3 59 0.5 0.2 0 *
4 4 10.0 0.9 0.1 0 0 *
5 5 11.0 3.7 0.3 0.3 0.1 0.1
6 6 12.7 2.4 0.3 0.2 0.2 0.0
7 7 16.1 3.5 0.3 0.1 0.1 0.1
8 8 17.9 2.7 0.4 0.3 0.3 0.1
9 9 18.8 3.5 0. 3 0.4 10.3 0.

t = number of terminals on each side of each routing channel
:k> c












Table 2.2. Improvement (%) over Fang, I = 1

cell I t 21 Heriti Heriti __rs
[~ ~~ cel 2f~ Hersi J Hersi 3__eur=ti 4
10 10 9.4 14.0 14.0 14.2 14.3
20 10 10.5 15.4 15.4 15.6 15.6
50 10 10.6 16.1 16.1 16.2
100 10 9.1 16.5 16.5 16.6
100 20 if 9.3 18.3 18.0

excessive run time




28



Table 2.3. Time taken, I

cells I t Fang Heuristic Heuristic Heuristic3 Optimal
1 2 k 2 k 3 k 4
3 3 0.0 0.01 0.02 0.02 0.01
4 4 0.0 0.02 0.02 0.04 0.09 0.05
5 5 0.0 0.03 0.04 0.09 0.36 1.2 0.77
6 6 0.0 0.03 0.08 0.21 0.66 3.7 13
7 7 0.0 0.04 0.14 0.50 3.4 27 278
8 8 0.0 0.07 0.24 0.84 4.4 34 1.8t
9 9 0.0 0.09 0.42 1.8 16 72.3 47t
10 10 0.0 0.14 0.78 3.1 19 334
20 10 0.01 0.32 5.8 16 117 1196
50 10 0.01 1.4 93 127 896
100 10 0.03 4.4. 782 590 4873
100 20 0.06 11 3022 4087

Times are in seconds.
t : Times are in hours.


(say up to 7). While Fang is significantly faster than the heuristics proposed here, the quality of the solutions generated by our heuristics is superior.

Table 2.4 is the analog of table 2.1 for the case of two layers. Again, our heuristics performed considerably better than did Fang. Table 2.5 gives the improvement in area due to increasing the number of routing layers from one to two. This is influenced somewhat by the width of cells which in our case ranged from 5 to 30 times the track separation. With narrower cells, the impact of the second layer would have been greater and with wider cells, it would have been less. Also, the impact of the second layer is more when more routing tracks are needed. For the smaller instances of table 2.4, for example, the optimal solutions with 1 = 2 required, on average, only

1.8% less area than when I = 1.




29









Table 2.4. Error rate (%) over optimal, 1 = 2

cells I Fang Heuristic Heuristic Heuristic3

3 3 6.3 0.5 0 0
4 4 9.8 1.1 0.1 0 0
5 5 11.0 2.7 0.1 0.2 0.1 0.0
6 6 13.0 1.4 0.1 0.0 0.0 0.1
7 7 15.9 2.3 0.1 0 0 0
8 8 16.8 1.7 0.2 0 0.1 0.0
9 9 18.4 2.2 0.1 0.1 0.1 0.0














Table 2.5. Improvement (%) over I = 1 cases

cells Fang Heuristic Heuristic Heuristic3
1 2 V=-k=[=
~~I I_________ k = 2k=_3 1 k = 4
10 10 4.1 4.7 3.3 3.3 3.0 3.0
20 10 4.1 5.8 3.4 3.4 3.3 3.2
50 10 4.6 5.2 3.4 3.2 3.2
100 10 4.7 6.3 3.5 3.4 3.4
100 20 7.1 10.4 4.8 5.2 -




30



Table 2.6. Improvement (%) over Fang, I = 2


[cells t Heuristic I Heuristic Heurisfic3
1 2 k = 2 1 k = 3 k =- 4
10 10 10.0 13.3 13.2 13.2 13.3
20 10 12.2 14.8 14.8 14.9 14.8
50 10 11.2 15.1 14.9 15.0
100 10 10.7 15.4 15.4 15.4
100 20 12.5 16.3 16.3 :__J



Table 2.6 is the analog of table 2.2 for the case of two layers. The results are similar to those in table 2.2. 'Fable 2.7 gives the average computing times for the two layer instances. These are less than for the one layer case as the constraint graph has fewer edges.

For large c, we recommend the use of heuristic 2 or 3 (with k = 2) and for small c we recommend using heuristic 3 (with k = 3 or 4).


2.6 Conclusion

vVe have considered the problem of joining a row of compacted cells and developed heuristics to stretch cells and river-route the nets so that the layout area is minimized. Our proposed heuristic was compared, experimentally, with Fang [81 and found to produce layouts with less area. However, Fang is faster. We recommend the use of our Heuristic 3 with k = 3 or 4 in practice.




31

















Table 2.7. Time taken, 1 = 2

~cellsj IJt Fanj Heuristic ersi1k=leuristic3 =_ Optimal

3 3 0.0 0.01 0.02 0.02 *0.0
4 4 0.0 0.02 0.02 0.03 0.06 0.04
5 5 0.0 0.02 0.02 0.07 0.25 0.75 0.60
6 6 0.0 0.02 0.04 0.13 0.38 2.0 11
7 7 0.0 0.04 0.06 0.29 1.9 14 223
8 8 0.0 0.06 0.12 0.53 2.5 19 15
9 9 0.0 0.06 0.20 0.96 7.5 34 39
10 10 0.01 0.09 0.32 1.7 9.9 15320 10 0.01 0.24 2.7 9.9 6.5 65150 10 0.01 1.1 42 85 586
100 10 0.02 3.7 357 472 3069 -1 00 20 0,05 7.9 1377 3166 -Times are in seconds.
t:Times are in hours.















CHAPTER 3
A NEW WEIGHT BALANCED BINARY SEARCH TREE


3.1 Introduction


A dictionary is a set of elements on which the operations of search, insert, and delete are performed. Many data structures have been proposed for the efficient representation of a dictionary [141. These include direct addressing schemes such as hash tables and comparison schemes such as binary search trees, AVL-trees, red/black trees [121, trees of bounded balance [21], treaps [1], deterministic skip lists [20], and skip lists [26]. Of these schemes, AVL-trees, red/black trees, and trees of bounded balance (WB(a)) are balanced binary search trees. When representing a dictionary with n elements, using one of these schemes, the corresponding binary search tree has height O(log n) and individual search, insert, and delete operations take O(log n) time. When (unbalanced) binary search trees, treaps, or skip lists are used, each operation has an expected complexity of O(log n) but the worst case complexity is O(n). When hash tables are used, the expected complexity is 0(1) per operation. However, the worst case complexity is 0(n). So, in applications where a worst case complexity guarantee is critical, one of the balanced binary search tree schemes is to be performed.






:32




33


In this chapter, we develop a new balanced binary search tree called fl-BBST (/3balanced binary search tree). Like WB(a) trees, this achieves balancing by controlling the relative number of nodes in each subtree. However, unlike WB(a) trees, during insert and delete operations, rotations are performed along the search path whenever they reduce the internal path length of the tree (rather than only when a subtree is out of balance). As a result, the constructed trees are expected to have a smaller internal path length than the corresponding WB(a) tree. Since the average search time is closely related to the internal path length, the time need to search in a /3-BBST is expected to be less than that in a WB(a) tree.

In Section 3.2, we define the total search cost of a binary search tree and show that the rebalancing rotations performed in AVL and red/black trees might increase this metric. We also show that while similar rotations in WB(a) trees do not increase this metric, insert and delete operations in WB(a) trees do not avail of all opportunities to reduce the metric. In Section 3.3, we define O-BBSTs and show their relationship to WB(a) trees. Search, insert, and delete algorithms for /3-BBSTs are developed in Section 3.4. A simplified version of 83-BBSTs is developed in Section 3.5. Search, insert and delete operations for this version also take O(log n) time each. An even simpler version of f3-BBSTs is developed in Section 3.6. For this version, we show that the average cost of an insert and search operation is O(log n) provided no deletes are performed.

An experimental evaluation of f3-B BSTs and competing schemes for dictionaries (AVL, red/black, skip lists, etc.) was done and the results of this are presented in




34



Section 3.7. This section also compares the relative performance of P-BBSTs and the two simplified versions of Sections 3.5 and 3.6.

3.2 Balanced Trees and Rotations

Following an insert or delete operation in a balanced binary search tree (e.g., AVL, red/black, WB(a), etc.), it may be necessary to perform rotations to restore balance. The rotations are classified as LL, RR, LR, and Rb [14]. Lb and RR rotations as well as bR and Rb rotations are symmetric. While the conditions under which the rotations are performed vary with the class of balanced tree considered, the node movement patterns are the same. Figure 3.1 shows the transformation performed by an bb and an bR rotation. In this figure, nodes whose subtrees have changed as a result of the rotation are designated by a prime. So, p' is the original node p however its subtrees are different.

bet h(x) be the height of the subtree with root x. bet s(x) be the number of nodes in this subtree. When searching for an element x, x is compared with one element at each of I(x) levels, where I(x) is the level at which x is present (the root is at level 1). So, one measure of the "goodness" of the binary search tree, T, for search operations (assuming each element is searched for with equal probability) is its total search cost defined as:



C(T) =Z1(x).
xET




35




9P
P

P d
qI 9P I

q
c a A d
b c

a b

(a) LL rotation


9P
qI

p d
I
P 9P
a q A. X
a b c d

b c

(b) LR rotation Figure 3.1. LL and RL rotations


Notice that C(T) = I(T) + n where I(T) is the internal path length of T and n is the number of elements/nodes in T. The cost of unsuccessful searches is equal to the external path length E(T). Since E(T) = I(T) + 2n, minimizing C(T) also minimizes E(T).

Total search cost is important as this is the dominant operation in a dictionary (note that insert can be modeled as an unsuccessful search followed by the insertion of a node at the point where the search terminated and deletion can be modeled by




36


a successful search followed by a physical deletion; both operations are then followed by a rebalancing/restructuring step).

Observe that in an actual implementation of the search operation in programming languages such as C++, C, and Pascal, the search for an x at level I(x) will involve upto two comparisons at levels 1, 2,..., 1(x). If the code first checks x = ej where ej is the element at level i to be compared and then x < e to decide whether to move to the left or right subtree, then the number of element comparisons is exactly 21(x) 1. In this case, the total number of element comparisons is



NC(T) = 2 E 1(x) n = 2C(T) n
XET


and minimizing C(T) also minimizes NC(T). If the code first checks x < ej and then x = ei (or > ei), the number of element comparisons done to find x is 1(x) + r(x) + 1 where r(x) is the number of right branches on the path from the root to x. The total number of comparisons is bounded by 2C(T). For simplicity, we use C(T) to motivate our data structure.

In an AVL tree, when an LL rotation is performed, h(q) = h(c)+1 = h(d)+1 (see Figure 3.1(a)). At this time, the balance factor at gp is h(p) h(d) = 2. The rotation restores height balance which is necessary to guarantee O(log n) search, insert, delete operations in an n node AVL tree. The rotation may, however, increase the total search cost. To see this, notice that an LL rotation affects the level numbers of only those nodes that are in the subtree with root gp prior to the rotation. We see that l(q') = l(q) 1, (p') = l(p) 1, 1(gp') = l(gp) + 1, the total search cost of the subtree




37


with root a is decreased by s(a) as a result of the rotation, etc. Hence, the increase in C(T) due to the rotation is:



1(p') 1(p) + l(q') I(q) + l(gp') l(gp) s(a) s(b) + s(d)



= -1 1 + 1 s(q) + 1 + s(d) = s(d) s(q). A similar analysis shows that an LR rotation increases C(T) by s(d) s(q).

If the LL rotation was triggered by an insertion, s(q) is at least one more than the minimum number of nodes in an AVL tree of height t = h(q)- 1. So, s(q) > 0t+2/v/5 where 4 = (1 + V/5)/2. The maximum value for s(d) is 2' 1. So, an LL rotation has the potential of increasing total search cost by as much as



2t 1 Ot+2/,/-5 -- 2t 1 1.62t+2/2.24.



This is negative for t < 2 and positive for t > 2. When t = 10, for example, an LL rotation may increase total search cost by as much as 877. As t gets larger, the potential increase in search cost gets much greater. This analysis is easily extended to the remaining rotations and also to red/black trees. Definition (WB(a) [21]) The balance, B(p), of a node p in a binary tree is the ratio (s(l) + 1)/(s(p) + 1) where I is the left child of p. For a E [0, 1/2], a binary tree T is in WB(a) iff a < B(p) < 1 a for every node p in T. By definition, the empty tree is in WB(a) for all a.




38


Lemma 1 (1) The maximum height, hmaz(n), of an n node tree in WB(a) is ~ log.s (n + 1) [21]

(2) Inserts and deletes can be performed in an n node tree in WB(a) in O(log n) time for 2/11 < a < 1 V2-/2 [4].

(3) Each search operation in an n node tree in WB(a) takes O(log n) time [21].


In the case of weight balanced trees WB(a), an LL rotation is performed when B(gp) = 1 a and B(p) > a/(1 a) (see Figure 3.1(a)) [21]. So,


s(p)+1 s(p) + 1
1 -as
s(gp) + 1 s(p) + s(d) + 2


or
a 2a- 1
s(d) s(p) a
1 a 1 a and

a s(q) + 1
-- < B(p) = o1-a B _s(p)+1

or
a 20 1
s(q) > s(p)-2- + 2a1-a 1-a

So, LL rotations (and also RR) do not increase the search cost. For LR rotations [211, B(gp) ,~ 1 a and B(p) < a/(1 a). So, s(d) ~ s(p)_ + and with respect to Figure 3.1(b),


a s(p) s(q)
> B(p)1 -a s(p) +1




39


or

s(q) >s(p) 1 2a a


For a < 1/3, s(q) > s(d) and LR (RL) rotations do not increase search cost. Thus, in the case of WB(a) trees, the rebalancing rotations do not increase search cost. This statement remains true if the conditions for LL and LR rotation are changed to those in Blum and Mehlhorn [4].

While rotations do not increase the search cost of WB(a) trees, these trees miss performing some rotations that would reduce search cost. For example, it is possible to have a < B(gp) < 1 a, B(p) > -a, and s(q) > s(d). Since B(gp) isn't high enough, an LL rotation isn't performed. Yet, performing such a rotation would reduce search cost.

3.3 U-BBSTs

Definition A cost optimized search tree (COST) is a binary search tree whose search cost cannot be reduced by performing a single LL, RR, LR, or RL rotation. Theorem 8 If T is a COST with n nodes, its height is at most log,(v/5(n + 1)) 2. Proof Let Nh be the minimum number of nodes in a COST of height h. Clearly, No = 0 and N1 = 1. Consider a COST Q of height h > 2 having the minimum number of nodes Nh. Q has one subtree R whose height is h 1 and another, S, whose height is < h 1. R must be a minimal COST of height h 1 and so has Nh-1 nodes. R, in return, must have one subtree, U, of height h 2 and another, V, of height < h 2. Both U and V are COSTs as R is a COST. Since R is a minimal




40


COST, U is a minimal COST of height h 2 and so has Nh-2 nodes. Since Q is a COST, IS max{IjUI, IVI}. We may assume that Nh is a nondecreasing function of h. So, ISI Nh-2. Since Q is a minimal COST of height h, ISI = Nh-2. So, Nh = Nh-1 + Nh-2 + 1, h > 2



N0 =0, N = 1.


This recurrence is the same as that for the minimum number of nodes in an AVL tree of height h. So, Nh = Fh+2 -1 where F is the i'th Fibbonacci number. Consequently, Nh a qh+2/v/5 1 and h < log,(v5(n + 1)) 2. O




Corollary 1 The maximum height of a COST with n nodes is the same as that of an AVL tree with this many nodes.


Definition Let a and b be the root of two binary trees. a and b are 3-balanced,

0

(a) fl(s(a) 1) < s(b)

(b) fl(s(b) 1) < s(a)


A binary tree T is fl-balanced iff the children of every node in T are 6-balanced.



A full binary tree is 1-balanced and a binary tree whose height equals its size (i.e., number of nodes) is 0-balanced.




41


Lemma 2 If the binary tree T is /-balanced, then it is 1-balanced for 0 < 7 < /.


Proof Follows from the definition of balance. 0




Lemma 3 If the binary tree T is #-balanced, 0 < /# 5 1/2, then it is in WB(a) for a = #/(1 + ). Proof Consider any node p in T. Let 1 and r be node p's left and right children.


s(l) + 1 1
B(p) = ---------- = 1
B(p) = s(1) + s(r) + 2 1 + s(L)+1


Since T is #-balanced, s(l) 1 < s(r)/4 or s(l) + 1 < s(r)/ + 2. So, s()+1 < 1/3 + 21-1
s(r) + 1 (s(r) + 1) 1/


or
s(r) + 1 >
s(l) + 1

So, B(p) < 1/(1 + /). Further, s(r) 1 < s(1)/3. So, s(r) + 1
< 1/#t.
s(1) + 1


And, B(p) 1/(1 + 1//) = #/(1 + f). Hence 8/(1 + #) < B(p) < 1/(1 + /) for every p in T. So, T is in WB(a) for a = 8/(1 + /3). 0




42




00
00
0/O O
0 0


Figure 3.2. A tree in WB(1/4) that is not --balanced


Remark While every #-balanced tree, 0 < < 1/2, is in WB(a) for a = /(1+ /), there are trees in WB(a) that are not 3-balanced. Figure 3.2 shows an example of a tree in WB(1/4) that is not -balanced.


Lemma 4 If T is a COST then T is -balanced. Proof If T is a COST, then every subtree of T is a COST. Consider any subtree with root p, left child 1, and right child r. If neither 1 nor r exist, then s(1) = s(r) = 0 and p is -balanced. If s(l) = 0 and s(r) > 1, then r has a nonempty subtree with root t and s(t) > s(1). So p is not a COST. Hence, s(r) < 1 and p is -balanced. The same is true when s(r) = 0. So, assume s(l) > 0 and s(r) > 0.

If s(1) = 1, then s(r) < 3 as otherwise, one of the subtrees of r has m > 2 nodes and m > s(l) implies p is not a COST. Since s(r) 5 3, 1(s(r) 1) < s(1) and (s(l)- 1) s(r). So, p is -balanced. The same proof applies when s(r) = 1. When s(1) > 1 and s(r) > 1, let a and b be the roots of the left and right subtrees of 1. Since p is a COST, s(a) < s(r) and s(b) 5 s(r). So, s(l) = s(a) + s(b) + 1 < 2s(r) + 1 and (s(l) 1) s(r). Similarly, (s(r) 1) s(l). So, !-(l,r). Since this proof applies to every nodes in T, the children of every p are -balanced and T is -balanced. O




43




~0
0 0 0
0 --0 0__Figure 3.3. '-balanced tree that is not a COST


Remark There are !-balanced trees that are not COSTs (see Figure 3.3).



While a COST is in WB(1/3) and WB(a) trees can be maintained efficiently only for 2/11 < a < 1 1/V2- z 0.293, a COST is better balanced than WB(a) trees with a in the usable range. Unfortunately, we are unable to develop O(log n) insert/delete algorithms for a COST.

In the next section, we develop insert and delete algorithms for #-balanced binary search trees (P-BBST) for 0 < 0 V2 1. Note that every (v'2- 1)-BBST is in WB(a) for a = 1 1/v/r- which is the largest permissible a. Since our insert and delete algorithms perform rotations along the search path whenever these result in improved search cost, BBSTs are expected to have better search performance than WB(a) trees (for a = //(1 + fl)).

Each node of a /-BBST has the fields LeftChild, Size, Data, and RightChild. Since every /-BBST, P3 > 0, is in WB(a), for a > 0, fl-BBSTs have height that is logarithmic in n, the number of nodes (provided # > 0).




44



3.4 Search, Insert, and Delete in a j8-BBST

To reduce notational clutter, in the rest of the chapter, we abbreviate s(a) by a (i.e., the node name denotes subtree size).

3.4.1 Search

This is done exactly as in any binary search tree. Its complexity is 0(h) where h is the height of the tree. Notice that since each node has a size field, it is easy to perform a search based on index (i.e., find the 1O'th smallest key). Similarly, our insert and delete algorithms can be adapted to indexed insert and delete.

3.4.2 Insertion


To insert a new element x into a /3-BBST, we first search for x in the /3-BBST. This search is unsuccessful (as x is not in the tree) and terminates by falling off the tree. A new node y containing x is inserted at the point where the search falls off the tree. bet p' be the parent (if any) of the newly inserted node. We now retrace the path from p' to the root performing rebalancing rotations.

There are four kinds of rotations LL, LR, Rb, and RR. Lb and RR rotations are symmetric and so also are bR and Rb rotations. The typical configuration before an Lb rotation is performed is given in Figure 3.4(a). p' denotes the root of a subtree in which the insertion was made. bet p be the (size of the) subtree before the insertion. Then, since the tree was a /3-BBST prior to the insertion, 03-(p, d). Also, for the Lb rotation to be performed, we require that (q > c) and (q > d). Note that q > d implies q > 1. 'Ae shall see that #~-(q, c) follows from the fact that the insertion is




45




9P p

LL
d q gp


q c c d

(a) before (b) after

Figure 3.4. LL rotation for insertion


made into a /3-BBST and from properties of the rotation. Following an LL rotation, p is updated to be the node p".


Lemma 5 [LL insertion lemma] If [f6-(p, d) A /3-(q, c) A (q > c) A (q > d) for 0 < < 1/2 before the rotation, then 0-(q, gp') and 0-(c,d) after the rotation. Proof Assume the before condition.

(a) f(q- 1) < c (as f-(q,c)) < gp'. Also, f(gp'- 1) =/(c+d) < 2/)3q (as />0, q > c and q > d) < q (as 3 < 1/2). So, 0-(q, gp').

(b) d < q :, d- 1 < q- 1 =, 0(d- 1) <5 (q- 1) < c (as #-(q,c)). Also, ,8(c-1)


In an LR. rotation, the before configuration is as in Figure 3.4(a). However, this time q < c. Figure 3.4(a) is redrawn in Figure 3.5(a). In this, the node labeled c in Figure 3.4(a) has been labeled q and that labeled q in Figure 3.4(a) has been labeled




46




gp
q'

P d LR(i)


a q
a b c d

b c

(a) before (b) after substep (i)

Figure 3.5. Substep (i) of insertion LR rotation


a. With respect to the labelings of Figure 3.5(a), rotation LR is applied when



[(q > a) A (q > d)].



The other conditions that apply when an LR rotation is performed are



[/-(p, d) A /3-(a,q) A /3-(b, c)].



Here p denotes the (size of the) left subtree of gp prior to the insertion. An LR rotation is accomplished in two substeps (or two subrotations). The first of these is shown in Figure 3.5(b). Following an LR rotation, p' is updated to be node q'. Lemma 6 [LR substep(i) insertion lemma] If [f-(p,d) A 3-(a,q) A /3-(b,c) A (q > a)A(q > d)]forO < P < 1/2 before the subrotation, then [3-(p",gp')A{(3-(a,b)A
(cd V ( -(ab) A -(cd0 after the subrotation. (c, d)) V (--(a, b) A 03-(c, d))}] after the subrotation.




47



Proof Assume the before condition. First, we show that 0-(p",gp') after the rotation. Note that O(p"- 1) = #(a + b) = #(a + b+ c + 1)- 3(c + 1) = p(p'- 1) /3(c+1) = 8(p-1)-8c < d-Oc < d< gp'. Also, /3(gp'- 1) = 8(c+d) b+#+#d (as #-(b,c)) 5 b + Sq (as q > d) 5 b + a + P (as #-(a,q)) < p" (as # 5 1/2 and p" = a + b + 1). So, #-(p", gp').

Next, we prove two properties that will be used to complete the proof. Pl: #(b- 1) < a.

To see this, note that fl(b 1) /3(q 1) < a (as 0-(a, q)). P2: (c 1) < d.

For this, observe that p' 1 = a + q 8(q 1) + q (as /-(a, q)) = ( + 1)(q 1) + 1. So, q- 1 < 2 = ~+. Similarly, q- 1 = b + c > fl(c- 1) + c (as #-(b,c))
--P--OP-1 <) d (---aq )< o
= (/3 + 1)(c- 1) + 1. So, 9(c- 1) < 'L(q 2) (q 1) 5 < 1 (as

/-(p, d)) < d.

To complete the proof of the lemma, we need to show


{(-(ab)A -(cd)) V( (a,b)A #-(c,d))}.
1+/ 1+/


We do this by considering the two cases b > c and b < c. Case b > c: Since a < q = b + c + (a -1) 0(b + c) < 2b < b. This and P1 imply 3-(a, b). Also, d < q = b + c + 1. So, +O(d 1) 5 + (b + c -) = c + -(b- 1) < -Lc + (as #-(b,c)) = c. This, together with P2 implies

1-(c, d). So, 0-(a, b) A -(c, d).

Case b < c: Since a < q = b + c + 1,a- 1 < b + c. So, a- 1 < b+c- 1 or




48


A- 1+0 1 13 + 1 + 0 1s-+b

Also, d- 1< q-2= b+c-1. So, fl(d- 1) : 3(b+c-1) <3(2c- 1) c. This, together with P2 implies #3-(c, d). So, ---- (a, b)A/3-(c, d). 0



Since an LR(i) rotation can cause the tree to lose its #-balance property, it is necessary to follow this with another rotation that restores the /-balance property. It suffices to consider the two cases of Figures 3.6 and 3.7 for this follow up rotation. The remaining cases are symmetric to these. In Figures 3.6 and 3.7, p and d denote the nodes that do not satisfy #-(p, d). Note, however, that these nodes do satisfy

1 1 16(p, d).

Since the follow up rotation to LR(i) is done only when


d~, ) A d)),
1+


either #(p- 1) > d or 3(d-1) > p. When fl(p- 1) > d, the second substep rotation is one of the two given in Figures 3.6 and 3.7. When 93(d 1) > p, rotations symmetric to these are performed. In the following, we assume /(p 1) > d. Further, we may assume d > 0, as d = 0 and 1+--(p, d) imply p < 1. Hence, /3(p, d). Also, d > 0 and fl(p- 1)> dimplyp> 1.

The LR(ii) LL rotation is done when the condition



A=(q>d)A(c<(1+3)q+(1-3))AB where




49




gp p
LR(ii)
d
d q gp'
LL q


q c c d

(a) before (b) after
Figure 3.6. Case LL for LR(ii) rotation


B = (p,d) A (-#-(p,d)) A -(q,c)A(fl(p 1) > d > 0).
1+0

Lemma 7 [Case LR(ii) LL rotation] If A holds before the rotation of Figure 3.6, then 0-(q, gp') and #-(c, d) after the rotation provided 0 < < V2 1. Proof (a) fl-(q,gp'):

O(q-1) c (as 0-(q,c)) < gp'. Also, 0(gp'-1) = 0(c+d) < #((1+P)q+(1-0)+d) 5 #(1 + O)q + 3(1 0) + #(q 1) (as q> d) = 0(2+ 8)q /2 < q (as 0(2+ 0) < 1 for 0 < p < v'2 1). So, fl-(q, gp').

(b) f-(c, d):

fl(d 1) < #(q 1) 5 c (as P-(q, c)). And, fl(c 1) = (c 1) + 0+(c
q + (c 1) = (q + c 1) = (P- 2) < (p 1) < d (as '3+-(p, d)).


So, fl-(c, d). O




Lemma 8 If (c < (1 + fl)q + (1 0)) A (fl(p 1) > d) in Figure 3.6, then d < q provided 0 < 3 < V2 1.




50


Proof Sinced < 0(p-1) = O(q+c) < O(q+(1+O)q+1-0) = (3+2)q+0(1-3) < q + 1 (as (# + 2) 1 and 0(1- ) < 1 for 0 < /_ v/2-1). So, d q. O



So, the only time an LR(ii) LL rotation is not done is when C = (Cl V C2) A B holds where

C1, = (q= d) A (c < (1 + #)q+1 ) C2=c> (1 +/3)q+(1- #).


At this time, the LR rotation of Figure 3.7 is done. In terms of the notation of Figure 3.7, the condition C becomes D = (D1 V D2) A E where



D = (a d) A (q < (1 + ~3)a + 1 /0)


D2=q > (1 + #)a + 1-/0


E =1 (p,d) A -0-(p,d) A fl-(a,q) A fl-(b,c) A (i(p- 1) > d > 0).
1+#


Lemma 9 When an LR(ii) LR rotation is performed and/3 < V2 1, q > d and so search cost is reduced.


Proof If D1, thensince d d/l3 d > d as S5 V 1. If D2, then d < (p- 1) = (a + q) < (q- + q) = q <
q q 1+(a s-)3 1 ). O
@A< : (as 0< Vf2--1). []
1+0 -




51




gp


pd LR(ii)

LR
a qxx
a b c d

b c

(a) before (b) after

Figure 3.7. Case LR for LR(ii) rotation


Lemma 10 When (d = a) A #3-(b, c) A (fl(p 1) > d) A (3 V2- 1) (see Figure S. 7), /3(a 1) < b and 03(d 1):5 c. Proof Since f(p -1) >dand d =a, O(p -1) > aorl#(a +q) > aor a( -8)


1(a -1) < 2(+b#,1-fl /3

1-0l 1-0l
#- + Il)b + #(#2 +# Ii+ p)
1-0l 1-0l
#_ +l I)b + 0(#2 2 -1)
1-f 1-fl





52


Since O(c- 1) < b,c < + 1. So,


#" f1" b #(3( + 1)b 302O(a-1)< 12(b+c+1)< # (b+ -2)_ 1- + 1- .
1-01- $1-13 1-13


So,
S+ 1 3# -1
a-1< b+ .
1-0 1- 0

However, since #12 + 2# 1 < 0 for f< vf 1, (1 + #)/(1 #) < and

(30-1)/(1- ) <10. So,a-l < b/y+#. Ifa>c+1, thenc



Lemma 11 [Case LR(ii) LR rotation] If D holds before the rotation of Figure 3.7, then #-(p', gp'), #-(a, b), and #-(c, d) following the rotation provided 0 < 1 5 1. Proof (a) -(p',gp'):

1(gp' 1) = #(c + d) 5 b + 3 + 8d (as 0-(b,c)) < b + 0 + 3q (from Lemmas 9 and 10, q >d) < b+ + a+ 13= a+ b+ 20 < a+b+ 1 = p'. Also, since +--(p,d) and q d,,3(p 1) 5 (1 + 1)d or #(a + q) 5 (,8 + 1)d or a + q < (1 + )d or a<(1+ -)d-q5(1+ )d-d=d/f. So, #(p'-1)=/3(a+b)
(as 0-(b,c)) < d+ c + 1 = gp'.

(b) 0-(a, b):

Since b < q and 13-(a,q),(b- 1) B (q- 1) < a.




53


When DI, 13(a 1) b was proved in Lemma 10. So, f3- (a, b). When D2, q a( +3)+ 1 So,







So,



#(aI 1fib +1c+/ O b+3 1+-P13 b



So, #-(a, b).

(c) 0-(c, d):

Note that 13(c 1) < 8(q 1) < (q- 1< '3(p -1) <5d. When DI, fl(d 1) c was proved in Lemma 10. So, 03-(c, d). When D2, if d < b +1, then d< b and /3(d -1) 138(b- 1) < c. So, assume d> b +1. Now, b < d 1 < O(p 1) 1. So,



b < 3(a+b+c )-i

q1 + # +b+c+1)-1


1P (b~c13(1 3)(b+c+ 1)) -1
1+3

1 + 1 c+ +0+ ( +3 + c+ 1)) 1


c+13(1+fl)c+213-1

= (2+13)c+3,3-1 <(2+L)c+3(asI3 V2--1)




54


f< c+ (as < V2- 1).


Also, from d < f/(p 1) and the above derivation, we get


d < (b+c+#+(1+#)(b+c+1))

< 1 + + C+, +T(1+0)( + 0 +c+ l))
1+ #1

j ,/311 2/32 1+13
(2+/ + + 22c) + ( + 1)
= (2+/)c+ 1+c+ (1 ) +1) 1++ 0
2/32 1 2+3+ + 2
= (2 1 f+)c + 1+c+

S(2 + )c+ + 42 + 1

(2+03)c+ 1 (aS#334/32/3< 1 +# for# v'2- 1).


So, f(d- 1) # (2 + P)c < c (as P V 1). So, -(c,d). 0



Theorem 9 If T is fl-balanced, 0 < #3 < V/ 1, prior to insertion, it is so following the insertion.

Proof First note that since all binary search trees are balanced for 1 = 0, the rotations (while unnecessary) preserve 0-balance. So, assume /3> 0. Consider the tree T' just after the new element has been inserted but before the backward restructuring pass begins.

If the newly inserted node, z, has no parent in T', then T was empty and T' is /-balanced. If z has a parent but no grandparent, then T has at most one nonempty




55


subtree X. Since T is fl-balanced, fl([XI 1) < 0. So, IXf < 1. Following the insertion, T' has one subtree with < 1 nodes and one with exactly one. So, T' is #l-balanced. We may therefore assume that z has a grandparent in T'.

From the downward insertion path, it follows that all nodes u in T' that have children 1 and r for which -#fl-(l, r) must lie on the path from the root to z. During the backward restructuring pass, each node on this path (other than z and its parent) play the role of gp in Figures 3.4 and 3.5. The #-property cannot be violated at z as z has no children. It cannot be violated at the parent, s, of z as s satisfied the fl-property prior to insertion. As a result its other subtree has < 1 element. So, following the insertion, s satisfies the fl-property. As a result, each node in T' that might possibly violate the fl-property becomes the gp node during the restructuring pass. Consider one such gp node. It has children in T' denoted by p' and d. Its children in T are p and d. Figures 3.4 and 3.5 show the case when d is the right subtree of gp in both T and T'. The cases RR and RL arise when d is the left subtree.

During the restructuring pass, gp begins at the grandparent of z and moves up to the root of T'. If z is at level r in T', (the root being at level 1), then gp takes on r 2 values during the restructuring pass. We shall show that at each of these r 2 positions either

(a) no rotation is performed and all descendants of gp satisfy the fl-property or

(b) a rotation is performed and following this, all descendants of node p" (Figure 3.4) or of node q' (Figure 3.5) satisfy the fl-property.




56


As a result, following the rotation (if any) performed when gp becomes the root of T', the restructured tree is /-balanced. The proof is by induction on r. When r = 3 (recall, we assume z has a grandparent), gp begins at the root of T' and its descendants satisfy the fl-property.

Without loss of generality, assume that the insertion took place in the left subtree of gp. With respect to Figure 3.4, we have three cases: (i) q c and q > d, (ii) q < c and c > d, and (iii) q < d and c < d. In case (i), all conditions for an LL rotation hold and such a rotation is performed. In case (ii), an LR rotation is performed. Following either rotation, T' is /-balanced. In case (iii), /3(p' 1) = /3(q + c) < 2/3d < d (as ,6 _< v/2- -1). Also, fl(d 1) < p < p + 1 = p'. So, fl(d 1) < p'. Hence, #-(p', d) and T' is /-balanced.

For the induction hypothesis, assume (a) and (b) whenever r < k. In the induction step, we show (a) and (b) for trees T with r = k + 1. The subtree in which the insertion is done has r = k. So, (a) and (b) hold for all gp locations in the subtree. We need to show (a) and (b) only when gp is at the root of T'. This follows from Lemmas 5, 6, 7, and 11.

The theorem now follows. 0




Lemma 12 The time needed to do an insertion in an n node /3-BBST is O(log n) provided 0



57



Proof Follows from the fact that insertion takes 0(h) time where h is the tree height and h = O(log n) when /3> 0 (Lemmas 1 and 3). 0




3.4.3 Deletion

To delete element x from a /3-BBST, we first use the unbalanced binary search tree deletion algorithm of Horowitz and Sahni [14] to delete x and then perform a series of rebalancing rotations. The steps are: Step 1 [Locate x] Search the f3-BBST for the node y that contains x. If there is no

such node, terminate.


Step 2 [Delete x] If y is a leaf, set d to nil, gp, to the parent of y, and delete node

y. If y has exactly one child, set d' to be this child; change the pointer from the parent (if any) of y to point to the child of y; delete node y; set gp to be the parent of d'. If y has two children, find the node z in the left subtree of y that has largest value; move this value into node y; set y = z; go to the start

of Step 2. { note that the new y has either 0 or 1 child }


Step 3 [Rebalance] Retrace the path from d' to the root performing rebalancing

rotations.


There are four rebalancing rotations LL, LR, RR, and Rb. Since Lb and RR as well as LR and Rb are symmetric rotations, we describe LL and LR only. The discussion is very similar to the case of insertion. The differences in proofs are due to the fact that a deletion reduces the size of encountered subtrees by 1 while an




58







g p p 1
LL
pd' qgp


q c c 9

(a) before (b) after

Figure 3.8. LL rotation for deletion insertion increases it by 1. In an LL rotation, the configuration just before and after the rotation is shown in Figure 3.8. This rotation is performed when q > c and q > d'. Following the rotation, d' is updated to the node p'.

Let d denote the size of the right subtree of gp before the deletion. So, d = d + 1. Since prior to the deletion the I3-BBST was #-balanced, it follows that f-(p, d) and 0-(q, c).


Lemma 13 [LL deletion lemma] If [f#-(p,d)A 0-(q,c)A(q c)A(q > d)A(1/3 < 3 1/2)] before the rotation, then [#-(q, gp') A 8-(c, d')] after the rotation. Proof (a) fl-(q, gp'):

/3(q -1) !5 c (as 8)-(q,c)) < gp'. Also, B3(gp' -1) = fl(c+ d') < 2,8q (as c < q and d' < q) < q (as )3 < 1/2). So, 0-(q, gp').

(b) #-(c,d'):

d' < q = d'-1 < q-1 =: (d'-1) < fl(q-1) < c. Also, when c < 1, 0(c-1) < 0 < d' (as d' > 0). When c > 1,q > c =z q > 2 and p = q+c+ 1 > c+3. So, 0(c- 1) < /(p- 1)- 3/3< d- 3/3 (as 0-(p,d)) < d- 1 (as fl > 1/3) = d'. Hence,




59


/3-(c, d'). 0



In an LR rotation, the before configuration is as in Figure 3.8(a). However, this time q < c. Figure 3.8(a) is redrawn in Figure 3.9(a). In this, the node labeled c in Figure 3.8(a) has been relabeled q and that labeled q in Figure 3.8(a) has been relabeled a. With respect to the labelings of Figure 3.9(a), rotation LR is applied when

[(q > a) A (q > d')].


The other conditions that apply when an LR rotation is performed are



[fl-(p, d) A f8-(a, q) A fl-(b, c)J.



Here d denotes the (size of) right subtree of gp prior to the deletion. As in the case of insertion, an LR rotation is accomplished in two substeps (or two subrotations). The first of these is shown in Figure 3.9. Following an LR rotation, d' is updated to node q'.


Lemma 14 [LR substep(i) deletion lemma] If [f3-(p, d) A #3-(a, q) A #-(b, c) A (q > a) A (q > d')] before the subrotation LR(i), then [3-(p', gp') A{ (/3-(a, b) A '3--(c, d))V( (a, b) A /-(c, d'))}] after the subrotation provided 1/3 < 3 < 1/2 Proof Assume the before condition.

(a) If b =c= 0, then q b+c+ 1 = 1. Furthermore, (q > a) and (q > d') imply a d'= 0. So, gp' = p'= 1. Hence, [!-(p',gp') A !-(a,b) A !-(c,d')]




60




gp
q'

P d' LR(i)


a q
a b c d'

b c

(a) before (b) after substep (i)

Figure 3.9. LR rotation for deletion


(b)Ifb=landc=0,thenq=2,a<1, andd'<1. So, 1
(c) Ifb=O0andc= 1, thenq=2, a<1, andd' < 1. So, 1 1 and c > 1. So, q >3, a> 1 (as -(a,q) i 6(q- 1) < a or a> 20 > 0), p = a + q + 1 > 5, d >2 (as #-(p, d) = (p 1) < d and f > 1/3), and d' = d 1 > 1.

First, we show that #-(p', gp'). For this, note that a + b + c + 1 = p 1. From #-(p,d), it follows that #(a + b+ c + 1) = /(p- 1) < d. So, 8(a + b) 5 d- c c- f. From Figure 3.9(b), we see that (p'- 1) = 3(a + b). Hence, I(p'- 1) 5 d- flc- = d'-,c+1-0 5 d'+1-2,8


(gp'- 1) =,(c + d') 5< b + + +d' (as ,-(b,c)) < b+Oq+O(asq>d')




61


< b + a + 2 (as -(a,q))

< p.



So, 0-(p', gp').

Next, we prove two properties that will be used to complete the proof. PI: /(b- 1) < a.

To see this, note that fl(b 1) < /(q 1) !5 a (as #-(a, q)). P2: l(c- 1) < d'.

For this, observe that fl(c 1) < P(q 2) (as c < q 1) < #(p 4) (as q = p a 1 and a > 1) = #/(p- 1)- 3/3 1/3) = d'.

To complete the proof of the lemma, we need to show



1043-(a, b) A # d)) V -(a, b) A #-(c, ))}.



For this, consider the two cases b > c and b < c (as in Lemma 6). Case b > c: Since a < q = b + c + 1,/#(a 1) < /(b + c) < 203b < b. This, together with P1 implies 3-(a, b). Also, d' < q =b+c+ 1. So, -6+(d' -1)< --(b +c- 1) =
--g-c~+ + +0(-+1+c
+ --- 1) c + c = c. This, together with P2 implies a --(c, d'). So,


/-(a,b) A 1 -(c, d').

Case b < c: Since a < q = b + c + 1,a 1 < b + c. So, a 1 < b + c 1 or or1+ + 1+ 1+ /3+1/3
d- 1 < q 2 = b + c 1. So, fl(d 1) < P(b + c 1) < 3l(2c 1) < c. This and




62


P2 imply 3-(c, d'). Hence, +-(a, b) A 3-(c,d'). 0



The substep(ii) rotations are the same as for insertion.


Theorem 10 If T is #-balanced, then following a deletion the resulting tree T' is also 3-balanced provided 1/3 < fl < -' 1.


Proof Similar to that of Theorem 9. 0



When 0 < # < 1/3, we need to augment the LL rotation by a transformation for the cased' = 0. Whend'= 0,/(p- 1) < d = d' + 1 = 1. So, p < 1/fl + 1 and gp = p + d' + 1 < 1/ + 2. To fl-balance at gp, the at most 1/ + 2 nodes in gp are rearranged into any fl-BBST in constant time (as 1/# + 2 is a constant). When d' > 0, the proof of Lemma 13 part (b) can be changed to show 0(c 1) < d' for 0 < < v2 1. The new proof is: since c < q,c < (p 1)/2 and fl(c 1) < (p- 1)/2 !5 d/2- # = d- d/2 0 < d- 1 0 < d'. The LR rotation needs to be augmented by a transformation for the case d' = d- 1 < 1. At this time, fl(p- 1) 5 d < (2) So, gp = p + d< 2( )+ 1 + To #-balance at

gp, we rearrange the fewer than + 1 + nodes in the subtree, in constant

time, into any #-balanced tree. When d' > 1 1, the proof for 3(c 1) 5 d' in 0(2+0)
Lemma 14 needs to be changed to show that the LR substep(i) lemma holds. The new proof is:



d > 01(p 1) = (a + b + c + 1) >(1(q-1) + b + c + 1bc+1)




63


= 1(j3(b+c)+b+c+1)

N + ((1+/)(c- 1) + (1 + O)c+ 1)

= #((1 +/3)2(c- 1)+2+/).


So, 1) d-2-2 So ~~~(1+,6)2-d-1(ad>

Also, note that when = 0, all trees are 8-balanced so the rotations (while not needed) preserve balance.


Theorem 11 With the special handling of the case d' = 0, the tree T' resulting from a deletion in a fl-BBST is also #-balanced for 0 < fl < V2- 1. Lemma 15 The time needed to delete an element from an n node fl-BBST is O6og n) provided 0 < /3 < V 1.


3.4.4 Enhancements

Since our objective is to create search trees with minimum search cost, the rebalancing rotations may be performed at each positioning of gp during the backward restructuring pass so long as the conditions for the rotation apply rather than only at gp positions where the tree is unbalanced.

Consider Figure 3.4(a). If p' < d, then the conditions of Lemmas 5 and 6 cannot apply as q < p' < d. However, it is possible that e > p' where e is the size of either the left or right subtree of d. In this case, an RR or RL rotation would reduce the total search cost. The proofs of Lemmas 5 and 6 are easily extended to show that these rotations would preserve balance even though no insertion was done in the subtree




64



d. The same observation applies to deletion. Hence the backward restructuring pass for the insert and delete operations can determine the need for a rotation at each gp location as below (I and r are, respectively, the left and right children of gp).

if s(l) > s(r) then check conditions for an Lb and LR rotation

else check conditions for an RR and RL rotation.

The enhanced restructuring procedure used for insertion and deletion is given in Figure 3.10. In the RR and RL cases, we have used the relation '>' rather than'> as this results in better observed run time.

Since it can be shown that the rotations preserve balance even when there has been no insert or delete, we may check the rotation conditions during a search operation and perform rotations when these improve total search cost.

Finally, we note that it is possible to use other definitions of 18-balance. For example, we could require 83(s(a) 2) < s(b) and 86(s(b) 2) < s(a) for 83-(a, b). One can show that the development of this chapter applies to these modifications also. Furthermore, when this new definition is used, the number of comparisons in the second substep of the LR and Rb rotations is reduced by one.

3.4.5 Top Down Algorithms


As in the case of red/black and WB(a) trees, it is possible to perform, in O(log n) time, inserts and deletes using a single top to bottom pass. The algorithms are similar to those already presented.




65



procedure Restructuring ; begin
while (gp) do
begin
if (s(gp.left) > s(gp.right)) then
begin {check conditions for an LL and LR rotation}
p = gp.left ;
if (s(p.left) > s(p.right)) then
begin if (s(p.left) > s(gp.right)) then do LL rotation; end
else
begin
if (s(p.right) > s(gp.right)) then {LR}
begin
do LR rotation ;
{ now notations a, b,c, and d follow from figure 3.1(b) }
if ((s(a) 1) > s(b)) then
if ((s(a.right) < (1 + )s(a.left) + 1 8) and (s(b) < s(a.left))) then do LL rotation else do LR rotation else if (P(s(d) 1) > s(c)) then
if ((s(d.left) < (1 + )s(d.right) + 1 4) and (s(c) < s(d.right))) then do RR rotation else do RL rotation; end
end
end
else {check conditions for an RR and RL rotation}
begin
p = gp.right;
if (s(p.left) > s(p.right)) then
begin
if (s(p.left) > s(gp.left)) then {RL}
do symmetric to the above LR case;
end
else
begin if (s(p.right) > s(gp.left)) then do RR rotation; end
end ;
gp = gp.parent;
end;
end;

Figure 3.10. Restructuring procedure




66


3.5 Simple 9-BBSTs

The development of Section 3.4 was motivated by our desire to construct trees with minimal search cost. If instead, we desire only logarithmic performance per operation, we may simplify the restructuring pass so that rotations are performed only at nodes where the fl-balance property is violated. In this case, we may dispense with the LL/RR rotations and the first substep of an LR/RL rotation. Only LR/RL substep (ii) rotations are needed. To see this, observe that Lemmas 7 and 11 show that the second substep rotations rebalance at gp (see Figures 3.6 and 3.7) provided '- l+13

(p, d) (The remaining conditions are ensured by the bottom-up nature of restructuring and the fact the tree was fl-balanced prior to the insert or delete).

If the operation that resulted in loss of balance at gp was an insert, then fl(p 2) < d (as p > d, the insert took place in subtree p and gp was 8l-balanced prior to the insert) and fl(p 1) > d (gp is not f-balanced following the insert). For the substep (ii) rotation to restore balance, we need fl(p 1) < (1 + f3)d. This is assured if d+/ : (fl+ 1)d (asfl(p-2) < d). So, we need d> 1. Ifd < 1, then d= 0. Now 0(p 2) < d and f(p 1) > d imply p = 2. One may verify that when p = 2, the LR(ii) rotations restore balance.

If the loss of f-balance at gp is the result of a deletion (say from its right subtree), then 0(p 1) < d + 1 (as gp was f-balanced prior to the delete). For the substep

(ii) rotation to accomplish the rebalancing, we need fl(p 1) < (f8 I- 1)d. This is guaranteed if d + 1 < (f8 + 1)d or d > 1/fl. When d < 1/fl and 8 > 1/3, d < 2. Since 0(p-1) < d+1 and 9> 1/3, when d=2, p < 10; whend= 1, p< 7; and when




67


d =0, p < 4. We may verify that for all these cases, the LR(ii) rotations restore balance. Hence, the only problematic case is when P3 < 1/3 and d < 1/f8.

When /3 < 1/3, an Lb rotation fails to restore balance only when d = 0 (see discussion following Theorem 10). So we need to rearrange the at most 1/0l + 2 nodes in gp into any fl-balanced tree when d = 0. An LR rotation fails only when d < -1. To see this, note that in the terminology of Lemma 14, d is d.
0(2+03)

The proof of P2 is extended to the case #3 1/3 when d'> 1 _1. Also, since d <1/fl, for the case b > c, we get fl(d'- 1) < 1-f < C (as c > 1). For the case b < c, we need to show fl(a 1) < b. Since an LR rotation is done only when condition D1 V D2 holds, from Lemmas 10 and 11, it follows that 8l(a 1) < b. So, an LR rotation rebalances when P3 < 1/3 provided d > 1 1. For smaller d, the at most 12(2+,) 0 1(2+0) nodess in the subtree gp may be directly rearranged into a fl-balanced tree.

The restructuring algorithm for simple fl-BBSTs is given in Figures 3.11 and 3.12. The algorithm of Figure 3.11 is used following an insert and that of Figure 3.12 after a delete.

Simple fl-BBSTs are expected to have higher search cost than the fl-BBSTs of Section 3.4. However, they are a good alternative to traditional WB(a) trees as they are expected to be "better balanced". To see this, note that from the proof of Lemma 3, the balance, B(p), at any node p in a fl-balanced tree satisfies


1 1+s(r)+ 1
B(p) ls(1) 1




68



procedure Restructuring2; begin
while (gp) do
begin
if (fl(s(gp.left) 1) > s(gp.right)) then {do an LL or LR rotation}
begin
p = gp.left;
if ((s(p.right) < (1 + )s(p.left) + 1 /3) and
(s(gp.right) < s(p.left))) then
do LL rotation
else do LR rotation ;
end
else
do symmetric to the above L case;
gp = gp.parent ;
end;
end;

Figure 3.11. Simple restructuring procedure for insertion


procedure Restructuring3 begin
while (gp) do
begin
if (fl(s(gp.left) 1) > s(gp.right)) then
if (/3 < 1/3) and (s(gp.right) < 1//3(2 + /3) 1) then
rearrange the subtree rooted at gp into any /3-balanced tree
else {do an LL or LR rotation}
begin
p = gp.left ;
if ((s(p.right) < (1 + fl)s(p.left) + 1 /3) and
(s(gp.right) < s(p.left))) then
do LL rotation
else do LR rotation ;
end
end
else
do symmetric to the above L case;
gp = gp.parent;
end;
end ;

Figure 3.12. Simple restructuring procedure for deletion




69


1
S1+ 1/3 + p(s(r)+1)
1 + 2)3-+
0 03((r)+x) 120-1
i +
13 1(3)1)


So,
1
B(p) < 1i + ((r)+1)

Also, sinces(r)- 1 < s(1)/fl, s(r) + 1 < s(l)/p + 2. Hence, 1 + < 1 + (( +
8Q+ 03((1)+'
2 So s(1)+1' So


1
B(p) 1 + +

1
-(1) T1 (s(l)+1)


Consequently,


1 1
I1 2-1 < B(p) < 1 + + 2-1
1 + (s(I)+1) -1 ( (r)+1)


When 3 = 0- 1,


1 1
< B(p) < 1
2 + v/ + 2 + v/ + 1 -+v


If s(p) 5 10, 0.296 < B(p) < 1 0.296. So, every #-balanced subtree with 10 or

fewer nodes is in WB(a) for a 0.296. Similarly, every subtree with 100 or fewer

nodes is in WB(a) for a 0.293. In fact, for every fixed k, subtrees of size k or less




70


procedure Restructuring4
begin
while (gp) do
begin
if (s(gp.left) > s(gp.right)) then
begin {check conditions for an LL and LR rotation}
p = gp.left ;
if (s(p.left) > s(p.right)) and (s(p.left) > s(gp.right)) then
do LL rotation
else if (s(p.left) 5 s(p.right)) and (s(p.right) > s(gp.right)) then
do LR rotation ;
end
else {check conditions for an RR and RL rotation}
do symmetric to the above L case;
gp = gp.parent;
end;
end;

Figure 3.13. Simple restructuring procedure without a f8 value

are in WB(a) for a slightly higher than 1 0.2929 which is the largest value

of a for which WB(a) trees can be maintained.

3.6 BBSTs without Deletion

In some applications of a dictionary, we need to support only the insert and

search operations. In these applications, we can construct binary search trees with

total cost

C(T) nlogo(V(n + 1))


by using the simpler restructuring algorithm of Figure 3.13.


Theorem 12 When the only operations are search and insert and restructuring is done

as in Figure 3.13, C(T) < nlog,(v'5(n + 1)).




71


Proof Suppose T currently has m 1 elements and a new element is inserted. Let u be the level at which the new element is inserted. Suppose that the restructuring pass performs rotations at q < u of the nodes on the path from the root to the newly inserted node. Then C(T) increases by at most v = u q as a result of the insertion. The number of nodes on the path from the root to the newly inserted node at which no rotation is performed is also v. Let these nodes be numbered 1 through v bottom to top. Let Si denote the number of elements in the subtree with root i prior to the restructuring pass. We see that S, > 1 and S2 > 2. For node i, 2 < i < v, one of its subtrees contains node i 1. Without loss of generality, let this be the left subtree of i. Let the root of the right subtree of i be d. So,


Si >! Si_ + s(d) + 1.



If i 1 is not the left child of i, then since no rotation is done at i, s(d) > Si-1. If i 1 is the left child of i, then consider node i 2. This is in one of the subtrees of i. Since no rotation is performed at i 1, s(d) > Si-2. Since Si-I > Si_2, we get



Si > Si-1 + S,-2 2 1.



Hence, S,, N where N,, is the minimum number of elements in a COST of height v. So, v < log,(v"5(m + 1)). So, when an element is inserted into a tree that has m 1 elements, its cost C(T) increases by at most logo(v'5(m + 1)). Starting with an empty tree and inserting n elements results in a tree whose cost is at most




72


n logj(v' (n + 1)). 0




Corollary 2 The expected cost of a search or insert in a BBST constructed as above is O(log n).

Proof Since C(T) < nlogj(v'5(n + 1)), the expected search cost is C(T)/n < log,(V5(n + 1)). The cost of an insert is the same order as that of a search as each insert follows the corresponding search path twice (top down and bottom up). 0



3.7 Experimental Results

For comparison purposes, we wrote C programs for BBSTs, SBBSTs (simple BBSTs), BBSTDs (BBSTs in which procedure Restructuring4 (Figure 3.13) is used to restructure following inserts as well as deletes), unbalanced binary search trees (BST), AVL-trees, top-down red-black trees (RB-T), bottom-up red-black trees (RBB) [31], weight balanced trees (WB), deterministic skip lists (DSL), treaps (TRP), and skip lists (SKIP). For the BBST and SBBST structures, we used P = 207/500 while for the WB structure, we used a = 207/707. While these are not the highest permissible values of # and a, this choice permitted us to use integer arithmetic rather than the substantially more expensive real arithmetic. For instance, ,l-(a, b) for 3 = 207/500 can be checked using the comparisons 207(s(a) 1) > 500s(b) and 207(s(b) 1) > 500s(a). The randomized structures TRP and SKIP used the same random number generator with the same seed. SKIP was programmed with probability value p = 1/4 as in Pugh [26].




7 3



To minimize the impact of system call overheads on run time measurements, we programmed all structures using simulated pointers (i.e., an array of nodes with integer pointers [271. Skip lists use variable size nodes. This requires more complex storage management than required by the remaining structures which use nodes of the same size. For our experiments, we implemented skip lists using fixed size nodes, each node being of the maximum size. As a result, our run times for skip lists are smaller than if a space efficient implementation had been used. In all our tree structure implementations, null pointers were replaced by a pointer to a tail node whose data field could be set to the search /insert/ delete key and thus avoid checking for falling off the tree. Similar tail pointers are part of the defined structure of skip and deterministic skip lists. Each tree also had a head node. WB(a) trees were implemented with a bottom-up restructuring pass. Our codes for SKIP and DSL are based on the codes of Pugh [261 and Papadakis [22], respectively. Our AVL and RB-T codes are based on those of Papadakis [22] and Sedgewick [28]. The treap structure was implemented using joins and splits rather than rotations. This results in better performance. Furthermore, AVL, RB-B, WB, and BBST were implemented with parent pointers in addition to left and right child pointers. For BBSTs, the enhancements described in Section 3.4.4 for insert and delete (see Figure 3.10) were employed. No rotations were performed during a search when using any of the structures.

For our experiments, we tried two versions of the code. These varied in the order in which the 'equality' and 'less than' or 'greater than' check between x and e (where x is the key being searched/inserted/deleted and e is the key in the current




74


node) is done. In version 1, we conducted an initial experiment to determine if the total comparison count is less using the order L:

if x < e then move to left child

else if x e then move to right child

else found

or the order R:

if x > e then move to right child

else if x je then move to left child

else found.

Our experiment indicated that doing the 'left child' check first (i.e. order L) worked better for AVL, BBST, BBSTD, and DSL structures while R worked better for the RB-T, RB-B, WB, SBBST, and TRP structures. No significant difference between L and R was observed for BSTs. For skip lists, we do not have the flexibility to change the comparison order. The version 1 codes performed the comparisons in the order determined to be better. For BSTs, the order R was used.

In the version 2 codes the comparisons in each node took the standard form

if x =e then found

else if x < e then move to left child

else move to right child.

The version 2 restructuring code for BBSTs differed from that of Figure 3.10 in that the '>' test in the second, third, and forth if statements was changed to ''




75


No change was made in the corresponding if statements for RR and RL rotations. While this increased the number of comparisons, it reduced the run time.

We experimented with n = 10,000, 50,000, 100,000, and 200,000. For each n, the following experiments were conducted:

(a) start with an empty structure and perform n inserts;

(b) search for each item in the resulting structure once; items are searched for in the order they were inserted

(c) perform an alternating sequence of n inserts and n deletes; in this, the n elements inserted in (a) are deleted in the order they were inserted and n new elements are inserted

(d) search for each of the remaining n elements in the order they were inserted

(e) delete the n elements in the order they were inserted.

For each n, the above five part experiment was repeated ten times using different random permutations of distinct elements. For each permutation, we measured the total number of element comparisons performed and then averaged these over the ten permutations.

First, we report on the relative performance of SBBSTs, BBSTDs, and 1313STs. For this comparison, we used only version I of the code. Table 3.1 gives the average number of key comparisons performed for each of the five parts of the experiment. The three versions of our proposed data structure are very competitive on this measure. BBSTI)s and BBSTs generally performed fewer comparisons than did SBBSTs. All three structures had a comparison count within 2% of one another.




76



Table 3.1. The number of key comparisons on random inputs (version 1 code)

n operation 11SBBST IBBSTD IBBST]
insert 212495 212223 212111
search 194661 191599 191578
10,000 ins/del 416125 416967 416862 search 194957 191666 191676
delete 168033 166441 166487
insert 1241080 1236682 1236114 search 1152137 1135131 1134969 50,000 ins/del 2437918 2438083 2437639 search 1153821 1134277 1134062
_____ delete 1018675 1007766 1007688 insert 2635913 2624829 2623792 search 2458079 2423988 2423613 100,000 ins/del 5183619 5180383 5179653 search 2461221 2420282 2419990
____ delete 2190798 2168049 2168110 insert 5580139 5555190 5553256 search 5223989 5148220 5147698 200,000 ins/del 10981441 10969578 10968053 search 5229172 5144808 5144148 delete 4692447 4641349 4641389



However, when we used ordered data rather than random data (Table 3.2), SBBSTs

performed noticeably inferior to BBSTDs and BBSTs; the later two remained very

competitive.

Tables 3.3 and 3.4 give the average heights of the trees using random data and

using ordered data, respectively. The first number gives the height following part (a)

of the experiment and the second following part (c). The numbers are identical for

BBSTDs and BBSTs and slightly higher (lower) for SBBSTs using random (ordered)

data.




77



Table 3.2. The number of key comparisons on ordered inputs (version 1 code)

n operation~ SBBST IBBSTD I BBST
insert 170182 150554 150554
search 188722 185530 185530
10,000 ins/del 425305 315177 314998
search 191681 184155 184155
____ delete 215214 135311 135131
insert 991526 872967 872967
search 1117174 1101481 1101481 50,000 ins/del 2472808 1806346 1805439 search 1116390 1098065 1098065 delete 1277756 792717 791815 insert 2103808 1850548 1850548 search 2384327 2354757 2354757 100,000 ins/del 5249194 3823415 3821594 search 2382759 2346118 2346128
_____ delete 2738294 1686397 1684584
insert 4449143 3903083 3903083 search 5068632 4946753 4946753 200,000 ins/del 11105525 8051695 8048058 search 5065496 5001967 5001967
_____ delete 5842168 3580856 13577223



Table 3.3. Height of the trees on random inputs (version 1 code)

n SBBST 1 BBSTD I BBST 10,000 fi17,17 ~f16,16 16,16
50,000 j~20,20 19,19 19,19
100,000 21,21 20,20 20,20
200,000 22,23 21,21 21,21



Table 3.4. Height of the trees on ordered inputs (version 1 code)

II 33SBBST I BBSTD [ BBST 10,000 3]16,15 17,17 17,17
50,000 20,20 20,20 20,20 100,000 21,21 21,21 21,21
200,000 22,22 23,22 23,22




78


The average number of rotations performed by each of the three structures is given in Tables 3.5 and 3.6. A single rotation (i.e., LL or RR) is denoted 'S' and a double rotation (i.e., LR or Rb) denoted 'D'. In the case of BBSTs, double rotations have been divided into three categories: D = LR and Rb rotations that do not perform a second substep rotation; DS = LR and Rb rotations with a second substep rotation of type Lb and RR; DD = LR and Rb rotations with a second substep rotation of type bR and Rb. BBSTIs and BBSTs performed a comparable number of rotations on both data sets. However, on random data SBBSTs performed about half as many rotations as did BBSTIs and BBSTs. On ordered data, SBBSTs performed 15 to 20% fewer rotations on part (a), 34% fewer on part (c), and 51% fewer on part (e).

The run-time performance of the structures is significantly influenced by compiler and architectural features as well as the complexity of a key comparison. The results we report are from a SUN SPARC-5 using the UNIX C compiler cc with optimization option. Because of instruction pipelining features, cache replacement policies, etc., the measured run times are not always consistent with the compiler and architecture independent metrics reported in Tables 3.1 through 3.6 and later in Tables 3.11 through 3.16. For example, since the search codes for all tree based methods are essentially identical, we would expect methods with a smaller comparison count to have a smaller run time for parts (b) and (d) of the experiment. This was not always the case.

Tables 3.7 and 3.8 give the run times of the three BBST structures using integer keys and Tables 3.9 and 3.10 do this for the case of real (i.e., floating point) keys. The





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Table 3.6. The number of rotations on ordered inputs (version 1 code)



n operation S1B1U S B TD ~ S D B DS IDDIinsert 9984 0 9985 2387 9985 2387 0 0
10,000 ins/del 14997 0 16567 6130 16644 5797 25 154
delete 4989 0 6570 3726 6647 3392 26 154
insert 49980 0 49983 11956 49983 11956 0 0
50,000 ins/del 74996 0 82862 30659 83247 28982 137 770
delete 24987 0 32859 18686 33242 17018 136 766
insert 99979 0 99983 23917 99983 23917 0 0 100,000 ins/del 149996 0 165738 61327 166504 57969 280 1540
_____ delete 49986 0 65733 37392 66505 34040 278 1536
insert 199978 0 199982 47839 199982 47839 0 0 200,000 ins/del 299996 0 331473 122653 333012 115938 559 3078
_____ delete 999 85 0 131478 74795 133016 68086 1557 13076J



sum of the run time for parts (a) (e) of the experiment is graphed in Figure 3.14. For random data, SBBSTs significantly and consistently outperformed BBSTDs and BBSTs. On ordered data, however, BBSTDs were slightly faster than BBSTs and both were significantly faster than SBBSTs.

Since BBSTs generated trees with the least search cost, we expect BBSTs to outperform SBBSTs and BBSTDs in applications where the comparison cost is very high relative to that of other operations and searches are done with a much higher frequency than inserts and deletes. However, with the mix of operations used in our tests, SBBSTs are the clear choice for random inputs and BBSTDs for ordered inputs.

In comparing with the other structures, our tables repeat the data for BBSTs. The reader may make the comparison with SBBSTs and BBSTDs.




81














Table 3.7. Run time on random inputs using integer keys (version 1 code)

n~ operation SBBST BBSTD [BBSTI insert 0.27 0.30 0.34
search 0.06 0.06 0.07
10,000 ins/del 0.57 0.62 0.70
search 0.06 0.06 0.06
delete 0.22 0.25 0.26
insert 1.48 1.61 1.75
search 0.35 0.36 0.37
50,000 ins/del 2.90 3.47 3.84
search 0.36 0.38 0.39
delete 1.13 1.47 1.62
insert 3.00 3.5 380
search 0.78 0.83 0.84
100,000 ins/del 6.28 7.78 8.41
search 0.83 0.87 0.88
delete 2.54 3.31 3.58
insert 6.56 7.74 8.37
search 1.80 1.89 1.89
200,000 ins/del 13.89 17.32 18.57 search 1.86 1.98 1.98
delete 5.64 7.41 8.02

Time Unit :sec




89















Table 3.8. Run time on ordered inputs using integer keys (version 1 code)

n operation SBBS BBSTD BBST
insert 0.32 0 -20 0.27
search 0.05 0.03 0.05
10,000 ins/del 0.58 0.43 0.57
search 0.07 0.03 0.03
delete 0.20 0.17 0.23
insert 1.8 1.20 1.10
search 0.25 0.20 0.20
50,000 ins/del 2.63 2.18 2.40 search 0.25 0.20 0.20
delete 0.95 0.92 1.05
insert 3- .43 2.23 2.53
search 0.72 0.45 0.42
100,000 ins/del 5.97 4.70 5.13
search 0.55 0.47 0.42
delete 2.10 1.98 2.15
insert 6.65 4.95 5.25
search 1.20 0.92 0.90
200,000 ins/del 13.13 10.23 10.88 search 1.17 0.90 0.90
delete 4.63 4.25 4.58

Time Unit: see




83













Table 3.9. Run time on random real inputs (version 1 code)

n operation II SBBSTT-[ BBSTD BBST
insert 0.23 0.34 0.36
search 0.07 0.10 0.10
10,000 ins/del 0.44 0.75 0.79
search 0.08 0.10 0.10
delete 0.17 0.29 0.30
insert 1.43 1.76 1.93
search 0.47 0.53 0.52
50,000 ins/del 2.76 3.89 4.22
search 0.50 0.54 0.55
delete 1.13 1.62 1.76
insert 2.96 3.94 4.36
search 1.08 1.17 1.16
100,000 ins/del 6.11 8.58 9.30
search 1.12 1.20 1.22
delete 2.50 3.66 3.95
insert 6.85 8.92 9.33
search 2.41 2.58 2.57
200,000 ins/del 13.86 19.49 20.46
search 2.49 2.69 2.66
delete 5.61 8.25 8.80

Time Unit: sec




84














Table 3.10. Run time on ordered real inputs (version 1 code)

n Operation SBBST BBSTD) BBST]
insert 0.27 0.23 0.20
search 0.08 0.07 0.07
10,000 ins/del 0.53 0.50 0.43
search 0.08 0.07 0.05
delete 0.18 0.23 0.20
insert 1.43 1.25 1.12
search 0.40 0.30 0.30
50,000 ins/del 2.80 2.17 2.37
search 0.40 0.30 0.30
_____ delte 1.07 0.90 0.97
insert 3.28 2.58 2.77
search 0.90 0.62 0.63
100,000 ins/del 6.15 4.70 5.13
search 0.87 0.62 0.63
delete 2.35 1.93 2.10
insert 7.37 4.55 4.92
search 1.85 1.32 1.32
200,000 ins/del 13.35 10.03 10.93
search 1.87 1.33 1.33
_____ delete 5.08 4.17 4.43

Time Unit: sec




85




Time is sum of time for parts (a)-(e) of the experiment 45

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15 10

50000 100000 150000 200000
n

Figure 3.14. Run time on real inputs (version 1 code)



The average number of comparisons for each of the five parts of the experiment are given in Table 3.11 for the version 1 implementation. On the comparison measure, AVL, RB-B, WB, and BBSTs are the front runners and are quite competitive with one another. On parts (a) (insert n elements) and (c) (insert n and delete n elements), AVL trees performed best while on the two search tests ((b) and (d)) and the deletion test (e), BBSTs performed best.

Table 3.12 gives the number of comparisons performed when ordered data (i.e., the elements in part (a) are 1, 2,... ,n and are inserted in this order) and those in part (c) are n + 1,.. ,2n (in this order) is used instead of random permutations of distinct elements. This experiment attempts to model realistic situations in which the inserted elements are in "nearly sorted order". BSTs were not included in this test as they perform very poorly with ordered data taking 0(n') time to insert n




86


times. The computer time needed to perform this test on BSTs was determined to be excessive. This test exhibited greater variance in performance. Among the deterministic structures, BBSTs outperformed the others in parts (a) (d) while AVL trees were ahead in part (e). For part (a), BBSTs performed approximately 45% fewer comparisons than did "L trees and approximately 12% fewer than WB trees. The randomized structure TRP was the best of the eight structures reported in Table 3.12 for part (a). It performed approximately 10% fewer comparisons than did BBST trees. However, the BBST remained best overall on parts (b), (c), and

(d).

The heights of the trees (number of levels in the case of DSL and SKIP) for the experiments with random and ordered data are given in Tables 3.13 and 3.14 respectively. The first number in each table entry is the tree height after part (a) of the experiment and the second, the height after part (c). In all cases, the number of levels using skip lists is fewest. However, among the tree structures, AVL and BBST trees have least height on random data and AVL has least with ordered data.

Tables 3.15 and 3.16, respectively, give the number of rotations performed by each of the deterministic tree schemes for experiment parts (a), (c), and (e). Note that none of the schemes performs rotations during a search.

On ordered data, BBSTs perform about 25% more rotations than do the remaining structures. These remaining structures perform about the same number of rotations. On random data, "L trees, bottom-up red-black trees and WB trees perform a comparable number of rotations. Top-down red-black trees and BBST trees






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EFFICIENT ALGORITHMS AND DATA STRUCTURES FOR VLSI CAD
By
SEONGHUN CHO
A DISSERTATION PRESENTED TO THE GRADUATE SCHOOL
OF THE UNIVERSITY OF FLORIDA IN PARTIAL FULFILLMENT
OF THE REQUIREMENTS FOR THE DEGREE OF
DOCTOR OF PHILOSOPHY
UNIVERSITY OF FLORIDA
1996

ACKNOWLEDGMENTS
My heartfelt appreciation goes to my advisor Professor Sartaj Sahni for giving
me continued guidance in my thesis work. I thank him for the help, patience and sup¬
port he provided throughout my stay in the University of Florida. Weekly meetings
and discussions with him have spawned many ideas, for which I am thankful.
I would like to thank other members in my supervisory committee, Dr Li-Min
Fu, Dr Theodore Johnson, Dr Sanguthevar Rajasekaran, and Dr Paul W. Chun for
their interest and comments.
Thanks go to Venkat Thanvantri for his willingness to discuss the general subject
of algorithms.
Thanks go to my wife Joungyim for her love, encouragement and patience.
Finally, I would like to thank my parents for the love and support, without which I
could not have pursued my doctoral studies. To them I dedicate this work.

TABLE OF CONTENTS
ACKNOWLEDGMENTS ü
LIST OF TABLES vii
LIST OF FIGURES x
ABSTRACT xi
CHAPTERS
1 INTRODUCTION 1
1.1 Background 1
1.2 Dissertation Outline 2
2 MINIMUM AREA JOINING OF COMPACTED CELLS 4
2.1 Introduction 4
2.2 Mayer River Routing 9
2.3 Constraint Graph Representation 16
2.4 Heuristics to Minimize Area 20
2.4.1 Heuristic 1 20
2.4.2 Heuristic 2 21
2.4.3 Heuristic 3 23
2.5 Experimental Results 25
2.6 Conclusion 30
3 A NEW WEIGHT BALANCED BINARY SEARCH TREE 32
3.1 Introduction 32
3.2 Balanced Trees and Rotations 34
3.3 /3-BBSTs 39
3.4 Search, Insert, and Delete in a /J-BBST 44
3.4.1 Search 44
3.4.2 Insertion 44
3.4.3 Deletion 57
3.4.4 Enhancements 63
3.4.5 Top Down Algorithms 64
iii

3.5 Simple /3-BBSTs 66
3.6 BBSTs without Deletion 70
3.7 Experimental Results 72
3.8 Conclusion 98
4 WEIGHT BIASED LEFTIST TREES AND MODIFIED SKIP LISTS 105
4.1 Introduction 105
4.2 Weight Biased Leftist Trees 106
4.3 Modified Skip Lists 109
4.4 MSLs As Priority Queues 118
4.5 Experimental Results For Priority Queues 122
4.6 Conclusion 129
5 CONCLUSIONS
135
A ABBREVIATIONS .
136
REFERENCES 138
BIOGRAPHICAL SKETCH 141
IV

LIST OF TABLES
2.1 Error rate (%) over optimal, / = 1 27
2.2 Improvement (%) over Fang, / = 1 27
2.3 Time taken, / = 1 28
2.4 Error rate (%) over optimal, 1 = 2 29
2.5 Improvement (%) over / = 1 cases 29
2.6 Improvement (%) over Fang, 1 = 2 30
2.7 Time taken, 1 = 2 31
3.1 The number of key comparisons on random inputs (version 1 code) . 76
3.2 The number of key comparisons on ordered inputs (version 1 code) . 77
3.3 Height of the trees on random inputs (version 1 code) 77
3.4 Height of the trees on ordered inputs (version 1 code) 77
3.5 The number of rotations on random inputs (version 1 code) 79
3.6 The number of rotations on ordered inputs (version 1 code) 80
3.7 Run time on random inputs using integer keys (version 1 code) .... 81
3.8 Run time on ordered inputs using integer keys (version 1 code) .... 82
3.9 Run time on random real inputs (version 1 code) 83
v

3.10 Run time on ordered real inputs (version 1 code) 84
3.11 The number of key comparisons on random inputs (version 1 code) . 87
3.12 The number of key comparisons on ordered inputs (version 1 code) . 88
3.13 Height of the trees on random inputs (version 1 code) 89
3.14 Height of the trees on ordered inputs (version 1 code) 89
3.15 The number of rotations on random inputs (version 1 code) 90
3.16 The number of rotations on ordered inputs (version 1 code) 91
3.17 Run time on random inputs using integer keys (version 1 code) .... 93
3.18 Run time on ordered inputs using integer keys (version 1 code) .... 94
3.19 Run time on random real inputs (version 1 code) 95
3.20 Run time on ordered real inputs (version 1 code) 96
3.21 The number of key comparisons on random inputs (version 2 code) . 99
3.22 The number of key comparisons on ordered inputs (version 2 code) . 100
3.23 Run time on random real inputs (version 2 code) 101
3.24 Run time on ordered real inputs (version 2 code) 103
4.1 The number of key comparisons 117
4.2 Number of levels 118
4.3 Run time 119
4.4 The number of key comparisons 124
4.5 Height/level of the structures 126
4.6 Run time using integer keys 127
4.7 The number of key comparisons 130
vi

4.8
Height/level of the structures
141
132
4.y
A.l
itun time using miegei Keys
137
vii

LIST OF FIGURES
2.1 Cell joining 5
2.2 Mayer river routing 8
2.3 Round robin and greedy layer assignments 11
2.4 Minimizing the number of tracks or layers 14
2.5 Constraint graph representation 18
2.6 Merge in constraint graph 19
2.7 Heuristic 1 20
2.8 Heuristic 2 22
2.9 Heuristic 3 24
3.1 LL and RL rotations 35
3.2 A tree in WB(l/4) that is not ^-balanced 42
3.3 —balanced tree that is not a COST 43
3.4 LL rotation for insertion 45
3.5 Substep (i) of insertion LR rotation 46
3.6 Case LL for LR(ii) rotation 49
3.7 Case LR for LR(ii) rotation 51
viii

3.8 LL rotation for deletion 58
3.9 LR rotation for deletion 60
3.10 Restructuring procedure 65
3.11 Simple restructuring procedure for insertion 68
3.12 Simple restructuring procedure for deletion 68
3.13 Simple restructuring procedure without a /3 value 70
3.14 Run time on real inputs (version 1 code) 85
3.15 Run time on random real inputs (version 1 code) 92
3.16 Run time on ordered real inputs (version 1 code) 97
3.17 Run time on random real inputs (version 2 code) 102
3.18 Run time on ordered real inputs (version 2 code) 102
4.1 Example min-VVBLTs 108
4.2 min-WBLT Insert 109
4.3 min-WBLT Delete-min 110
4.4 Skip Lists Ill
4.5 Modified Skip Lists 113
4.6 MSL Search 113
4.7 MSL Insert 114
4.8 MSL Delete 114
4.9 Run time 120
4.10 TMSL Insert 121
4.11 TMSL Delete-min 121
IX

4.12TMSL Delete-max
. 122
4.13 Run time on randoml 128
4.14 Run time on random2 129
4.15 Run time on randoml 131
4.16 Run time on random2 133
x

Abstract of Dissertation Presented to the Graduate School
of the University of Florida in Partial Fulfillment of the
Requirements for the Doctor of Philosophy
EFFICIENT ALGORITHMS AND DATA STRUCTURES FOR VLSI CAD
By
Seonghun Cho
May 1996
Chairman: Dr. Sartaj Sahni
Major Department: Computer and Information Science and Engineering
In this dissertation, we develop efficient algorithms and data structures for prob¬
lems that arise in electronic computer aided design (ECAD).
We consider the problem of joining a row of compacted cells so as to minimize the
area occupied by the cells and the interconnects. The cell joining process includes
cell stretching and river routing. We propose several heuristics to join a row of
cells in such a way that area is minimized. The proposed heuristics are compared
experimentally with the previously proposed heuristic.
We develop a new class of weight balanced binary search trees called /^-balanced
binary search trees (/3-BBSTs). /2-BBSTs are designed to have reduced internal path
xi

length. As a result, they are expected to exhibit good search time characteristics.
Individual search, insert, and delete operations in an n node /2-BBST take O(logn)
time for 0 < /? < y/2 — 1. Experimental results comparing the performance of /?-
BBSTs, WB(q) trees, AVL-trees, red/black trees, treaps, deterministic skip lists and
skip lists are presented. Two simplified versions of /TBBSTs are also developed.
We propose the weight biased leftist tree as an alternative to traditional leftist
trees for the representation of mergeable priority queues. A modified version of
skip lists that uses fixed size nodes is also proposed. Experimental results show
our modified skip list structure is faster than the original skip list structure for the
representation of dictionaries. Experimental results comparing weight biased leftist
trees and competing priority queue structures as well as experimental' results for
double ended priority queues are presented.
Xll

CHAPTER 1
INTRODUCTION
1.1 Background
In VLSI layout, we are concerned with transforming a circuit from its logical de¬
sign to a physical implementation. The layout problem for VLSI circuits is generally
decomposed into smaller problems such as partitioning, floorplanning, placement,
routing and compaction.
The partitioning process decomposes a large circuit/module into a collection of
smaller sub-circuits/modules. In floorplanning, logical components of a circuit are
assigned relative positions on a chip. The physical realization of each component (i.e.,
its area and aspect ratio) is also selected. The objectives of floorplanning include
overall area minimization, minimization of power consumption, etc. The precise
locations for the components of a design are then determined during the placement
process to optimize the area and the timing. After the components are placed, the
pins are connected during the routing process. During the process of compaction,
the components and interconnections are moved so as to further optimize the layout
in terms of area and delay.
The routing process is usually divided into three smaller sub-problems of global
routing, detailed routing, and specialized routing. Global routing decomposes the
1

2
complex routing problem into small and manageable sub-problems. It assigns each
net to a set of routing regions such as channel, switchbox, and planar routing to
minimize a combination of criteria such as area, circuit delay, etc. Steiner trees
and spanning trees are the commonly used approaches for net connection in global
routing. Specialized routing is used to connect power-ground nets or clock nets.
Detailed routing has two types of routing, general and restricted and there are three
types of detailed restricted routing, channel, switchbox, and planar.
In channel routing, all terminals of nets are located in two parallel rows across
a routing region called channel. In switchbox routing, terminals of nets are located
on the four sides of the routing region. Planar routing is a problem in which the
interconnection topology of the nets is planar. That is, all connections can be realized
on a single layer. Vias allow wires to change layers but the presence of vias reduces
reliability and performance of a circuit. Single layer routing is not always possible.
River routing is a special case of planar routing in which all nets have exactly
two terminals, one on each side of the channel, and the net sequence on each side
of the channel is the same. River routing is used in PCB routing, particularly in
dataflow architectures with multi-bit buses connecting a series of logic blocks, and
symbolic IC design systems.
1.2 Dissertation Outline
This dissertation is divided into four chapters. In Chapter 2, we consider the
problem of joining a row of compacted cells so as to minimize the area occupied by
the cells and the interconnects. The cell joining process includes cell stretching and

3
river routing. We propose several heuristics to join a row of cells in such a way that
area is minimized. The proposed heuristics are compared experimentally with the
previously proposed one.
VLSI Physical Design Automation is essentially the study of algorithms and data
structures related to the physical design process. Specific data structures can be used
to improve the performance of algorithms. For example, maze routing algorithms and
line-probe algorithms in global routing [30] use search structures, retiming algorithms
[6, 29] use priority queue structures, and layout compaction algorithms [13] use radix
priority search trees.
Chapters 3 and 4 are about search and priority queue structures. As specific
data structures could be used to produce better performance of algorithms in VLSI
design automation, new search and priority queue structures are proposed and thor¬
oughly compared with other data structures.
Finally, in the last chapter, we present conclusions of this work.

CHAPTER 2
MINIMUM AREA JOINING OF COMPACTED CELLS
2.1 Introduction
When designing circuits with compacted symbolic sticks basic cells, the circuit
is realized by a collection of compacted cells that tile a two-dimensional area. The
intercell interconnects are such that each interconnect connects two terminals that
are on adjacent boundaries of neighboring cells. So, for example, if cells A and B
(Figure 2.1(a)) are neighboring cells of the circuit, then the right boundary of A
is adjacent to the left boundary of B. The number of terminals on each of these
boundaries will be the same and the Pth terminal (from the bottom) on the right
boundary of A is to be connected to the i’th terminal (from the bottom) on the left
boundary of B.
Since the cells are available in compacted form, it is not possible to reduce the
distance between any pair of terminals on any side of a cell. However, this distance
can be increased by stretching the cell. In the example of Figure 2.1(a), we can stretch
either cell vertically by defining a horizontal cut line at any position and pulling the
two cell pieces apart by any desired amount (the cell can also be stretched horizontally
by using a vertical cut line).
4

5
(a)Horizontal adjacent cells
(b)Joining by stretching
(c)Joining by river routing (d) Combination cell joining
Figure 2.1. Cell joining
The required interconnects between cells A and B of Figure 2.1(a) can be ac¬
complished by stretching cells A and B so that the terminals of A and B line up as
in Figure 2.1(b). The broken lines in Figure 2.1(a) indicate the cut lines used for
stretching. The stretching enables us to join cells A and B using no routing tracks
(by “join” we mean make the interconnects between cells A and B). This method of
joining cells is also called pitch matching.
Another way to join cells A and B is to river route the interconnects as in
Figure 2.1(c). This uses routing tracks in a channel between cells A and B but does
not increase cell height. The pitch matching and river routing approaches to cell
joining have been studied in Boyer [5] and Weste [33]. Algorithms for single-layer

6
river routing can be found in several works [15, 19, 23. 24] and those for multilayer
river routing can be found in Baratz [3]. Single-layer gridless river routing is studied in
Tompa [32]. Two applications of river routing are hybrid circuit design and structured
design (DSP).
Cell stretching (or pitch matching) increases the height of the layout while river
routing increases its width. Both affect the layout area. The layout of Figure 2.1(b)
has area 150. To compute the area of the layout of Figure 2.1(c), we assume tracks
have unit separation. So, the layout width is 14 and height is 11. The layout has area
154. Cheng and Despain [S] have proposed using a combination of cell stretching and
river routing so as to obtain layouts with smaller area than possible when only one
of these joining methods is used. Figure 2.1(d) shows the result of joining cells A
and B using both stretching and river routing. The area of this layout is 144. This
is minimum for the instance of Figure 2.1(a).
Cheng and Despain [8] have proposed a heuristic for single layer joining of
compacted cells. At each step of their heuristic either a row or column of compacted
cells is joined. Following this, the row or column of joined cells is replaced by a
composite cell that represents the result of joining. Notice that when a row (column)
of cells is joined, cells may be stretched vertically (horizontally) and river routing is
done in a vertical (horizontal) channel. To join a row of cells, Cheng and Despain [8]
bound the maximum height to which a cell may be stretched. This bound is
h-rnax + hlvg/(4 * h max)

7
where hmax is the height of the tallest compacted cell being joined and havg is the
average height of the cells being joined.
Using this bound, cells are joined one-at-a-time using a penalty/reward scheme
to determine if a pair of terminals is to be joined by stretching or by river routing.
Lim, Cheng, and Sahni [17] have considered the case when only two cells are to
be joined. They develop fast polynomial time algorithms to obtain the minimum area
join of two cells. In addition, they are able to obtain, in low order polynomial time,
minimum area joins that minimize the length of the longest wire or the total wire
length. Lim [16] has proposed an 0(n(7z/c)c-1) algorithm to find the minimum area
join of c cells having a total of n terminals. This algorithm does an exhaustive search
over all possible numbers of tracks in the c — 1 routing channels between adjacent
cells. A constraint graph is used to determine the minimum height layout for each
assignment of number of tracks to routing channels. The time required per track
assignment is 0(n) and the worst case number of track assignments is 0((n/c)c_1).
The algorithm of Lim [16] is flawed as it handles channels with zero routing tracks
by joining the adjacent cells using minimum height cell stretching and then considers
the joined cells as one. This problem is easily fixed, however, by combining, in the
constraint graph, pairs of vertices that represent corresponding terminals of the two
cells (i.e., ¿’th terminals of each cell) with zero routing tracks in between.
In this chapter, wre consider the case when / > 1 routing layers are available to
river route the inter cell connections. Note that while multiple layers do not affect
layout area when cell stretching alone is used, a reduction in area is possible when

8
(a) 1-layer (b) 2-layer
Figure 2.2. /-layer river routing
cell stretching is combined with river routing or when river routing alone is used.
We assume that in each layer of each routing channel, the interconnects are to be
accomplished using river routing. An alternative is to use HV routing when / = 2,
HVH or VHV routing when / = 3, and extensions of HVH and VHV routing for
/ > 3. However, for river routing instances, using routing layers in this way has
no advantage over river routing in each layer (see Theorem 5, Section 2). When
the number of layers available for river routing is increased, one may see a dramatic
reduction in the number of routing tracks needed per layer. Figure 2.2 shows an
instance that needs n tracks when routed in one layer but only one track/layer when
routed in two layers.
We begin, in section 2, by stating the necessary and sufficient conditions for
a river routing instance to be routable in / layers using at most t tracks per layer
and stating how to perform /-layer river routing when such a routing is possible. In

9
this section, we also show that HV style routing has no advantage over river routing
in each layer. In Section 2.3, we describe the constraint graph used to determine
minimum height stretching of c cells. Heuristics for the minimum area joining of
c cells are proposed in Section 2.4 and the results of experiments with these are
provided in Section 2.5. Our conclusions appear in Section 2.6.
2.2 Mayer River Routing
Let (A{,B¡), 1 < t < m, be a set of terminal pairs such that the /4,’s are on
one side (say left or top) of a routing channel and the Bfs are on the other (right
or bottom) side. Terminal is to be connected to terminal B{, 1 < i < m. For
this channel routing instance to be an instance of river routing, it must be the case
that ai < a2 < ... < am and b\ < b2 < ... < bm where a, and 6,, respectively, give
the positions of terminals A, and J5,-, 1 < i < m. We may assume an underlying
grid with each terminal being at a grid position. In the case of a horizontal (vertical)
channel the g,s and 6,,s are grid column (row) numbers. Leiserson and Pinter [15] have
obtained the following necessary and sufficient condition for a river routing instance
to be routable in a single layer using at most t > 0 tracks.
Theorem 1 [15] The river routing instance defined above is routable in a single layer
using at most t > 0 tracks if and only if
(a) ai+t - bi>t
(b) b,+t - ai>t
for every i < m — t.

10
For the general case of / > 1 layers, we obtain the necessary and sufficient
condition of Theorem 2.
Theorem 2 The river routing instance defined above is routable in l > 1 layers (each
layer routing whole nets) using at most t > 0 tracks per layer if and only if
(a) ai+it bi >t
(b) bi+it — cti > t
for every i < m — It.
Proof: First, we establish that (a) and (b) are necessary for l layer routing. Since
the proofs for (a) and (b) are similar, we provide that for (a) only. Suppose that
a1+/t — bi < t for some i. Consider the It + 1 terminal pairs (Aj, Bj), i < j < i + It.
When routing these on / layers, at least one layer has to be assigned > t + 1 terminal
pairs. So, suppose that terminal pairs (A\,B[), (A'2,B'2), ..., (A't+l, B't+l), ... are
assigned to the same layer for river routin g. We may assume that a[ < a'2 < ... <
a'(+1 and b\ < b'2 < ... < 6'i+1. Since a't+1 < a,+/t and b\ > 6,, a't+1 — b[ < ai+¡t — 6, < t.
From Theorem 1, it follows that the terminal pairs (/!' , Z^), 1 < j < t + 1, cannot be
river routed on a single layer. Hence, (Aj, Bj), i < j < i + /¿, cannot be river routed
on / layers. So, (Aj, Bj), 1 < j < m, cannot be river routed on / layers. As a result,
(a) is a necessary condition.
To show that (a) and (b) are sufficient conditions for routability, we present two
algorithms (RoundRobin and Greedy) that assign the nets to layers in such a way
that each layer is river routable when both (a) and (b) are satisfied. The correctness

11
procedure RoundRobin ;
{ Assign the m nets to / layers. }
begin
for i := 1 to m do
assign net (A,, Bf) to layer (i mod /) + 1 ;
end ;
procedure Greedy ;
{ Assign the m nets to l layers. }
begin
for i := 1 to m do
assign net (Ai,B{) to layer q such that q is the smallest integer < /
for which the conditions of Theorem 1 axe not violated on layer q
(if there is no such q, then fail) ;
end ;
Figure 2.3. Round robin and greedy layer assignments
of these algorithms is established in Theorems 3 and 4, respectively. â–¡
We later discovered that Baratz [3] has not only obtained the same condition
but also proposed the same two algorithms for /-layer river routing. One assigns
nets to layers in a round robin fashion and the other uses a greedy strategy. The
corresponding procedures are given in Figure 2.3.
Theorem 3 The layer assignment produced by the RoundRobin procedure is river
routable if
(a) Oi+/í b¡ > t
(b) bi+it - a,>t
for all i < m — It.

12
Proof: Let (A[,B[), (A',^), • • • = (Aj,Bj), (Aj+l,Bj+l), (A>+2/, BJ+„), ... be the
nets assigned to layer (j mod /) + 1, j < l. So, a\ = a,+(,_|)/ and b[ =
Hence,
a', + t ~K~ aj + (i + t-l)l - fy + («-l)i
= aj+(i-i)i+ti - fy+(«-l)/
> t (from (a))
Similarly, b'i+t — a[ > t. So, the layer assignment satisfies the conditions of Theorem 1
and is river routable using t tracks. â–¡
Theorem i If (a) al+¡t - b, > t and (b) bl+¡t — ax>t for all i < m —It, then procedure
Greedy assigns nets to layers such that the assignment to each layer is routable using
t tracks.
Proof: If procedure Greedy is able to assign each of the m nets to a layer, then the
layer assignments satisfy the conditions of Theorem 1 and so are routable using t
tracks. Suppose the algorithm fails while trying to assign net (Ar, Br) to a layer. At
this time nets (A,, £?,), 1 < i < r, have been assigned to layers so as to satisfy the
conditions of Theorem 1 and the assignment of net (AT,BT) to each of these layers
violates these conditions. Consider first those layers, La, on which condition (a) is
violated. For a layer s € La, suppose that the assigned nets are ..., (A'_t,B':_t),
..., (A'j_x,B'j_j). Let (Aj,Bj) = (Ar,Br). Since s € La, we have a'- — 6'_t < t.
Now, if 6r_/( > 6'_(, then aT — 6r_/t < t which violates condition (a) of this theorem.
So, 6r_ít < Since, ór_í£ < b'-_t < ... < &'_2 < ,, t of the It - 1 nets

13
(Ar-tt+\, Br-it+1(i4r_i, 5r_!) have been assigned to layer s £ La. Consequently,
the layers in La account for t\La\ of these It — 1 nets.
In a similar way, we can show that the remaining / — \La\ layers account for
another t(l — |L0|) of these nets. This gives us a total of tl nets, whereas we had only
tl— 1. This contradiction implies that procedure Greedy cannot fail unless conditions
(a) and (b) are not satisfied. Q
Procedure RoundRobin is easily seen to have complexity of 0(m). A straightfor¬
ward implementation of procedure Greedy will have complexity of 0(ml). However,
by using priority search trees [18] the complexity can be reduced to O(mlog/). In
practice, since / is quite small, it is unlikely that the priority search tree implemen¬
tation will run faster than the straightforward implementation in which the / layers
are checked in sequence. The actual routing for all l layers can be done in 0(mt)
time using the computed layer assignment and the single layer routing algorithm of
Leiserson and Pinter [15].
Using Theorem 2, we can develop a linear time algorithm to determine the
minimum number of tracks needed to route an instance in / layers as well as to find
the minimum number of layers needed for a t track routing. The algorithm for the
former is given in Figure 2.4. This figure also shows the changes needed in case t is
given and we wish to determine the minimum number of layers. The correctness of
the algorithm follows from that of Theorem 2 and the fact that t (or /) is increased
only if the current t (/) is found to be infeasible. The complexity is 0(m) as neither

14
procedure MinimizeTracks ; {or MinimizeLayers}
{ Determine the minimum number of tracks per layer (or minimum number
of layers) needed for multilayer river routing }
begin
t 0 ; {or / := 1}
* := 1 ;
while (i < m — It) do
if (ai+u - b{ < t) or (6,+/t - a, < t)
then t := t -f 1 {or / := / + 1}
else i := t + 1 ;
end ;
Figure 2.4. Minimizing the number of tracks or layers
i nor t (/) can exceed m. So, neither clause of the if statement can be executed more
than m — 1 times.
Using the multilayer river routing results of Baratz [3], one can trivially extend
all the results of Lim, Cheng and Sahni [17] to the case of multilayer joining of
compacted cells. So, the multilayer minimum area join of two compacted cells with
m nets can be obtained in 0(m2) time. If we wish to minimize the maximum wire
length while keeping area minimum, the asymptotic time complexity is still 0(m2).
The total wire length can be minimized while keeping area minimum in O(m2logm)
time.
In HV style routing, each routing layer is assigned a routing direction (either
H or V). In an H (V) layer only horizontal (vertical) wire segments can be laid out.
Horizontal segments on one layer connect to vertical segments, of the same net, on
another layer by means of vias. In the case of river routing instances, one can see
that there is no advantage to having more than two V-layers (i.e., two V-layers are
sufficient to route all river routing instances).

15
Let RR(/,¿) be the set of all river routing instances that can be routed in /
layers, using t tracks per layer and using river routing in each layer. Let HV(/, t) be
all river routing instances that can be routed using HV style routing, l layers, and
t tracks per layer. Note that HV(/, t) includes instances routable with 0, 1, and 2
V-layers. Let HVV(/, t) be all river routing instances using HV style routing, / — 2
H-layers, and 2 V-layers. Theorems 5 and 6 below hold for both the knock-knee [25]
and directional HV models. Theorem 7 holds only for the directional model.
Theorem 5 HV(/, t) C RR(/,<) for every l > 1 and every t > 1.
Proof: HV(/,í) C RR(l,t) follows from a more general result obtained by Baratz [3].
Baratz [3] has shown that, for river routing instances, there is no advantage to using
any routing scheme that wires a net on more than one layer. Since it is easy to
construct river routing instances X such that X £ RR(/,f) and X ^ HV(/, <), it
follows that HV(/, t) C RR(/,<).
We provide a simpler proof of HV(/, t) C RR(/, t). This proof will also establish
our next result. We shall show that if X is a river routing instance such that X ^
RR(/,i), then X £ HV(/,0- Hence, HV(/,t) C RR(/,<).
Suppose that X ^ RR(/, t). From Theorem 2, it follows that a,+/t — 6, < t or
bl+it — a,- < t for some i. Suppose that a,-+/< — 6, < t (the proof is similar when
bi+n — a, < t). So, a,+/t < b, -f t. Since X is a river routing instance, at least nets
i + f,..., i + It intersect a vertical cut line drawn at a,+/t. Hence, the density of X
at a,+it is > i -f It — (i + t) T 1 = (/ — l)t -f 1. When HV style routing is used with

16
/, / > 1, layers, at most / — 1 layers are available for horizontal routes. With t tracks
per layer, densities of at most (/ — 1)/ can be accommodated. So, X HV(/,t). □
Theorem 6 HVV(/, t) C RR(/ — 1, i) for every l > 2 and every t > 1.
Proof: As in Theorem 5, suppose that X (f RR(/— 1, t). Let i be such that —
b, < t. The net density at is > (/ - 2)t + 1. In HVV routing, two layers
are V-layers. So only 1 — 2 layers are available for horizontal segments. This is not
enough as the horizontal segment density is > (/ — 2)t + 1 at a,+(/_i)t. Hence, X $
HVV(M).
One may easily construct river routing instances that are in RR(/ — 1, t) but not
in HVV(/,i). â–¡
Theorem 7 RR(2, t) (f HV(/, t) - HVV(/, t) for every l > 1 and every t > 1.
Proof: Consider the RR instance (ai,6i) = (1,2) and (a2,¿>2) = (2,3). This is in
RR(2,t) for every t > 1 but is not in HV(/, t) — HVV(/,t) for any /. □
As remarked earlier, Theorem 7 holds only for the directional model. For the
knock-knee model, one can show that RR(/,t) C HV(/ + l,i) for every / > 1 and
every t > 1.
2.3 Constraint Graph Representation
Lim [16] has proposed the use of a constraint graph to determine the terminal
positions in a row of compacted cells. This is for the case when the number of tracks

17
in each routing channel is given and we wish to minimize the layout height. In the
constraint graph, each cell is represented by a directed chain of vertices. Each cell
terminal is represented by a vertex. The exception is when a compacted cell has
terminals at the same ^-position on both sides of the cell. In this case, the two
terminals at the same y-position are represented by a single vertex. The vertex chain
is linked in the direction of increasing y-position. The chain edges are labeled by the
minimum allowable terminal separation. In addition, the constraint graph contains
a source vertex that represents the bottom of the layout and a sink vertex that
represents the layout top. The source vertex connects to the bottom of each chain
and the top of each chain is connected to the sink vertex.
Figure 2.5(b) shows the chains (solid edges) for the four cell row of Figure 2.5(a).
To complete the constraint graph, directed edges are added to introduce the channel
routing constraints of Theorem 2. These are represented by the broken edges of
Figure 2.5(b). Figure 2.5(b) is for the two layer case.
Lim [16] has shown that the constraint graph is acyclic provided the number
of tracks in each routing channel is > 0. He has proposed handling channels with
zero tracks by finding first the minimum area joining of the adjacent cells (only cell
stretching is permitted now) and then combining these two cells into one. I.e., the two
cells are replaced by their minimum area join. This strategy can be shown to result
in non-optimality of the algorithm proposed in Lim [16]. To preserve optimality, it
is necessary to merge the vertices that represent terminals that are the endpoints of

18
sink
Figure 2.5. Constraint graph representation
nets that are to be routed using no tracks as in Figure 2.6. The resultant constraint
graph is also acyclic.
It is easy to see that the number of vertices and edges in the constraint graph
is 0(n) where n is the total number of terminals. Furthermore, the graph can be
constructed in 0(n) time given the number of routing layers and the number of tracks
in each channel. The constraint graph described by us is identical to that of Lim [16]
except in the way channels with zero tracks are handled and in that our graph is
defined for / > 1 routing layers while that of Lim [16] is only for / = 1.

19
sink
The length of the longest path from the source vertex of the constraint graph
to each of the remaining vertices can be computed in O(n) time by doing this in
topological order [14, Section 6.5]. It is easy to see that if each terminal is placed
at a vertical position given by the longest path length from the source, then all nets
can be routed in the given number of tracks (as the conditions of Theorem 2 are
satisfied in each routing channel). Furthermore, Lim [16] has shown that such a
positioning of terminals results in a stretched layout of minimum height for the given
channel widths. As a result, when channel widths are known, cells can be stretched
to minimize area in 0(n) time. The channel widths that result in minimum area can
be determined in 0(n(n/c)c-1) time where c is the number of cells by trying out all

20
procedure Heuristicl ;
begin
for i := 1 to c — 1 do
begin
determine the minimum area join of each pair of adjacent cells ;
select the pair that has minimum area and replace it with its
minimum area join ;
end ;
end ;
Figure 2.7. Heuristic 1
possible channel widths [16]. Since this is feasible only for small c, we propose several
heuristics in the next section.
2.4 Heuristics to Minimize Area
We formulate three greedy heuristics to obtain the minimum area join of a row
of c compacted cells that have a total of n terminals.
2.4.1 Heuristic 1
The heuristic is described in Figure 2.7.
In each iteration of the for loop we examine every pair of adjacent cells. For
each pair, the minimum area join is found using the algorithm of Lim, Cheng and
Sahni [17] extended to the multilayer case as discussed in Section 2.2. The pair which
has the minimum area join is replaced by a single cell that represents this join. So
following each iteration of the for loop the number of cells decreases by one. When
the for loop terminates, we are left with a single cell that represents the join of
all c cells. The time needed to determine the minimum area join of a pair of cells
with n, nets between them is 0(n?). The time to do this for all pairs of adjacent

21
cells is 0(£,n2) = 0(n2). So, the for loop iteration with i = 1 takes 0(n2) time.
On subsequent iterations, only the two pairs that include the cell introduced in the
previous iteration need to have their minimum area join computed. Since each cell
pair being considered includes at least one composite cell, the minimum area join
is computed by considering the portion of the constraint graph that represents all
the basic cells in the cell pair. Channel widths for channels within a composite cell
are not changed while obtaining the minimum area join of the cell pair. However,
as different channel widths for the channel between the two (composite) cells being
joined are tried, the constraint graph is used to determine the minimum height of the
combined cell. So, the time to combine two (composite) cells with n, terminals in the
channel between them is 0(nn,). Hence the time for the remaining c — 2 iterations
is 0(n£, n,) = 0(n2). The overall complexity of Heuristic 1 is therefore 0(n2). In
case the terminals are uniformly distributed over the cells, n, = 0(n/c) for all i. The
time for the first iteration of the for loop is now 0(n2/c) and that for each of the
remaining iterations is 0(n2/c). The overall time is 0(n2).
2.4.2 Heuristic 2
In this heuristic, we begin by assigning each channel the number of tracks needed
to route the channel with no cell stretching. This number can be determined in 0(n,)
time for a channel with n, nets as described in Section 2.2. The time taken to do
this for all c — 1 channels is 0(n). The configuration obtained in this way is the
maximum width layout. Starting from this configuration, we reduce the total number
of tracks available across all c — 1 channels by one on each iteration. For this, the

22
procedure Heuristic2 ;
begin
for each channel determine the number of tracks,
needed to route with no stretching, 1 < i < c ;
t := Ed U ;
set up the constraint graph using t¿ tracks in channel i, 1 < i < c ;
compute layout area, A ;
for tracks := t downto 1 do {reduce by 1}
begin
for i := 1 to c — 1 do
begin
reduce the number of tracks in channel i by 1 ;
modify the constraint graph to reflect this ;
determine the length of the longest path in the graph
and from this the layout area, a, ;
end ;
select j such that aj — min{ a, } ;
reduce the number of tracks in channel j by 1 ;
A = min{ A,aj } ;
end ;
end ;
Figure 2.8. Heuristic 2
effect of a one track reduction is computed for each channel. The minimum layout
height is determined by computing the length of the longest path in the constraint
graph of Section 2.3. The track reduction is done in the channel that results in the
smallest layout height (hence the minimum area for the given number of tracks). The
algorithm is stated more formally in Figure 2.8.
When the algorithm terminates, A is the area of the minimum area join found
by the heuristic. To reconstruct the layout, it is necessary to store the tracks per
channel each time A is updated in the statement A = mm{ A, a} }.
For the time complexity, we see that the steps that precede the outer for loop
take 0(n) time. Each iteration of the outer loop takes O(nc) time. Hence this loop

23
contributes a total of 0(net) to the time. Since t = 0(n), the overall time complexity
of Heuristic 2 is 0(n2c).
2.4.3 Heuristic 3
Unlike Heuristic 2 which attempts to minimize the layout height for each value
of t, the total number of tracks, Heuristic 3 attempts to minimize the width (i.e.,
total number of tracks) for each choice of layout height. The heuristic begins with a
layout height, ht, equal to the height of the tallest compacted cell. At each iteration,
the next layout height to use is computed as described later. During each iteration,
cells are combined in groups of at most k (k > 1 is a parameter to the heuristic).
Each group of combined cells is replaced by its minimum area join subject to the
constraint that the height of the join does not exceed ht. This joining of < k cells at
a time continues until only one cell remains. Its area is computed and recorded. The
minimum area obtained over all heights tried is then reported as the best. Heuristic
3 is given in Figure 2.9.
In our implementation of Heuristic 3, the minimum area join of k cells is found
by considering the portion of the constraint graph for all the basic cells included in
these k cells. So, for this purpose composite cells are not handled as single cells.
Rather, as in Heuristic 1, the basic cells they are composed of are considered and
channel widths previously assigned to the associated channels are not changed. Track
assignment is done only for the k — 1 channels between the k composite cells. We
found this to give better results than when composite cells were regarded as atomic.
For the case k = 2, the minimum area is determined by a binary search over the

24
procedure Heuristic3 ;
begin
ht := height of the tallest cell ;
repeat { minimize width subject to height < ht }
repeat { do this by combining k cells at a time }
select k adjacent cells such that the minimum height cell is selected
and the height of the tallest selected cell is minimum
(if there are fewer than k cells, then select all of them) ;
obtain the minimum area layout for the selected cells under
the constraint that the layout height does not exceed ht ;
during the preceding step record the next value of ht
that is possible for a layout ;
until one cell remains ;
compute the area of the remaining cell and record it
if it is less than the minimum area found so far ;
if there is no next height then terminate ;
ht := next height ;
until false ;
end ;
Figure 2.9. Heuristic 3
number of tracks in the single channel. This takes O(nlogn,) time where n, is the
number of nets in channel i. Thus the time needed for the inner repeat loop when
k — 2 is O(cnlogn) (for uniform terminal distribution it is 0(cn log(n/c)). During
the binary search, the heights corresponding to channel widths that require height
> ht are recorded. The minimum of these heights yields the next value of ht.
When k > 2, all track combinations for the k — 1 channels are tried as in
Section 2.3. Again, each composite cell is broken up into its basic cells. As different
track combinations are tried, we record the minimum height > ht that results from
any track combination. This gives the next value of ht. The time for the inner
repeat loop is 0((c/(/r — l))n(n/¿)Á:“1) (or 0((c/(k — l))n(n/c)fc_1) when terminals
are uniformly distributed).

25
In all our experiments, the outer repeat loop was iterated fewer than (k — l)n
times. To ensure that the number of iterations is 0(kn), one may adopt the following
scheme. When the number of iterations first reaches (k — l)n, compute a set of at
most n new heights by beginning with the current constraint graph. This uses the
current assignment of number of tracks in each channel. Heuristic 2 is next used to
reduce the total number of available tracks by one and determine the height needed
to complete the routing with the reduced number of tracks. This process gives us
at most n new heights hi < h2 < ... < hp. Heuristic 3 is now resumed with hi as
the next height. Only two iterations are performed. Then Heuristic 3 is resumed
with max{ h2, ht } as the next height. Again two iterations of the outer repeat loop
are done. Next the heuristic is resumed with max{ h$,ht } as the next height. This
continues until we have gone through p resumptions of the heuristic. With this scheme
to limit the number of iterations, the complexity of Heuristic 3 becomes 0(cn2 log n)
when k — 2 and 0((c/(k — 1 ))fcn2(n/¿)*-1) = 0(cn2(n/¿)k_1) when k > 2. For the
case when the n terminals are uniformly distributed over the c cells, the complexity
is 0(cn2 log(n/c)) when k = 2 and 0((c/(k— 1 ))kn(n/c)k~1) = 0(cn(n/c)*-1) when
k > 2. One may verify that since Heuristic 3 tries the maximum useful height (i.e.,
the height needed when no routing tracks are available), it generates optimal solutions
when k = c.
2.5 Experimental Results
We programmed our three heuristics as well as the heuristic Fang [8] in C and
ran tests on a single KSR processor. Optimal solutions for instances with up to nine

26
cells were obtained using the corrected version of the exhaustive search algorithm of
Lim [16]. Our test set consisted of instances that had a number of cells, c, equal to
one of the numbers in the set {3, ..9, 10, 20, 50, 100}. For each value of c, there
were twenty instances and the results were averaged over these instances. An instance
with c cells had c — 1 routing channels. The number, i, of terminals on either side of
each routing channel was equal to c for 3 < c < 9 and was 10 for the other values of
c. In addition, when c = 100, we also had instances with 20 terminals on either side.
In our experiments, we considered only single layer and two layer routing.
Table 2.1 gives the average percentage by which the area of the single layer
solutions generated by each of the heuristics exceeded the area of the single layer
optimal solution. As is evident, each of the heuristics proposed in this chapter gave
noticeably better solutions than did Fang. This table is only for the cases 3 < c < 9
as for c > 9 the optimal algorithm of Lim [16] required too much time to complete.
When k > c, Heuristic 3 is guaranteed to generate an optimal solution. So, we did
not run these cases.
In table 2.2, we have used the single layer solution produced by Fang as the
benchmark against which the solutions obtained by our three heuristics are compared.
This table gives the average percentage by which the area of the solutions produced
by our heuristics is less than that of the solutions produced by Fang. Our solutions
have area 9 to 18% less.
Table 2.3 compares the computing time requirements of the various algorithms
for the case of one layer. The optimal algorithm is useful only for small values of c

Table 2.1. Error rate (%) over optimal, / = 1
t = number of terminals on each side of each routing channel
* : k > c
Table 2.2. Improvement (%) over Fang, / = 1
cells
t
Heuristic
1
Heuristic
2
Heuristic3
Jfc = 2
k = 3
Ar = 4
10
10
9.4
14.0
14.0
14.2
14.3
20
10
10.5
15.4
15.4
15.6
15.6
50
10
10.6
16.1
16.1
16.2
-
100
10
9.1
16.5
16.5
16.6
-
100
20
9.3
18.3
18.0
-
-
: excessive run time

28
Table 2.3. Time taken, l — \
cells
t
Fang
Heuristic
1
Heuristic
2
Heuristic3
Optimal
k = 2
ifc = 3
II
3
3
0.0
0.01
0.02
0.02
*
*
0.01
4
4
0.0
0.02
0.02
0.04
0.09
*
0.05
5
5
0.0
0.03
0.04
0.09
0.36
1.2
0.77
6
6
0.0
0.03
0.08
0.21
0.66
3.7
13
7
7
0.0
0.04
0.14
0.50
3.4
27
278
8
8
0.0
0.07
0.24
0.84
4.4
34
1.8*
9
9
0.0
0.09
0.42
1.8
16
72.3
47*
10
10
0.0
0.14
0.78
3.1
19
334
-
20
10
0.01
0.32
5.8
16
117
1196
-
50
10
0.01
1.4
93
127
896
-
-
100
10
0.03
4.4.
782
590
4873
-
-
100
20
0.06
11
3022
4087
-
-
-
Times are in seconds,
f : Times are in hours.
(say up to 7). While Fang is significantly faster than the heuristics proposed here,
the quality of the solutions generated by our heuristics is superior.
Table 2.4 is the analog of table 2.1 for the case of two layers. Again, our heuristics
performed considerably better than did Fang. Table 2.5 gives the improvement in
area due to increasing the number of routing layers from one to two. This is influenced
somewhat by the width of cells which in our case ranged from 5 to 30 times the track
separation. With narrower cells, the impact of the second layer would have been
greater and with wider cells, it would have been less. Also, the impact of the second
layer is more when more routing tracks are needed. For the smaller instances of
table 2.4, for example, the optimal solutions with / = 2 required, on average, only
1.8% less area than when / = 1.

29
Table 2.5. Improvement (%) over / = 1 cases
cells
t
Fang
Heuristic
1
Heuristic
2
Heuristic3
k = 2
ib = 3
k = 4
10
10
4.1
4.7
3.3
3.3
3.0
3.0
20
10
4.1
5.8
3.4
3.4
3.3
3.2
50
10
4.6
5.2
3.4
3.2
3.2
-
100
10
4.7
6.3
3.5
3.4
3.4
-
100
20
7.1
10.4
4.8
5.2
-
—

30
Table 2.6. Improvement (%) over Fang, / = 2
cells
t
Heuristic
1
Heuristic
2
Heuristic3
k = 2
k = 3
k = 4
10
10
10.0
13.3
13.2
13.2
13.3
20
10
12.2
14.8
14.8
14.9
14.8
50
10
11.2
15.1
14.9
15.0
-
100
10
10.7
15.4
15.4
15.4
-
100
20
12.5
16.3
16.3
-
-
Table 2.6 is the analog of table 2.2 for the case of two layers. The results are
similar to those in table 2.2. Table 2.7 gives the average computing times for the two
layer instances. These are less than for the one layer case as the constraint graph has
fewer edges.
For large c, we recommend the use of heuristic 2 or 3 (with k = 2) and for small
c we recommend using heuristic 3 (with i = 3 or 4).
2.6 Conclusion
We have considered the problem of joining a row of compacted cells and de¬
veloped heuristics to stretch cells and river-route the nets so that the layout area is
minimized. Our proposed heuristic was compared, experimentally, with Fang [8] and
found to produce layouts with less area. However, Fang is faster. We recommend the
use of our Heuristic 3 with k = 3 or 4 in practice.

31
Table 2.7. Time taken, / = 2
cells
t
Fang
Heuristic
1
Heuristic
2
Heuristic3
Optimal
k = 2
k = 3
= 4
3
3
0.0
0.01
0.02
0.02
*
*
0.0
4
4
0.0
0.02
0.02
0.03
0.06
*
0.04
5
5
0.0
0.02
0.02
0.07
0.25
0.75
0.60
6
6
0.0
0.02
0.04
0.13
0.38
2.0
11
7
7
0.0
0.04
0.06
0.29
1.9
14
223
8
8
0.0
0.06
0.12
0.53
2.5
19
1.5*
9
9
0.0
0.06
0.20
0.96
7.5
34
39*
10
10
0.01
0.09
0.32
1.7
9.9
153
-
20
10
0.01
0.24
2.7
9.9
65
651
-
50
10
0.01
1.1
42
85
586
-
-
100
10
0.02
3.7
357
472
3069
-
-
100
20
0.05
7.9
1377
3166
—
—
—
Times are in seconds,
f : Times are in hours.

CHAPTER 3
A NEW WEIGHT BALANCED BINARY SEARCH TREE
3.1 Introduction
A dictionary is a set of elements on which the operations of search, insert, and
delete are performed. Many data structures have been proposed for the efficient
representation of a dictionary [14]. These include direct addressing schemes such as
hash tables and comparison schemes such as binary search trees, AVL-trees, red/black
trees [12], trees of bounded balance [21], treaps [1], deterministic skip lists [20], and
skip lists [26]. Of these schemes, AVL-trees, red/black trees, and trees of bounded
balance (WB(a)) are balanced binary search trees. When representing a dictionary
with n elements, using one of these schemes, the corresponding binary search tree
has height O(logn) and individual search, insert, and delete operations take O(logn)
time. When (unbalanced) binary search trees, treaps, or skip lists are used, each
operation has an expected complexity of O(log n) but the worst case complexity is
0(n). When hash tables are used, the expected complexity is 0(1) per operation.
However, the worst case complexity is 0(n). So, in applications where a worst case
complexity guarantee is critical, one of the balanced binary search tree schemes is to
be performed.
32

33
In this chapter, we develop a new balanced binary search tree called /?-BBST (/?-
balanced binary search tree). Like WB(a) trees, this achieves balancing by controlling
the relative number of nodes in each subtree. However, unlike WB(o) trees, during
insert and delete operations, rotations are performed along the search path whenever
they reduce the internal path length of the tree (rather than only when a subtree is
out of balance). As a result, the constructed trees are expected to have a smaller
internal path length than the corresponding WB(a) tree. Since the average search
time is closely related to the internal path length, the time need to search in a /3-BBST
is expected to be less than that in a WB(a) tree.
In Section 3.2, we define the total search cost of a binary search tree and show
that the rebalancing rotations performed in AVL and red/black trees might increase
this metric. We also show that while similar rotations in WB(a) trees do not increase
this metric, insert and delete operations in WB(o) trees do not avail of all oppor¬
tunities to reduce the metric. In Section 3.3, we define /3-BBSTs and show their
relationship to WB(a) trees. Search, insert, and delete algorithms for /?-BBSTs are
developed in Section 3.4. A simplified version of /3-BBSTs is developed in Section 3.5.
Search, insert and delete operations for this version also take O(logn) time each. An
even simpler version of /?-BBSTs is developed in Section 3.6. For this version, we
show that the average cost of an insert and search operation is O(log n) provided no
deletes are performed.
An experimental evaluation of /?-BBSTs and competing schemes for dictionaries
(AVL, red/black, skip lists, etc.) was done and the results of this are presented in

34
Section 3.7. This section also compares the relative performance of /?-BBSTs and the
two simplified versions of Sections 3.5 and 3.6.
3.2 Balanced Trees and Rotations
Following an insert or delete operation in a balanced binary search tree (e.g.,
AVL, red/black, WB(ct), etc.), it may be necessary to perform rotations to restore
balance. The rotations are classified as LL, RR, LR, and RL [14]. LL and RR
rotations as well as LR and RL rotations are symmetric. While the conditions under
which the rotations are performed vary with the class of balanced tree considered,
the node movement patterns are the same. Figure 3.1 shows the transformation
performed by an LL and an LR rotation. In this figure, nodes whose subtrees have
changed as a result of the rotation are designated by a prime. So, p' is the original
node p however its subtrees are different.
Let h(x) be the height of the subtree with root x. Let s(x) be the number of
nodes in this subtree. When searching for an element x, x is compared with one
element at each of /(x) levels, where /(x) is the level at which x is present (the root
is at level 1). So, one measure of the “goodness” of the binary search tree, T, for
search operations (assuming each element is searched for with equal probability) is
C(T) = •£ '(*)•
x£T
its total search cost defined as:

35
(a) LL rotation
(b) LR rotation
Figure 3.1. LL and RL rotations
Notice that C(T) = I(T) + n where I(T) is the internal path length of T and
n is the number of elements/nodes in T. The cost of unsuccessful searches is equal
to the external path length E(T). Since E(T) = I(T) + 2n, minimizing C(T) also
minimizes E(T).
Total search cost is important as this is the dominant operation in a dictionary
(note that insert can be modeled as an unsuccessful search followed by the insertion
of a node at the point where the search terminated and deletion can be modeled by

36
a successful search followed by a physical deletion; both operations are then followed
by a rebalancing/restructuring step).
Observe that in an actual implementation of the search operation in program¬
ming languages such as C++, C, and Pascal, the search for an x at level /(x) will
involve upto two comparisons at levels 1,2,..., /(x). If the code first checks x = e,
where e, is the element at level i to be compared and then x < e, to decide whether to
move to the left or right subtree, then the number of element comparisons is exactly
2/(x) — 1. In this case, the total number of element comparisons is
NC(T) = 2 l(x) ~ n = 2C(r) - n
xeT
and minimizing C(T) also minimizes NC(T). If the code first checks x < e, and then
x = e, (or > et), the number of element compaxisons done to find x is /(x) + r(x) + 1
where r(x) is the number of right branches on the path from the root to x. The
total number of comparisons is bounded by 2C(T’). For simplicity, we use C(T) to
motivate our data structure.
In an AVL tree, when an LL rotation is performed, h(q) = /i(c) +1 = /i(d) +1 (see
Figure 3.1(a)). At this time, the balance factor at gp is h(p) — h(d) = 2. The rotation
restores height balance which is necessary to guarantee O(logn) search, insert, delete
operations in an n node AVL tree. The rotation may, however, increase the total
search cost. To see this, notice that an LL rotation affects the level numbers of only
those nodes that are in the subtree with root gp prior to the rotation. We see that
l(q') = l{
37
with root a is decreased by s(a) as a result of the rotation, etc. Hence, the increase
in C(T) due to the rotation is:
Kp) - Kp) + *W) -K = -1 - 1 + 1 - s(q) + 1 + s(d) = s(d) - s(g).
A similar analysis shows that an LR rotation increases C(T) by s(d) — s(q).
If the LL rotation was triggered by an insertion, s(q) is at least one more than the
minimum number of nodes in an AVL tree of height t = h(q) — 1. So, s(q) > t+2/>/5
where = (1 + \/5)/2. The maximum value for s(d) is 2J — 1. So, an LL rotation
has the potential of increasing total search cost by as much as
2‘ - 1 - i+2/^5 « 2‘ - 1 - 1.62t+2/2.24.
This is negative for t < 2 and positive for t > 2. When t = 10, for example, an
LL rotation may increase total search cost by as much as 877. As t gets larger, the
potential increase in search cost gets much greater. This analysis is easily extended
to the remaining rotations and also to red/black trees.
Definition (WB(a) [21]) The balance, B(p), of a node p in a binary tree is the
ratio (s(/) + l)/(s(p) + 1) where / is the left child of p. For a € [0,1/2], a binary tree
T is in WB(a) iff a < B{p) < 1 — a for every node p in T. By definition, the empty
tree is in WB(a) for all a.

38
Lemma 1 (1) The maximum height, hmax(n), of an n node tree in WB(a) is ~
log_!_(n + 1) [21]
1 —Q
(2) Inserts and deletes can be performed in an n node tree in WB(a) in O(logn) time
for 2/11 < a < 1 -y/2/2 [4].
(3) Each search operation in an n node tree in WB(a) takes 0(\ogn) time [21].
In the case of weight balanced trees WB(a), an LL rotation is performed when
B(gp) «1 — a and B(p) > a/( 1 — a) (see Figure 3.1(a)) [21]. So,
1 aro *(p) + 1 = s^ + 1
~ s(gp) + 1 s(p) + s(d) + 2
or
s{d) « s(p)- +
1 — a
2a - 1
1 — a
and
a
1 — a
< B(p) =
s(q) +1
s{p) +1
or
s(q) >
1 — a
+
2a - 1
1 — a
So, LL rotations (and also RR) do not increase the search cost. For LR rotations
[21], B(gp) « 1 - a and B(p) < a/( 1 — a). So, s(d) m s(p)j^ + ^7 and with
respect to Figure 3.1(b),
a
> B(P) =
s(p) ~ ¿(g)
s(p) + 1
1 — a

39
or
1 — 2a a
5(?) > s{p)- — - -■
1 — a 1 — a
For a < 1/3, s(q) > s(d) and LR (RL) rotations do not increase search cost. Thus,
in the case of WB(a) trees, the rebalancing rotations do not increase search cost.
This statement remains true if the conditions for LL and LR rotation are changed to
those in Blum and Mehlhorn [4].
While rotations do not increase the search cost of WB(a) trees, these trees
miss performing some rotations that would reduce search cost. For example, it is
possible to have a < B(gp) < 1 - a, B(p) > y£^, and s(q) > s(d). Since B(gp) isn’t
high enough, an LL rotation isn’t performed. Yet, performing such a rotation would
reduce search cost.
3.3 ff-BBSTs
Definition A cost optimized search tree (COST) is a binary search tree whose search
cost cannot be reduced by performing a single LL, RR, LR, or RL rotation.
Theorem 8 If T is a COST with n nodes, its height is at most log¿(\/5(« + 1)) — 2.
Proof Let Nh be the minimum number of nodes in a COST of height h. Clearly,
A^o = 0 and N\ = 1. Consider a COST Q of height h > 2 having the minimum
number of nodes Nh- Q has one subtree R whose height is A — 1 and another, S,
whose height is < h — 1. R must be a minimal COST of height h — 1 and so has
Nh~i nodes. R, in return, must have one subtree, U, of height h — 2 and another, V,
of height < h — 2. Both U and V are COSTs as R is a COST. Since R is a minimal

40
COST, U is a minimal COST of height h - 2 and so has 1Vfc_2 nodes. Since Q is a
COST, |S| > max{\U\, |V'|}. We may assume that Nh is a nondecreasing function of
h. So, |5| > Nh-2- Since Q is a minimal COST of height h, |S| = N^-2- So,
Nh = Nk.i + Nh-2 + 1, h > 2
N0 = 0, Ni = 1.
This recurrence is the same as that for the minimum number of nodes in an AVL tree
of height h. So, = Fh+2 ~ 1 where F, is the *’th Fibbonacci number. Consequently,
Nh « (f>h+2/V5 — 1 and h < log¿(\/5(« + 1)) — 2. □
Corollary 1 The maximum height of a COST with n nodes is the same as that of an
AVL tree with this many nodes.
Definition Let a and b be the root of two binary trees, a and b are /3-balanced,
0 < 0 < 1, with respect to one another, denoted /3-(a,6), iff
(a) /?(s(a) - 1) < s(6)
(b) 0(s{b) - 1) < s(a)
A binary tree T is ^-balanced iff the children of every node in T are /^-balanced.
A full binary tree is 1-balanced and a binary tree whose height equals its size
(i.e., number of nodes) is 0-balanced.

41
Lemma 2 If the binary tree T is ¡3-balanced, then it is -f-balanced for 0 < 7 < p.
Proof Follows from the definition of balance. â–¡
Lemma 3 If the binary tree T is /?-balanced, 0 < ft < 1/2, then it is in WB(a) for
a = 0/(1 +ft).
Proof Consider any node p in T. Let l and r be node p's left and right children.
B(p) =
*{l) + 1
s(l) + $(r) + 2
1
1 +
J(r)+1 '
,(/) + !
Since T is /9-balanced, s(l) — 1 < s(r)//3 or s(l) -f 1 < s(r)//3 + 2. So,
*(l) + 1
s(r) + 1
2/3-1
f3(s(r) + 1)
or
•s(r) + 1
40+T
>/?•
So, B(p) < 1/(1 + P). Further, s(r) — 1 < s(l)/p. So,
s(r) + 1
-(0 + 1
And, B(p) > 1/(1 + 1 IP) = P/( 1 + P). Hence P/( 1 + P) < B(p) < 1/(1 + P) for
every p in T. So, T is in WB(or) for o = P/(l + P). â–¡

42
O
Figure 3.2. A tree in WB(l/4) that is not |-balanced
Remark While every /^-balanced tree, 0 < /? < 1/2, is in WB(a) for a = ¡3/(1+/3),
there are trees in WB(a) that are not /^-balanced. Figure 3.2 shows an example of a
tree in WB(l/4) that is not ^-balanced.
Lemma i If T is a COST then T is --balanced.
Proof If T is a COST, then every subtree of T is a COST. Consider any subtree
with root p, left child /, and right child r. If neither l nor r exist, then s(l) = s(r) = 0
and p is |-balanced. If s(l) — 0 and s(r) > 1, then r has a nonempty subtree with
root t and s(t) > s(l). So p is not a COST. Hence, s(r) < 1 and p is ^-balanced. The
same is true when s(r) = 0. So, assume s(l) > 0 and s(r) > 0.
If s(l) = 1, then s(r) < 3 as otherwise, one of the subtrees of r has m > 2
nodes and m > s(l) implies p is not a COST. Since s(r) < 3, ^(^(r) — 1) < s(l) and
^(s(/) — 1) < s(r). So, p is ^-balanced. The same proof applies when s(r) = 1. When
s(l) > 1 and s(r) > 1, let a and b be the roots of the left and right subtrees of l. Since
p is a COST, s(a) < s(r) and s(b) < s(r). So, s(l) = s(a) -f s(b) + 1 < 2s(r) + 1 and
|(s(/) — 1) < s(r). Similarly, |(s(r) — 1) < s(l). So, |-(/, r). Since this proof applies
to every nodes in T, the children of every p are ^-balanced and T is |-balanced. â–¡

43
Figure 3.3. ^-balanced tree that is not a COST
Remark There are ^-balanced trees that are not COSTs (see Figure 3.3).
While a COST is in WB(l/3) and WB(a) trees can be maintained efficiently
only for 2/11 < a < 1 — l/\/2 « 0.293, a COST is better balanced than WB(ar)
trees with a in the usable range. Unfortunately, we are unable to develop O(logn)
insert/delete algorithms for a COST.
In the next section, we develop insert and delete algorithms for /^-balanced
binary search trees (/?-BBST) for 0 < < y/2 — 1. Note that every (y/2 — 1)-BBST is
in WB(a) for a = 1 — l/\/2 which is the largest permissible a. Since our insert and
delete algorithms perform rotations along the search path whenever these result in
improved search cost, BBSTs are expected to have better search performance than
WB(a) trees (for a = /?/(1 + ¡3)).
Each node of a /J-BBST has the fields LeftChild, Size, Data, and RightChild.
Since every /3-BBST, /? > 0, is in WB(o), for a > 0, /2-BBSTs have height that is
logarithmic in n, the number of nodes (provided (3 > 0).

44
3.4 Search. Insert, and Delete in a ff-BBST
To reduce notational clutter, in the rest of the chapter, we abbreviate s(a) by a
(i.e., the node name denotes subtree size).
3.4.1 Search
This is done exactly as in any binary search tree. Its complexity is 0(h) where
h is the height of the tree. Notice that since each node has a size field, it is easy
to perform a search based on index (i.e., find the 10’th smallest key). Similarly, our
insert and delete algorithms can be adapted to indexed insert and delete.
3.4.2 Insertion
To insert a new element x into a /3-BBST, we first search for x in the /?-BBST.
This search is unsuccessful (as x is not in the tree) and terminates by falling off the
tree. A new node y containing x is inserted at the point where the search falls off
the tree. Let p' be the parent (if any) of the newly inserted node. We now retrace
the path from p' to the root performing rebalancing rotations.
There are four kinds of rotations LL, LR, RL, and RR. LL and RR rotations are
symmetric and so also are LR and RL rotations. The typical configuration before an
LL rotation is performed is given in Figure 3.4(a). p' denotes the root of a subtree in
which the insertion was made. Let p be the (size of the) subtree before the insertion.
Then, since the tree was a /2-BBST prior to the insertion, 0-(p,d). Also, for the LL
rotation to be performed, we require that (q > c) and (q > d). Note that q > d
implies q > 1. We shall see that (3-(q,c) follows from the fact that the insertion is

45
(a) before (b) after
Figure 3.4. LL rotation for insertion
made into a /3-BBST and from properties of the rotation. Following an LL rotation,
p' is updated to be the node p".
Lemma 5 [LL insertion lemma] If ](3-(p, d) A f3-(q, c) A (q > c) A (q > d)J for 0 < ¡3 <
1/2 before the rotation, then ¡3-{q,gp') and /?-(c, d) after the rotation.
Proof Assume the before condition.
(a) f3(q — 1) < c (as f3-(q,c)) < gp'. Also, f3(gp' — 1) = 0(c + d) < 2f3q (as /3 > 0,q>c
and q > d) < q (as ¡3 < 1/2). So, /3-(q,gp').
(b) d < q => d — 1 < q — 1 => f3(d — 1) < /3(q — 1) < c (as f3-(q,c)). Also,
P(c~ 1) < P(q + c- 1) = f3(p' -2) =p{p-\) < d (as 0-{p,d)). So, /3-(c,d). â–¡
In an LR rotation, the before configuration is as in Figure 3.4(a). However, this
time q < c. Figure 3.4(a) is redrawn in Figure 3.5(a). In this, the node labeled c in
Figure 3.4(a) has been labeled q and that labeled q in Figure 3.4(a) has been labeled

46
(a) before
(b) after substep (i)
Figure 3.5. Substep (i) of insertion LR rotation
a. With respect to the labelings of Figure 3.5(a), rotation LR is applied when
[(? > a)A(q> d)].
The other conditions that apply when an LR rotation is performed are
[/?-(p,d)A /Ma> ?)A c)]-
Here p denotes the (size of the) left subtree of gp prior to the insertion. An LR
rotation is accomplished in two substeps (or two subrotations). The first of these is
shown in Figure 3.5(b). Following an LR rotation, p' is updated to be node q'.
Lemma 6 [LR substep(i) insertion lemma] If [fl-(p,d) A fi-(a,q) A /?-(6, c) A (q >
a)A(q>d)]forO (c, d)) V (7^3-(ai b) A /3-(c, d))}] after the subrotation.

Proof Assume the before condition. First, we show that (3-(p",gp') after the
rotation. Note that (3(p" — 1) = 0(a + 6) = (3(a + 6 + c + 1) — 0(c + 1) = (3(p' — 1) —
0(c + 1) = 0(p- 1) - (3c < d-(3c < d < gp'. Also, fl(gp' - 1) = 0(c + d) < 6 + 0 + 0d
(as 0-(6,c)) d) p" = a + 6 + 1). So, p-(pC,gp').
Next, we prove two properties that will be used to complete the proof.
PI: 0(6-1) < a.
To see this, note that 0(6 — 1) < 0( P2: 0(c- 1) < d.
For this, observe that p' — \ — a+ q> (3(q — 1) + q (as 0-(a, q)) — (0 + 1)(<7 — 1) + 1.
So, q - 1 < = j^y. Similarly, q - 1 = 6 + c > 0(c - 1) + c (as 0-(6,c))
= 0» + l)(c - 1) + 1. So, fle - 1) < jfj(« - 2) < jfjfa - 1) < (as
0-(p,d)) < d.
To complete the proof of the lemma, we need to show
{(/»-(«,*) A V A/»-M))}.
We do this by considering the two cases 6 > c and b < c.
Case b > c: Since a PI imply 0-(a, 6). Also, d < 9 = 6 + c + 1. So, ^fy(d — 1) < ^(6 + c - 1) =
i?+Tc ~ 1) — ]3+ic + (as 0-(6, c)) = c. This, together with P2 implies
i+/r(c,d). So, 0-(a,6) A y|^-(c,d).
Case 6 < c: Since a < 9 = 6 + c + 1, a — 1 < 6 + c. So, a — l<6+c— lor

48
+ + Ht,c)) = 6. This and PI imply
Also, d - 1 < q - 2 = 6 + c - 1. So, @(d - 1) < /?(& + c - 1) < /3(2c - 1) < c. This,
together with P2 implies /3-(c, d). So, y^j-(a,6) A 0-(c,d). â–¡
Since an LR(i) rotation can cause the tree to lose its /3-balance property, it is
necessary to follow this with another rotation that restores the /3-balance property.
It suffices to consider the two cases of Figures 3.6 and 3.7 for this follow up rotation.
The remaining cases are symmetric to these. In Figures 3.6 and 3.7, p and d denote
the nodes that do not satisfy /3-(p, d). Note, however, that these nodes do satisfy
Since the follow up rotation to LR(i) is done only when
A (-/»-&>, <0),
either /3(p— 1) > d or /3(d — 1) > p. When /3(p— 1) > d, the second substep rotation is
one of the two given in Figures 3.6 and 3.7. When /3(d — 1) > p, rotations symmetric
to these are performed. In the following, we assume /3(p — 1) > d. Further, we may
assume d > 0, as d — 0 and y^-(p, d) imply p < 1. Hence, f3-(p,d). Also, d > 0 and
/3(p — 1) > d imply p > 1.
The LR(ii) LL rotation is done when the condition
A = (q > d) A (c < (1 + /3)q + (1 — 0)) A B where

49
(a) before (b) after
Figure 3.6. Case LL for LR(ii) rotation
B = y|^-(p,d) A A0-{q,c) A (p(p - 1) > d > 0).
Lemma 7 [Case LR(ii) LL rotation] If A holds before the rotation of Figure 3.6, then
P-{q,gp') and ft-(c,d) after the rotation provided 0 < ¡3 < >/2 — 1.
Proof (a) p-(q,gp')\
P{q~ 1) < c(as0-(q,c)) < gp'. Also, p(gp'-l) = p(c+d) < P((l+P)q+(l~P) + d) <
P{ 1 + p)q + P( 1 - 0) + - 1) (as q > d) = P(2 + - /?2 < 9 (as /?(2 + P) < 1
for 0 < £ < \/2 - 1). So, p-{q,gp').
(b)
^( ffe? + i|^(c “ 1) = + c - 1) = - 2) < ^(p - 1) < d (as ^-(;p,d)).
So, P-(c,d). â–¡
Lemma 8 If (c < (1 + P)q + (1 — /?)) A (/?(p — 1) > d) in Figure 3.6, then d < q
provided 0 < P < \/2 — 1.

50
Proof Sinced < 0(p— 1) = 0(q+c) < 0(q + (l+0)q+l—0) = 0(0+2)q + 0(l — 0) <
9 + 1 (as 0(0 + 2) < 1 and 0(1 - 0) < 1 for 0 < 0 < y/2 - 1). So, d < q. â–¡
So, the only time an LR(ii) LL rotation is not done is when C = (C\ V C2) A B
holds where
cx = (q = d) A (c < (1 + 0)q + 1-0)
C2 = c > (1 + 0)q + (1 — 0).
At this time, the LR rotation of Figure 3.7 is done. In terms of the notation of
Figure 3.7, the condition C becomes D — (D\ V D?) A E where
Di = (a = d) A (q < (1 + 0)a + 1 - 0)
D2 — q > (1 + 0)o. + 1 — 0
E = d) A d) A #•(«» 9) A c) A (0ÍP - 1) > ¿ > 0).
Lemma 9 When an LR(ii) LR rotation is performed and 0 < \/2 — 1, q > d and so
search cost is reduced.
Proof If Di, then since d < 0(p — 1) = 0(a + q) = 0(d + 9), q > d/0 — d > d as
fi < V2 ~ 1. K D,, then i < ftp-1) = 0(a + ?) < + ,) = Hgpq
*s0
51
a
b
d
LR(ii)
LR
a
b
c
d
c
(a) before
(b) after
Figure 3.7. Case LR for LR(ii) rotation
Lemma 10 When (d = a) A 0-(b, c) A (0(p — 1) > d) A (¡3 < \/2 — 1) (see Figure 3.7),
0(a — 1) < b and 0(d — 1) < c.
Proof Since 0(p — 1) > d and d = a, 0(p — 1) > a or 0(a + q) > a or a(l — 0) < 0q
or a< So, 0(a - 1) < ^q - 0 = ^(b + c + 1) - 0.
If c < | + 0, then
/»(<■-1) < ^ + ^ +
1-0 1-0 P
0(0+1 )b 0(0> + 0-l+0)
1-0 + 1-0
0(0+1 )b 0(02 + 20-l)
1-0 + 1-0
< b (as 0(0 -I- 1) < 1 - 0 for 0 < V2 - 1 and 02 + 20 - 1 < 0 for 0 < V2 - 1).

52
Since /9(c — 1) < 6, c < | -f 1. So,
ffl+l)b 3/92-/9
1-/9 1-/9 '
So,
However, since /92 + 2/9—1 < 0 for /9 < y/2 — 1, (1 + /9)/(l — /9) < and
(3/3 - 1)/(1 — /9) < /9. So, a - 1 < b/P + p. If a > c + 1, then c < a - 1 < 6//3 + /9.
We have already shown that for c < b/P + /9, /9(a — 1) < b. So, assume a < c + 1-
Now, a < c and /9(a — 1) < P(c - 1) < b (as /9-(6,c)). So, ,3(a - 1) < b in all cases.
P(a — 1) < c may be shown in a similar way. Since a = d, we get /9(d — 1) < c. □
Lemma 11 [Case LR(ii) LR rotation] If D holds before the rotation of Figure 3.7,
then P-(p',gp'), P-(a,b), and P-(c,d) following the rotation provided 0 < /9 < \/2 — 1.
Proof (a) p-(p’,gp'):
P(gp' — 1) = P{c + d) < b + P + Pd (as P-(b,c)) < b + /9 + Pq (from Lemmas 9
and 10, q > d) and q > d,P(p — 1) < (P + 1 )d or p(a -f q) < (P + 1 )d or a + q < (1 + j¡)d or
a < (1 4- jj)d-q < (1 + ±)d-d = d/p. So, P{p' - 1) = p(a + 6) (as P-(b, c)) < d-\- c+ 1 = gp'.
(b) P-(a,b):
Since b < q and P-(a,q),P(b— 1) < P(q — 1) < a.

53
When Di, P(a — 1) < b was proved in Lemma 10. So, /3-(a,6).
When D2, q > a(l + /3) + 1 - /3. So,
^ q 1 - /3 _ 6 + c+ 1 1-/3
“ÍT/3-HÑ3~ l+P l+P'
So,
/3(a-l)<
pb + /3c 4- P
1+/3
1-/3
1+/3
/3-/3<
/3¿> + ¿> + 2/3
1+/3
1-/3
1+/3
P — P —
b.
So, /3-(a,6).
(c) P-(c,d):
Note that /3(c - 1) < P(q - 1) < j^{q - 1) < y+^ÍP - 1) < d.
When D\, P(d — 1) < c was proved in Lemma 10. So, f3-(c,d).
When D2, if d < b + 1, then d < b and P(d- 1) < P(b- 1) < c. So, assume d > b+ 1.
Now, b < d — 1 < /3(p — 1) — 1. So,
b < /3(a + 6 + C+1) — 1
< Hq~+ * + c+l)-l
= j^(4+c+/J+(l+/?)(6+c+l))-l
< j-^(|+1+c + ^ + (1+W(| + 1 + c+1))-1
= I~(^Ílc + (l+/?) + (l + /?)(iÍ^<:+2))-l
= c + P + (\ + P)c + 2/3 — 1
= (2 + P)c + 3/3 — 1 < (2 + P)c + P (as /3 < \/2 — 1)

54
< ^ + 0 (as P < v/2- 1).
Also, from d < P(p — 1) and the above derivation, we get
d <
<
(6+ c+ /? + (! + p)(b + c + 1))
1+0
rr^(|+y9+c+/9+(1+/*)
= TT^O+r&+^c)+w+1)
2Q2
= (2 + 0)c + + P{P + 1)
2/?2 + /92 + /?3 + p + /?2
— (2 + P)c +
= (2 + P)c +
l+P
P3 -f 4/?2 + P
1 +P
< (2 + P)c+l {asP3 + 4p2 + P< l+pfor p< v/2 - 1).
So, 0(d - 1) < p{2 + p)c < c (as P < v/2 - 1). So, /?-(c, d).
â–¡
Theorem 9 If T is P-balanced, 0 < P < \/2 — 1, prior to insertion, it is so following
the insertion.
Proof First note that since all binary search trees are balanced for P = 0, the rota¬
tions (while unnecessary) preserve 0-balance. So, assume P > 0. Consider the tree V
just after the new element has been inserted but before the backward restructuring
pass begins.
If the newly inserted node, z, has no parent in T\ then T was empty and T' is
/^-balanced. If z has a parent but no grandparent, then T has at most one nonempty

55
subtree X. Since T is /2-balanced, (d(\X\ — 1) < 0. So, |.Y| < 1. Following the
insertion, T' has one subtree with < 1 nodes and one with exactly one. So, T' is
/2-balanced. We may therefore assume that z has a grandparent in T'.
From the downward insertion path, it follows that all nodes u in T' that have
children / and r for which ->/2-(/, r) must lie on the path from the root to z. During
the backward restructuring pass, each node on this path (other than 2 and its parent)
play the role of gp in Figures 3.4 and 3.5. The /2-property cannot be violated at z
as z has no children. It cannot be violated at the parent, s, of z as s satisfied the
/^-property prior to insertion. As a result its other subtree has < 1 element. So,
following the insertion, s satisfies the /2-property. As a result, each node in T' that
might possibly violate the /2-property becomes the gp node during the restructuring
pass. Consider one such gp node. It has children in T' denoted by p' and d. Its
children in T are p and d. Figures 3.4 and 3.5 show the case when d is the right
subtree of gp in both T and T'. The cases RR and RL arise when d is the left
subtree.
During the restructuring pass, gp begins at the grandparent of z and moves up
to the root of T'. If z is at level r in T'’, (the root being at level 1), then gp takes on
r — 2 values during the restructuring pass. We shall show that at each of these r — 2
positions either
(a) no rotation is performed and all descendants of gp satisfy the /2-property or
(b) a rotation is performed and following this, all descendants of node p" (Figure 3.4)
or of node q' (Figure 3.5) satisfy the /2-property.

56
As a result, following the rotation (if any) performed when gp becomes the root
of T', the restructured tree is /^-balanced. The proof is by induction on r. When
r = 3 (recall, we assume z has a grandparent), gp begins at the root of T' and its
descendants satisfy the /9-property.
Without loss of generality, assume that the insertion took place in the left subtree
of gp. With respect to Figure 3.4, we have three cases: (i) q > c and q > d, (ii) q < c
and c > d, and (iii) q < d and c < d. In case (i), all conditions for an LL rotation hold
and such a rotation is performed. In case (ii), an LR rotation is performed. Following
either rotation, V is /9-balanced. In case (iii), /3(p' — 1) = 0(q + c) < 2/3d < d (as
fi < y/2 - 1). Also, fi(d — 1) < p < p + 1 = p\ So, /3(d - 1) < p'. Hence, /9-(p',d)
and T' is /9-balanced.
For the induction hypothesis, assume (a) and (b) whenever r < k. In the
induction step, we show (a) and (b) for trees T with r = k + 1. The subtree in
which the insertion is done has r = k. So, (a) and (b) hold for all gp locations in the
subtree. We need to show (a) and (b) only when gp is at the root of T'. This follows
from Lemmas 5, 6, 7, and 11.
The theorem now follows. â–¡
Lemma 12 The time needed to do an insertion in an n node 0-BBST is O(logn)
provided 0 < ¡3 < y/2 — 1.

57
Proof Follows from the fact that insertion takes 0(h) time where h is the tree
height and h = O(logn) when ¡3 > 0 (Lemmas 1 and 3). □
3.4.3 Deletion
To delete element x from a /LBBST, we first use the unbalanced binary search
tree deletion algorithm of Horowitz and Sahni [14] to delete x and then perform a
series of rebalancing rotations. The steps are:
Step 1 [Locate x] Search the /3-BBST for the node y that contains x. If there is no
such node, terminate.
Step 2 [Delete x\ If y is a leaf, set dl to nil, gp to the parent of y, and delete node
y. If y has exactly one child, set dl to be this child; change the pointer from
the parent (if any) of y to point to the child of y; delete node y; set gp to be
the parent of dl. If y has two children, find the node z in the left subtree of y
that has largest value; move this value into node y; set y = z; go to the start
of Step 2. { note that the new y has either 0 or 1 child }
Step 3 [Rebalance] Retrace the path from dl to the root performing rebalancing
rotations.
There are four rebalancing rotations LL, LR, RR, and RL. Since LL and RR
as well as LR and RL are symmetric rotations, we describe LL and LR only. The
discussion is very similar to the case of insertion. The differences in proofs are due
to the fact that a deletion reduces the size of encountered subtrees by 1 while an

58
(a) before (b) after
Figure 3.8. LL rotation for deletion
insertion increases it by 1. In an LL rotation, the configuration just before and after
the rotation is shown in Figure 3.8. This rotation is performed when q > c and
q > d!. Following the rotation, d' is updated to the node p'.
Let d denote the size of the right subtree of gp before the deletion. So, d = d'-fi 1.
Since prior to the deletion the /3-BBST was /^-balanced, it follows that /?-(p, d) and
/T(9, c).
Lemma 13 [LL deletion lemma] If [0-(p,d) A(3-(q,c) A(q > c)l\(q > A(1/3 < /? <
1/2)7 before the rotation, then [&-(q,gp') A 0-(c,d')] after the rotation.
Proof (a) /3-(q,gp'):
0{q — 1) < c (as /3-(q,c)) < gp'. Also, (3(gp' — 1) = /3(c + d!) < 20q (as c < q and
d' < q) < q (as P < 1/2). So, f3-(q,gp').
(b) /?-M'):
d! < q => d' — 1 < q — 1 => 0(d' — 1) < 0(q — 1) < c. Also, when c < 1, /?(c—1) < 0 < d'
(as d! > 0). When c > \,q > c =$■ q > 2 and p — q + c + 1 >c + 3. So,
f3{c - 1) < (){p - 1) - 3/? < d - 3(3 (as 0-(p, d)) < d - 1 (as > 1/3) = d!. Hence,

59
In an LR rotation, the before configuration is as in Figure 3.8(a). However, this
time q < c. Figure 3.8(a) is redrawn in Figure 3.9(a). In this, the node labeled c
in Figure 3.8(a) has been relabeled q and that labeled q in Figure 3.8(a) has been
relabeled a. With respect to the labelings of Figure 3.9(a), rotation LR is applied
when
[(? > a) A ( d!)\.
The other conditions that apply when an LR rotation is performed are
[^-(p, d) A q) A /?-(6, c)].
Here d denotes the (size of) right subtree of gp prior to the deletion. As in the case
of insertion, an LR rotation is accomplished in two substeps (or two subrotations).
The first of these is shown in Figure 3.9. Following an LR rotation, d1 is updated to
node q'.
Lemma If [LR substep(i) deletion lemma] If [/3-(p, d) A (3-(a, q) A /?-(6, c) A (q > a) A
(q > d')] before the subrotation LR(i), then [/3-(p',gp')A{(f3-(a, 6)Ay^j-(c, d'))V(y^j-
(a,b) A /?-(c, d'))}] after the subrotation provided 1/3 < R < 1/2 .
Proof Assume the before condition.
(a) If b = c = 0, then <7 = 6 + c + 1 = 1. Furthermore, (q > a) and (q > d') imply
a = d! = 0. So, gp' = p' = 1. Hence, \\-(p',gp') A |-(a,6) A |-(c,d')]

60
(a) before
(b) after substep (i)
Figure 3.9. LR rotation for deletion
(b) If b = 1 and c = 0, then q = 2, a < 1, and d! < 1. So, 1 < p' < 3 and 1 < gj/ < 2.
Hence, [§-(?', gjt) A ±-(a,6) A |-(c,d')]
(c) If b = 0 and c = 1, then q = 2, a < 1, and d! <1. So, 1 < p' < 2 and 1 < gp' < 3.
Hence, \\-{j/,gp') A \-(a,b) A \-{c,d')\
As a result of (a) - (c), to complete the proof, we may assume that b > 1 and c > 1.
So, q > 3, a > 1 (as /3-(a, q) f3(q — 1) < a or a > 2/3 > 0), p = a + q + 1 > 5,d > 2
(as fi-(p,d) => /9(p — 1) < d and 0 > 1/3), and d' = d — 1 > 1.
First, we show that (3-(p',gp'). For this, note that a + 6 + c+ l = p — 1. From
/3-(p, d), it follows that (3(a + 6 + c + 1) = /9(p — 1) < d. So, y9(a + b) < d — flc — 0.
From Figure 3.9(b), we see that 0(p' — 1) = 0{a + b). Hence, 0{p' — 1) < d — 0c — 0 =
d' — 0c + 1 — 0 < d! + 1 — 20 < gp'. Also,
0(gp'-\) = 0(c + d') < b 0 + 0d! (as 0-(b, c))
< b + 0q + 0 (as q > d')

61
< 6 + a + 2¿0 (as 0-(a,q))
< ?'â– 
So, P-(p',gp').
Next, we prove two properties that will be used to complete the proof.
PI: ¿0(6-1) To see this, note that ¿0(6 — 1) < ft(q — 1) < a (as /?-(a, 9)).
P2: P(c- 1) < d‘.
For this, observe that ¿9(c — 1) < (3(q — 2) (as c < q — 1) < fi(p — 4) (as q = p — a — 1
and a > 1) = ¿0(p - 1) — 3/0 < d - 1 (as ¿9-(p, d) and ¿3 > 1/3) = d'.
To complete the proof of the lemma, we need to show
{(H“, 4) A Jd')) V (“•(«, 4) A 0-(c, i))l
For this, consider the two cases 6 > c and b < c (as in Lemma 6).
Case b > c: Since a with PI implies ¿0-(a, 6). Also, d' ^+IC+ *) - ^+IC+ ¡3+1 = c’ This’ toSether with P2 imples ^-{c,d'). So,
/Ma>&) A d')-
Case b < c: Since a < q — 6 + c + 1,a — 1 < 6 + c. So, a — 1 < 6 + c— 1
°r < i% + TT0 = b. This and PI imply ^-(0,6). Also,
d' — 1 < q — 2 = b + c — 1. So, /?(d' — 1) < /3(6 + c — 1) < /?(2c — 1) < c. This and

62
P2 imply f3-(c,d'). Hence, j^-(a,b) A 0-(c,d'). â–¡
The substep(ii) rotations are the same as for insertion.
Theorem 10 IfT is ^-balanced, then following a deletion the resulting tree V is also
¡3-balanced provided 1/3 < ¡3 < y/2 — 1.
Proof Similar to that of Theorem 9. a
When 0 < ¡3 < 1/3, we need to augment the LL rotation by a transformation
for the case d! = 0. When d! = 0, /3(p - 1) < d = d! + 1 = 1. So, p < 1//2 + 1 and
gp = p + d! + 1 < 1//2 + 2. To /2-balance at gp, the at most 1//2 + 2 nodes in gp
are rearranged into any /2-BBST in constant time (as 1//2 + 2 is a constant). When
d! > 0, the proof of Lemma 13 part (b) can be changed to show /?(c — 1) < d! for
0 < ¡3 < y/2 — 1. The new proof is: since c < q,c < (p — l)/2 and /?(c — 1) <
/?(p — \)/2 — d < d/2 — /3 = d — d/2 — ¡3 < d — 1— /3 < d1. The LR rotation needs
to be augmented by a transformation for the case d! — d — 1 < p^+p) ~ T this
time, f3(p - 1) < d < j^py So, gp = p + d < pi^j) + 1 + p^+p]- To /2-balance at
gp, we rearrange the fewer than pi^+p) + 1 + p(2+pj no<^es >n subtree, in constant
time, into any ^-balanced tree. When d! > p¡j+pj — 1, the proof for /2(c — 1) < d' in
Lemma 14 needs to be changed to show that the LR substep(i) lemma holds. The
new proof is:
d > /3{p - 1) = /3(a + b + c + 1) > /3(/3(q - 1) -|- b + c-f 1)

63
= ${/?(&+ c) 4- b + c+ 1)
> /9((1 + /J)/J(c-1) + (1 + /J)c+1)
= 0((1+/?)’(<:-1)+ 2 + 0).
So, 0(c - 1) < < rf - 1 («« á > *,¿35) = *.
Also, note that when /3 = 0, all trees are /^-balanced so the rotations (while not
needed) preserve balance.
Theorem 11 With the special handling of the case d' = 0, the tree T' resulting from
a deletion in a /3-BBST is also (3-balanced for 0 < (3 < y/2 — 1.
Lemma 15 The time needed to delete an element from an n node f)-BBST is 0(\ogn)
provided 0 < ¡3 < \/2 — 1.
3.4.4 Enhancements
Since our objective is to create search trees with minimum search cost, the
rebalancing rotations may be performed at each positioning of gp during the backward
restructuring pass so long as the conditions for the rotation apply rather than only
at gp positions where the tree is unbalanced.
Consider Figure 3.4(a). If p' < d, then the conditions of Lemmas 5 and 6 cannot
apply as q < p' < d. However, it is possible that e > p' where e is the size of either the
left or right subtree of d. In this case, an RR or RL rotation would reduce the total
search cost. The proofs of Lemmas 5 and 6 are easily extended to show that these
rotations would preserve balance even though no insertion was done in the subtree

64
d. The same observation applies to deletion. Hence the backward restructuring pass
for the insert and delete operations can determine the need for a rotation at each gp
location as below (/ and r are, respectively, the left and right children of gp).
i MO > s(r) then check conditions for an LL and LR rotation
else check conditions for an RR and RL rotation.
The enhanced restructuring procedure used for insertion and deletion is given in
Figure 3.10. In the RR and RL cases, we have used the relation *>’ rather than ‘>’
as this results in better observed run time.
Since it can be shown that the rotations preserve balance even when there
has been no insert or delete, we may check the rotation conditions during a search
operation and perform rotations when these improve total search cost.
Finally, we note that it is possible to use other definitions of /3-balance. For
example, we could require fl(s(a) — 2) < s(b) and /3(s(2>) — 2) < s(a) for 0-(a,b).
One can show that the development of this chapter applies to these modifications
also. Furthermore, when this new definition is used, the number of comparisons in
the second substep of the LR and RL rotations is reduced by one.
3.4.5 Top Down Algorithms
As in the case of red/black and WB(a) trees, it is possible to perform, in 0(log n)
time, inserts and deletes using a single top to bottom pass. The algorithms are similar
to those already presented.

65
procedure Restructuring ;
begin
while (gp) do
begin
if (s(gp.left) > s(gp.right)) then
begin {check conditions for an LL and LR rotation}
p = gp.left ;
if (s(p.left) > s(p.right)) then
begin if (s(p.left) > s(gp.right)) then do LL rotation; end
else
begin
if (s(p.right) > s(gp.right)) then {LR}
begin
do LR rotation ;
{ now notations a, 6, c, and d follow from figure 3.1(b) }
if (/?(s(a) — 1) > s(b)) then
if ((s(a.right) < (1 + 0)s(a.left) + 1 — 0) and
(s(b) < s(a.left))) then
do LL rotation
else do LR rotation
else if (/?(s(d) — 1) > s(c)) then
if ((s(d.left) < (1 + P)s(d.right) +1—0) and
(s(c) < s(d.right))) then
do RR rotation
else do RL rotation ;
end
end
end
else {check conditions for an RR and RL rotation}
begin
P = gp-right ;
if (s(p.left) > s(p.right)) then
begin
if (s(p.left) > s(gp.left)) then {RL}
do symmetric to the above LR case ;
end
else
begin if (s(p.right) > s(gp.left)) then do RR rotation; end ;
end ;
gp = gp.parent ;
end ;
end ;
Figure 3.10. Restructuring procedure

66
3.5 Simple /3-BBSTs
The development of Section 3.4 was motivated by our desire to construct trees
with minimal search cost. If instead, we desire only logarithmic performance per
operation, we may simplify the restructuring pass so that rotations are performed
only at nodes where the /3-balance property is violated. In this case, we may dispense
with the LL/RR rotations and the first substep of an LR/RL rotation. Only LR/RL
substep (ii) rotations are needed. To see this, observe that Lemmas 7 and 11 show that
the second substep rotations rebalance at gp (see Figures 3.6 and 3.7) provided
(p, d) (The remaining conditions are ensured by the bottom-up nature of restructuring
and the fact the tree was ^-balanced prior to the insert or delete).
If the operation that resulted in loss of balance at gp was an insert, then /?(p —
2) < d (as p > d, the insert took place in subtree p and gp was ^-balanced prior
to the insert) and fi(p — 1) > d (gp is not ^-balanced following the insert). For the
substep (ii) rotation to restore balance, we need 0(p— 1) < (1 + fd)d. This is assured
if d + /? < (/? + l)d (as /3(p — 2) < d). So, we need d > 1. If d < 1, then d = 0. Now
/?(p — 2) < d and ¡3(p — 1) > d imply p = 2. One may verify that when p = 2, the
LR(ii) rotations restore balance.
If the loss of /3-balance at gp is the result of a deletion (say from its right subtree),
then /?(p — l) < d + 1 (as gp was /3-balanced prior to the delete). For the substep
(ii) rotation to accomplish the rebalancing, we need /?(p — 1) < (/? F l)d. This is
guaranteed if d+ 1 < ((3 + l)d or d > 1//3. When d < 1//3 and > 1/3, d < 2. Since
0(p — 1) < d + 1 and /3 > 1/3, when d = 2, p < 10; when d = 1, p < 7; and when

67
d = O, p < 4. We may verify that for all these cases, the LR(ii) rotations restore
balance. Hence, the only problematic case is when /? < 1/3 and d < 1/(3.
When /? < 1/3, an LL rotation fails to restore balance only when d = 0 (see
discussion following Theorem 10). So we need to rearrange the at most 1/(3 + 2
nodes in gp into any /^-balanced tree when d — 0. An LR rotation fails only when
d < p^+0) ~~ To see this, note that in the terminology of Lemma 14, d is dl.
The proof of P2 is extended to the case /? < 1/3 when dl > p^+p) ~ 1- Also, since
dl < 1//3, for the case b > c, we get /3(d' — 1) < 1 — (3 < c (as c > 1). For the
case b < c, we need to show (3(a — 1) < b. Since an LR rotation is done only when
condition Dl V D2 holds, from Lemmas 10 and 11, it follows that /?(a — 1) < b. So,
an LR rotation rebalances when 0 < 1/3 provided d > p^+p) ~ 1- F°r smaller d, the
at most pip+p) + 0(2+0) + 1 nodes in the subtree gp may be directly rearranged into
a /3-balanced tree.
The restructuring algorithm for simple /3-BBSTs is given in Figures 3.11 and 3.12.
The algorithm of Figure 3.11 is used following an insert and that of Figure 3.12 after
a delete.
Simple /S-BBSTs are expected to have higher search cost than the /3-BBSTs
of Section 3.4. However, they are a good alternative to traditional WB(a) trees as
they are expected to be “better balanced”. To see this, note that from the proof of
Lemma 3, the balance, B(p), at any node p in a ^-balanced tree satisfies
1 _ 1 , ^(r) + !
B(p) + s(l) + 1

68
procedure Restructuring2 ;
begin
while (gp) do
begin
if (0(s(gp.left) — 1) > s(gp.right)) then {do an LL or LR rotation)
begin
P = gp iefi;
if ((s(p.right) < (1 + 0)s(p.left) + 1—0) and
(s(gp.right) < s(p.left))) then
do LL rotation
else do LR rotation ;
end
else
do symmetric to the above L case ;
gp = gp.parent ;
end ;
end ;
Figure 3.11. Simple restructuring procedure for insertion
procedure Restructurings ;
begin
while (gp) do
begin
if (0(s(gp.left) — 1) > s(gp.right)) then
if (0 < 1/3) and (s(gp.right) < 1/0(2 + 0) — 1) then
rearrange the subtree rooted at gp into any /^-balanced tree
else {do an LL or LR rotation)
begin
p = gp.left ;
if ((s(p.right) < (1 + 0)s(p.left) + 1 — 0) and
(s(gp.right) < s(p.left))) then
do LL rotation
else do LR rotation ;
end
end
else
do symmetric to the above L case ;
gp = gp.parent ;
end ;
end ;
Figure 3.12. Simple restructuring procedure for deletion

69
>
1 +
1/0 +
2/3-1
«•(r)+l)
2/3-1
1 + 0 + 0(«(r)+l)
I 4. 2/3-1
P ^ /»(*(r)+l)
So,
B(p) < 1
1
1 + p +
2/3-1
/3(a(r)+l)
Also, since s(r) - 1 < a(/)/0, s(r) + 1 < s(l)/0 + 2. Hence, 1 + < 1 + +
*(0+i'
So,
B(p) | , 1 i . 2
1 + 0 0(s(l)+1) T i(/)+l
1
“ 1 . i I 2/3-1 •
1 + /3 + /3(*(0+l)
Consequently,
1 + i + 2/?~1
1 + P + ^(*(0+i)
< B(P) <
1
1 + j +
2/3 — 1 ’
/3(*(r)+l)
When 0= v^- 1,
~r < #(p) < 1
2 + ^+7^ 2 + v^+^ft
If s(p) < 10, 0.296 < B(p) < 1 — 0.296. So, every 0-balanced subtree with 10 or
fewer nodes is in WB(a) for a ss 0.296. Similarly, every subtree with 100 or fewer
nodes is in WB(a) for a « 0.293. In fact, for every fixed k, subtrees of size k or less

70
procedure Restructuring4 ;
begin
while (gp) do
begin
if (s(gp.left) > s(gp.right)) then
begin {check conditions for an LL and LR rotation}
p = gp left 5
if (s(p.left) > s(p.right)) and (s(p.left) > s(gp.right)) then
do LL rotation
else if (s(p.left) < s(p.right)) and (s(p.right) > s(gp.right)) then
do LR rotation ;
end
else {check conditions for an RR and RL rotation}
do symmetric to the above L case ;
gp = gp.parent ;
end ;
end ;
Figure 3.13. Simple restructuring procedure without a /? value
are in WB(a) for a slightly higher than 0.2929 which is the largest value
of a for which WB(a) trees can be maintained.
3.6 BBSTs without Deletion
In some applications of a dictionary, we need to support only the insert and
search operations. In these applications, we can construct binary search trees with
total cost
C(r) by using the simpler restructuring algorithm of Figure 3.13.
Theorem 12 When the only operations are search and insert and restructuring is done
as in Figure 3.13, C(T) < nlog¿(\/5(n + 1)).

71
Proof Suppose T currently has m — 1 elements and a new element is inserted. Let
u be the level at which the new element is inserted. Suppose that the restructuring
pass performs rotations at q < u of the nodes on the path from the root to the newly
inserted node. Then C{T) increases by at most v — u — q as a result of the insertion.
The number of nodes on the path from the root to the newly inserted node at which
no rotation is performed is also v. Let these nodes be numbered 1 through v bottom
to top. Let Si denote the number of elements in the subtree with root i prior to the
restructuring pass. We see that Si > 1 and S2 > 2. For node 1, 2 < i < v, one of its
subtrees contains node i — 1. Without loss of generality, let this be the left subtree
of 1. Let the root of the right subtree of t be d. So,
Si > Si-i + s(d) + 1.
If z — 1 is not the left child of t, then since no rotation is done at *, s(d) > 5,_ 1. If
z — 1 is the left child of 1, then consider node i — 2. This is in one of the subtrees of
i. Since no rotation is performed at i — 1, s(d) > Si-2■ Since Si-1 > Si-2, we get
Si > Si-1 + Si-2 + L
Hence, Sv > Nv where Nv is the minimum number of elements in a COST of
height v. So, v < log^(\/5(m + 1)). So, when an element is inserted into a tree that
has m — 1 elements, its cost C(T) increases by at most log¿(v/5(m + 1)). Starting
with an empty tree and inserting n elements results in a tree whose cost is at most

72
nlog*(V5(n +1)). D
Corollary 2 The expected cost of a search or insert in a BBST constructed as above
is O(\ogn).
Proof Since C(T) < nlog^(v/5(n + 1)), the expected search cost is C(T)/n <
l°g^(v/5(n + 1)). The cost of an insert is the same order as that of a search as each
insert follows the corresponding search path twice (top down and bottom up). â–¡
3.7 Experimental Results
For comparison purposes, we wrote C programs for BBSTs, SBBSTs (simple
BBSTs), BBSTDs (BBSTs in which procedure Restructuring4 (Figure 3.13) is used
to restructure following inserts as well as deletes), unbalanced binary search trees
(BST), AVL-trees, top-down red-black trees (RB-T), bottom-up red-black trees (RB-
B) [31], weight balanced trees (WB), deterministic skip lists (DSL), treaps (TRP),
and skip lists (SKIP). For the BBST and SBBST structures, we used /? = 207/500
while for the WB structure, we used a = 207/707. While these are not the highest
permissible values of /? and a, this choice permitted us to use integer arithmetic
rather than the substantially more expensive real arithmetic. For instance, /?-(a,6)
for /? = 207/500 can be checked using the comparisons 207(s(a) — 1) > 500s(6)
and 207(s(6) — 1) > 500s(a). The randomized structures TRP and SKIP used the
same random number generator with the same seed. SKIP was programmed with
probability value p = 1/4 as in Pugh [26].

73
To minimize the impact of system call overheads on run time measurements,
we programmed all structures using simulated pointers (i.e., an array of nodes with
integer pointers [27]. Skip lists use variable size nodes. This requires more complex
storage management than required by the remaining structures which use nodes of
the same size. For our experiments, we implemented skip lists using fixed size nodes,
each node being of the maximum size. As a result, our run times for skip lists are
smaller than if a space efficient implementation had been used. In all our tree struc¬
ture implementations, null pointers were replaced by a pointer to a tail node whose
data field could be set to the search/insert/delete key and thus avoid checking for
falling off the tree. Similar tail pointers are part of the defined structure of skip and
deterministic skip lists. Each tree also had a head node. WB(a) trees were imple¬
mented with a bottom-up restructuring pass. Our codes for SKIP and DSL are based
on the codes of Pugh [26] and Papadakis [22], respectively. Our AVL and RB-T codes
are based on those of Papadakis [22] and Sedgewick [28]. The treap structure was
implemented using joins and splits rather than rotations. This results in better per¬
formance. Furthermore, AVL, RB-B, WB, and BBST were implemented with parent
pointers in addition to left and right child pointers. For BBSTs, the enhancements
described in Section 3.4.4 for insert and delete (see Figure 3.10) were employed. No
rotations were performed during a search when using any of the structures.
For our experiments, we tried two versions of the code. These varied in the
order in which the ‘equality’ and ‘less than’ or ‘greater than’ check between x and e
(where x is the key being searched/inserted/deleted and e is the key in the current

74
node) is done. In version 1, we conducted an initial experiment to determine if the
total comparison count is less using the order L:
if x < e then move to left child
else if x ^ e then move to right child
else found
or the order R:
if x > e then move to right child
else if x / e then move to left child
else found.
Our experiment indicated that doing the ‘left child’ check first (i.e. order L) worked
better for AVL, BBST, BBSTD, and DSL structures while R worked better for the
RB-T, RB-B, WB, SBBST, and TRP structures. No significant difference between L
and R was observed for BSTs. For skip lists, we do not have the flexibility to change
the comparison order. The version 1 codes performed the comparisons in the order
determined to be better. For BSTs, the order R was used.
In the version 2 codes the comparisons in each node took the standard form
if x = e then found
else if x < e then move to left child
else move to right child.
The version 2 restructuring code for BBSTs differed from that of Figure 3.10 in
that the ‘>’ test in the second, third, and forth if statements was changed to *>’.

75
No change was made in the corresponding if statements for RR and RL rotations.
While this increased the number of comparisons, it reduced the run time.
We experimented with n = 10,000, 50,000, 100,000, and 200,000. For each n,
the following experiments were conducted:
(a) start with an empty structure and perform n inserts;
(b) search for each item in the resulting structure once; items are searched for in the
order they were inserted
(c) perform an alternating sequence of n inserts and n deletes; in this, the n elements
inserted in (a) are deleted in the order they were inserted and n new elements are
inserted
(d) search for each of the remaining n elements in the order they were inserted
(e) delete the n elements in the order they were inserted.
For each n, the above five part experiment was repeated ten times using different
random permutations of distinct elements. For each permutation, we measured the
total number of element comparisons performed and then averaged these over the ten
permutations.
First, we report on the relative performance of SBBSTs, BBSTDs, and BB-
STs. For this comparison, we used only version 1 of the code. Table 3.1 gives the
average number of key comparisons performed for each of the five parts of the ex¬
periment. The three versions of our proposed data structure are very competitive
on this measure. BBSTDs and BBSTs generally performed fewer comparisons than
did SBBSTs. All three structures had a comparison count within 2% of one another.

76
Table 3.1. The number of key comparisons on random inputs (version 1 code)
n
operation
SBBST
BBSTD
BBST
insert
212495
212223
212111
search
194661
191599
191578
10,000
ins/del
416125
416967
416862
search
194957
191666
191676
delete
168033
166441
166487
insert
1241080
1236682
1236114
search
1152137
1135131
1134969
50,000
ins/del
2437918
2438083
2437639
search
1153821
1134277
1134062
delete
1018675
1007766
1007688
insert
2635913
2624829
2623792
search
2458079
2423988
2423613
100,000
ins/del
5183619
5180383
5179653
search
2461221
2420282
2419990
delete
2190798
2168049
2168110
insert
5580139
5555190
5553256
search
5223989
5148220
5147698
200,000
ins/del
10981441
10969578
10968053
search
5229172
5144808
5144148
delete
4692447
4641349
4641389
However, when we used ordered data rather than random data (Table 3.2), SBBSTs
performed noticeably inferior to BBSTDs and BBSTs; the later two remained very
competitive.
Tables 3.3 and 3.4 give the average heights of the trees using random data and
using ordered data, respectively. The first number gives the height following part (a)
of the experiment and the second following part (c). The numbers are identical for
BBSTDs and BBSTs and slightly higher (lower) for SBBSTs using random (ordered)
data.

77
Table 3.2. The number of key comparisons on ordered inputs (version 1 code)
n
operation
SBBST
BBSTD
BBST
insert
170182
150554
150554
search
188722
185530
185530
10,000
ins/del
425305
315177
314998
search
191681
184155
184155
delete
215214
135311
135131
insert
991526
872967
872967
search
1117174
1101481
1101481
50,000
ins/del
2472808
1806346
1805439
search
1116390
1098065
1098065
delete
1277756
792717
791815
insert
2103808
1850548
1850548
search
2384327
2354757
2354757
100,000
ins/del
5249194
3823415
3821594
search
2382759
2346118
2346128
delete
2738294
1686397
1684584
insert
4449143
3903083
3903083
search
5068632
4946753
4946753
200,000
ins/del
11105525
8051695
8048058
search
5065496
5001967
5001967
delete
5842168
3580856
3577223
Table 3.3. Height of the trees on random inputs (version 1 code)
n
SBBST
BBSTD
BBST
10,000
17,17
16,16
16,16
50,000
20,20
19,19
19,19
100,000
21,21
20,20
20,20
200,000
22,23
21,21
21,21
Table 3.4. Height of the trees on ordered inputs (version 1 code)
n
SBBST
BBSTD
BBST
10,000
16,15
17,17
17,17
50,000
20,20
20,20
20,20
100,000
21,21
21,21
21,21
200,000
22,22
23,22
23,22

78
The average number of rotations performed by each of the three structures is
given in Tables 3.5 and 3.6. A single rotation (i.e., LL or RR) is denoted ‘S’ and a
double rotation (i.e., LR or RL) denoted ‘D\ In the case of BBSTs, double rotations
have been divided into three categories: D = LR and RL rotations that do not perform
a second substep rotation; DS = LR and RL rotations with a second substep rotation
of type LL and RR; DD = LR and RL rotations with a second substep rotation of
type LR and RL. BBSTDs and BBSTs performed a comparable number of rotations
on both data sets. However, on random data SBBSTs performed about half as many
rotations as did BBSTDs and BBSTs. On ordered data, SBBSTs performed 15 to
20% fewer rotations on part (a), 34% fewer on part (c), and 51% fewer on part (e).
The run-time performance of the structures is significantly influenced by com¬
piler and architectural features as well as the complexity of a key comparison. The
results we report are from a SUN SPARC-5 using the UNIX C compiler cc with
optimization option. Because of instruction pipelining features, cache replacement
policies, etc., the measured run times are not always consistent with the compiler
and architecture independent metrics reported in Tables 3.1 through 3.6 and later
in Tables 3.11 through 3.16. For example, since the search codes for all tree based
methods are essentially identical, we would expect methods with a smaller compar¬
ison count to have a smaller run time for parts (b) and (d) of the experiment. This
was not always the case.
Tables 3.7 and 3.8 give the run times of the three BBST structures using integer
keys and Tables 3.9 and 3.10 do this for the case of real (i.e., floating point) keys. The

Table 3.5: The number of rotations on random inputs (version 1 code)
n
operation
SBBST
BBSTD
BBST
S
D
S
D
S
D
DS
DD
insert
2341
2220
5045
4314
5025
3938
151
93
10,000
ins/del
4269
3216
10158
6311
10104
5849
232
103
delete
1607
1110
5235
2104
5201
2018
51
28
insert
11719
11120
25216
21596
25059
19732
754
455
50,000
ins/del
21330
16125
51238
31499
50979
29198
1161
531
delete
8058
5648
26214
10462
26068
10033
248
131
insert
23450
22262
50283
43230
50047
39461
1527
920
100,000
ins/del
42780
32203
102218
62967
101836
58491
2275
1046
delete
16095
11306
52227
21022
51943
20147
496
260
insert
46934
44525
100664
86605
100205
79013
3054
1840
200,000
ins/del
85283
64417
204459
125960
203568
116940
4593
2059
delete
32233
22551
104344
41884
103826
40157
990
523

80
Table 3.6. The number of rotations on ordered inputs (version 1 code)
n
operation
SBBST
BBSTD
BBST
S
D
S
D
S
D
DS
DD
insert
9984
0
9985
2387
9985
2387
0
0
10,000
ins/del
14997
0
16567
6130
16644
5797
25
154
delete
4989
0
6570
3726
6647
3392
26
154
insert
49980
0
49983
11956
49983
11956
0
0
50,000
ins/del
74996
0
82862
30659
83247
28982
137
770
delete
24987
0
32859
18686
33242
17018
136
766
insert
99979
0
99983
23917
99983
23917
0
0
100,000
ins/del
149996
0
165738
61327
166504
57969
280
1540
delete
49986
0
65733
37392
66505
34040
278
1536
insert
199978
0
199982
47839
199982
47839
0
0
200,000
ins/del
299996
0
331473
122653
333012
115938
559
3078
delete
99985
0
131478
74795
133016
68086
557
3076
sum of the run time for parts (a) - (e) of the experiment is graphed in Figure 3.14.
For random data, SBBSTs significantly and consistently outperformed BBSTDs and
BBSTs. On ordered data, however, BBSTDs were slightly faster than BBSTs and
both were significantly faster than SBBSTs.
Since BBSTs generated trees with the least search cost, we expect BBSTs to
outperform SBBSTs and BBSTDs in applications where the comparison cost is very
high relative to that of other operations and searches are done with a much higher
frequency than inserts and deletes. However, wfith the mix of operations used in
our tests, SBBSTs are the clear choice for random inputs and BBSTDs for ordered
inputs.
In comparing with the other structures, our tables repeat the data for BBSTs.
The reader may make the comparison with SBBSTs and BBSTDs.

81
Table 3.7. Run time on random inputs using integer keys (version 1 code)
n
operation
SBBST
BBSTD
BBST
insert
0.27
0.30
0.34
search
0.06
0.06
0.07
10,000
ins/del
0.57
0.62
0.70
search
0.06
0.06
0.06
delete
0.22
0.25
0.26
insert
1.48
1.61
1.75
search
0.35
0.36
0.37
50,000
ins/del
2.90
3.47
3.84
search
0.36
0.38
0.39
delete
1.13
1.47
1.62
insert
3.00
3.57
3.80
search
0.78
0.83
0.84
100,000
ins/del
6.28
7.78
8.41
search
0.83
0.87
0.88
delete
2.54
3.31
3.58
insert
6.56
7.74
8.37
search
1.80
1.89
1.89
200,000
ins/del
13.89
17.32
18.57
search
1.86
1.98
1.98
delete
5.64
7.41
8.02
Time Unit : sec

82
Table 3.8. Run time on ordered inputs using integer keys (version 1 code)
n
operation
SBBST
BBSTD
BBST
insert
0.32
0.20
0.27
search
0.05
0.03
0.05
10,000
ins/del
0.58
0.43
0.57
search
0.07
0.03
0.03
delete
0.20
0.17
0.23
insert
1.38
1.20
1.10
search
0.25
0.20
0.20
50,000
ins/del
2.63
2.18
2.40
search
0.25
0.20
0.20
delete
0.95
0.92
1.05
insert
3.43
2.23
2.53
search
0.72
0.45
0.42
100,000
ins/del
5.97
4.70
5.13
search
0.55
0.47
0.42
delete
2.10
1.98
2.15
insert
6.65
4.95
5.25
search
1.20
0.92
0.90
200,000
ins/del
13.13
10.23
10.88
search
1.17
0.90
0.90
delete
4.63
4.25
4.58
Time Unit : sec

83
Table 3.9. Run time on random real inputs (version 1 code)
n
operation
SBBST
BBSTD
BBST
insert
0.23
0.34
0.36
search
0.07
0.10
0.10
10,000
ins/del
0.44
0.75
0.79
search
0.08
0.10
0.10
delete
0.17
0.29
0.30
insert
1.43
1.76
1.93
search
0.47
0.53
0.52
50,000
ins/del
2.76
3.89
4.22
search
0.50
0.54
0.55
delete
1.13
1.62
1.76
insert
2.96
3.94
4.36
search
1.08
1.17
1.16
100,000
ins/del
6.11
8.58
9.30
search
1.12
1.20
1.22
delete
2.50
3.66
3.95
insert
6.85
8.92
9.33
search
2.41
2.58
2.57
200,000
ins/del
13.86
19.49
20.46
search
2.49
2.69
2.66
delete
5.61
8.25
8.80
Time Unit : sec

84
Table 3.10. Run time on ordered real inputs (version 1 code)
n
operation
SBBST
BBSTD
BBST
insert
0.27
0.23
0.20
search
0.08
0.07
0.07
10,000
ins/del
0.53
0.50
0.43
search
0.08
0.07
0.05
delete
0.18
0.23
0.20
insert
1.43
1.25
1.12
search
0.40
0.30
0.30
50,000
ins/del
2.80
2.17
2.37
search
0.40
0.30
0.30
delete
1.07
0.90
0.97
insert
3.28
2.58
2.77
search
0.90
0.62
0.63
100,000
ins/del
6.15
4.70
5.13
search
0.87
0.62
0.63
delete
2.35
1.93
2.10
insert
7.37
4.55
4.92
search
1.85
1.32
1.32
200,000
ins/del
13.35
10.03
10.93
search
1.87
1.33
1.33
delete
5.08
4.17
4.43
Time Unit : sec

85
Time is sum of time for parts (a)-(e) of the experiment
Figure 3.14. Run time on real inputs (version 1 code)
The average number of comparisons for each of the five parts of the experiment
are given in Table 3.11 for the version 1 implementation. On the comparison measure,
AVL, RB-B, WB, and BBSTs are the front runners and are quite competitive with
one another. On parts (a) (insert n elements) and (c) (insert n and delete n elements),
AVL trees performed best while on the two search tests ((b) and (d)) and the deletion
test (e), BBSTs performed best.
Table 3.12 gives the number of comparisons performed when ordered data (i.e.,
the elements in part (a) are l,2,...,n and are inserted in this order) and those in
part (c) are n + 1,... ,2n (in this order) is used instead of random permutations of
distinct elements. This experiment attempts to model realistic situations in which
the inserted elements are in “nearly sorted order”. BSTs were not included in this
test as they perform very poorly with ordered data taking 0(n2) time to insert n

86
times. The computer time needed to perform this test on BSTs was determined
to be excessive. This test exhibited greater variance in performance. Among the
deterministic structures, BBSTs outperformed the others in parts (a) - (d) while
AVL trees were ahead in part (e). For part (a), BBSTs performed approximately
45% fewer comparisons than did AVL trees and approximately 12% fewer than WB
trees. The randomized structure TRP was the best of the eight structures reported
in Table 3.12 for part (a). It performed approximately 10% fewer comparisons than
did BBST trees. However, the BBST remained best overall on parts (b), (c), and
(d).
The heights of the trees (number of levels in the case of DSL and SKIP) for
the experiments with random and ordered data are given in Tables 3.13 and 3.14
respectively. The first number in each table entry is the tree height after part (a) of
the experiment and the second, the height after part (c). In all cases, the number of
levels using skip lists is fewest. However, among the tree structures, AVL and BBST
trees have least height on random data and AVL has least with ordered data.
Tables 3.15 and 3.16, respectively, give the number of rotations performed by
each of the deterministic tree schemes for experiment parts (a), (c), and (e). Note
that none of the schemes performs rotations during a search.
On ordered data, BBSTs perform about 25% more rotations than do the re¬
maining structures. These remaining structures perform about the same number of
rotations. On random data, AVL trees, bottom-up red-black trees and WB trees per¬
form a comparable number of rotations. Top-down red-black trees and BBST trees

lable 3.11: The number of key comparisons on random inputs (version 1 code)
n
operation
BST
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
10,000
insert
search
ins/del
search
delete
264175
254175
516853
252200
215555
211401
193253
411220
193141
167312
262838
194606
515184
197399
200218
211886
194291
414990
195525
167455
211916
194153
414635
194442
167531
212111
191578
416862
191676
166487
276247
258089
923524
256578
526242
296866
258662
601137
254119
242743
224757
255072
519430
256124
231745
50,000
insert
search
ins/del
search
delete
1560958
1510958
3061868
1500504
1316917
1234911
1147273
2417733
1145808
1013535
1550701
1150466
3058045
1173662
1242426
1236968
1146754
2424944
1152764
1013144
1238628
1149970
2431281
1151578
1015988
1236114
1134969
2437639
1134062
1007688
1640660
1512093
5351715
1499657
3077266
1717037
1503452
3456045
1497081
1451835
1357076
1537547
2996512
1501731
1373858
100,000
insert
search
ins/del
search
delete
3329780
3229780
6537563
3208453
2839934
2623894
2445659
5137280
2443038
2181327
3305332
2451137
6564352
2502098
2692672
2626314
2446466
5154118
2457531
2177946
2631411
2453855
5170695
2456748
2185213
2623792
2423613
5179653
2419990
2168110
3513401
3244497
11545200
3229747
6561272
3632046
3247143
7476441
3310823
3177135
2919371
3188621
6399463
3225343
2981173
200,000
insert
search
ins/del
search
delete
7076132
6876132
13907058
6830718
6095324
5553640
5191730
10862426
5186737
4664876
7016676
5209189
13940982
5332771
5800203
5558174
5199786
10921880
5223154
4664344
5571133
5215568
10956496
5220965
4680768
5553256
5147698
10968053
5144148
4641389
7483199
6887196
24207106
6814733
13811271
7682439
6797942
15543559
6916150
6700557
6178596
6697223
13377747
6680642
6149268

Table 3.12: The number of key comparisons on ordered inputs (version 1 code)
n
operation
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
277234
376228
241383
171017
150554
435199
135989
247129
search
191917
188246
190106
188722
185530
262423
271087
256706
10,000
ins/del
421032
718040
508810
425843
314998
983676
390899
354566
search
195133
189494
190090
191681
184155
249694
269031
250538
delete
104038
276136
218216
214930
135131
468244
193080
84392
insert
1618930
2233658
1436225
995720
872967
2585557
825390
1422120
search
1120497
1117001
1120495
1117174
1101481
1509152
1540082
1467217
50,000
ins/del
2418422
4311748
3055100
2475487
1805439
6019215
2194668
1973416
search
1124001
1168633
1126126
1116390
1098065
1481819
1568903
1449810
delete
607478
1719212
1323918
1276262
791815
2785792
1181612
486498
insert
3437858
4767564
3072389
2112201
1850548
5521408
1724473
2925618
search
2390963
2383979
2390961
2384327
2354757
3218246
3564282
2970715
100,000
ins/del
5111850
9223606
6510188
5254541
3821594
12788447
4438266
4406427
search
2397971
2487243
2402224
2382759
2346128
3163554
3281308
3277089
delete
1289954
3737982
2847792
2735270
1684584
5971196
2403622
961283
insert
7275714
10135418
6544713
4465935
3903083
11743159
3428355
6403207
search
5081893
5067933
5081891
5068632
4946753
6836428
7174727
6448304
200,000
ins/del
10773706
19647336
13820364
11116226
8048058
27076911
9054078
9062233
search
5095909
5274461
5104418
5065496
5001967
6727017
7006341
6458321
delete
2729906
8075474
6095538
5836096
3577223
12741948
5094044
1995215
oo
00

89
Table 3.13. Height of the trees on random inputs (version 1 code)
n
BST
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
10,000
31,31
16,16
17,18
16,17
17,17
16,16
12,11
32,31
8,8
50,000
38,38
19,19
20,21
19,20
20,20
19,19
13,12
38,37
9,9
100,000
41,41
20,20
21,22
20,21
21,22
20,20
14,13
41,40
9,9
200,000
44,43
21,21
22,24
21,22
23,23
21,21
15,14
43,44
9,9
Table 3.14. Height of the trees on ordered inputs (version 1 code)
n
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
10,000
14,14
20,20
24,24
16,15
17,17
14,13
33,34
8,8
50,000
16,16
23,23
29,28
20,20
20,20
16,16
41,41
9,9
100,000
17,17
25,25
31,30
21,21
21,21
17,17
46,41
9,9
200,000
18,18
27,27
33,32
22,22
23,22
18,18
47,46
9,9
perform a significantly larger number of rotations. In fact, BBSTs perform about
twice as many rotations as AVL trees.
The average run times for the random data tests are given in Table 3.17 and
in Table 3.18 for the ordered data test. Both of these use integer keys. The times
using real keys are given in Tables 3.19 and 3.20. The sum of the run time for parts
(b) and (d) of the experiment is graphed in Figure 3.15 for random data and in
Figure 3.16 for ordered data. The graph of Figure 3.17 shows only one line MIX
for AVL, RB-T, RB-B, WB, and BBST while that of Figure 3.18 shows MIX for
AVL, RB-T, RB-B, and WB as the times for these are very close. With integer
keys and random data, unbalanced binary search trees (BSTs) outperformed each
of the remaining structures. The next best performance was exhibited by bottom-
up red-black trees. They did marginally better than AVL trees. The remaining

Table 3.15: The number of rotations on random inputs (version 1 code)
n
operation
AVL
RB-T
RB-B
WB
BBST
S
D
S
D
S
D
S
D
S
D
DS
DD
insert
2328
2322
1964
1955
1946
1933
2274
2065
5025
3938
151
93
10,000
ins/del
4343
3224
14773
8213
4053
2591
4256
2978
10104
5849
232
103
delete
1645
1120
9558
2678
1845
1166
1595
1022
5201
2018
51
28
insert
11664
11614
9822
9815
9710
9689
11355
10352
25059
19732
754
455
50,000
ins/del
21585
16214
81895
45180
20255
12979
21266
14975
50979
29198
1161
531
delete
8231
5630
54806
13431
9196
5844
7963
5194
26068
10033
248
131
insert
23316
23254
19593
19677
19340
19414
22723
20730
50047
39461
1527
920
100,000
ins/del
43243
32361
196769
103835
40618
25919
42567
29898
101836
58491
2275
1046
delete
16466
11264
119825
26953
18530
11708
16024
10420
51943
20147
496
260
insert
46631
46518
39290
39291
38797
38793
45458
41480
100205
79013
3054
1840
200,000
ins/del
86218
64712
394187
209941
80892
52030
84927
59911
203568
116940
4593
2059
delete
33047
22477
247905
54046
37083
23379
31984
20800
103826
40157
990
523

Table 3.16: Ihe number of rotations on ordered inputs (version 1 code)
AVL
RB-T
RB-B
WB
BBST
n
operation
S
D
S
D
S
D
S
D
S
D
DS
DD
insert
9986
0
9980
0
9976
0
9984
0
9985
2387
0
0
10,000
ins/del
14996
0
14999
0
14995
0
14997
0
16644
5797
25
154
delete
4990
0
4983
1
4989
0
4989
0
6647
3392
26
154
insert
49984
0
49977
0
49971
0
49980
0
49983
11956
0
0
50,000
ins/del
74994
0
75000
0
74994
0
74996
0
83247
28982
137
770
delete
24988
0
24978
1
24986
0
24987
0
33242
17018
136
766
insert
99983
0
99975
0
99969
0
99979
0
99983
23917
0
0
100,000
ins/del
149994
0
150000
0
149994
0
149996
0
166504
57969
280
1540
delete
49987
0
49977
1
49985
0
49986
0
66505
34040
278
1536
insert
199982
0
199973
0
199967
0
199978
0
199982
47839
0
0
200,000
ins/del
299994
0
300000
0
299994
0
299996
0
333012
115938
559
3078
delete
99986
0
99976
1
99984
0
99985
0
133016
68086
557
3076

92
Time is sum of time for parts (b) and (d) of the experiment
Figure 3.15. Run time on random real inputs (version 1 code)
structures have a noticeably inferior structure. For ordered integer keys, BSTs take
more time than we were willing to expend. Of the remaining structures, treaps
generally performed best on parts (a), (c), and (e) while BBSTs did best on parts
(b) and (d).
With real keys and random data, BSTs did not outperform the remaining struc¬
tures. Now, the five balanced binary tree structure became quite competitive with
respect to the search operations (i.e., parts (b) and (d)). RB-B generally outper¬
formed the other structures on parts (a), (c), and (e). Using ordered real keys, the
treap was the clear winner on parts (a), (c), and (e) while BBSTs handily outper¬
formed the remaining structures on parts (b) and (d).
Some of the experimental results using version 2 of the code are shown in Ta¬
bles 3.21- 3.24. On the comparison measure, with random data (Table 3.21), skip

Table 3.17: Run time on random inputs using integer keys (version 1 code)
n
operation
BST
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
0.08
0.12
0.15
0.12
0.20
0.34
0.19
0.18
0.24
search
0.05
0.05
0.05
0.06
0.05
0.07
0.09
0.09
0.18
10,000
ins/del
0.14
0.21
0.36
0.22
0.39
0.70
0.49
0.33
0.45
search
0.05
0.05
0.05
0.05
0.05
0.06
0.09
0.09
0.18
delete
0.05
0.08
0.12
0.09
0.16
0.26
0.20
0.08
0.16
insert
0.65
0.79
0.98
0.73
1.18
1.75
1.10
1.01
1.36
search
0.40
0.36
0.36
0.36
0.35
0.37
0.58
0.56
1.25
50,000
ins/del
1.04
1.48
2.50
1.26
2.22
3.84
2.77
1.86
2.73
search
0.40
0.41
0.44
0.36
0.36
0.39
0.57
0.56
1.16
delete
0.39
0.54
1.01
0.51
0.94
1.62
1.16
0.51
1.10
insert
1.34
1.57
2.10
1.54
2.54
3.80
2.46
2.23
2.84
search
0.88
0.80
0.80
0.83
0.78
0.84
1.36
1.30
2.63
100,000
ins/del
2.36
3.21
5.52
2.74
4.86
8.41
6.35
4.10
6.13
search
0.93
0.94
1.00
0.84
0.83
0.88
1.33
1.29
2.61
delete
0.88
1.24
2.26
1.14
2.11
3.58
2.64
1.23
2.41
insert
2.79
3.37
4.41
3.18
5.21
8.37
5.56
4.70
6.25
search
2.00
1.80
1.81
1.81
1.78
1.89
3.03
2.91
5.85
200,000
ins/del
5.24
6.99
12.51
5.99
10.54
18.57
14.29
8.95
13.29
search
2.08
2.12
2.25
1.91
1.87
1.98
3.04
2.93
5.81
delete
2.01
2.69
5.06
2.51
4.55
8.02
5.84
2.76
5.35
Time Unit : sec

Table 3.18. Run time on ordered inputs using integer keys (version 1 code)
n
operation
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
0.12
0.17
0.12
0.18
0.27
0.23
0.08
0.20
search
0.05
0.03
0.03
0.07
0.05
0.07
0.05
0.12
10,000
ins/del
0.18
0.32
0.20
0.35
0.57
0.42
0.17
0.20
search
0.05
0.05
0.05
0.05
0.03
0.07
0.05
0.13
delete
0.05
0.10
0.07
0.13
0.23
0.15
0.05
0.07
insert
0.75
1.02
0.92
1.25
1.10
0.98
0.47
0.92
search
0.32
0.27
0.27
0.28
0.20
0.33
0.32
0.62
50,000
ins/del
1.28
2.17
1.25
2.20
2.40
2.03
0.80
1.07
search
0.28
0.28
0.27
0.28
0.20
0.30
0.37
0.62
delete
0.30
0.75
0.37
0.85
1.05
0.65
0.30
0.27
insert
1.50
2.52
1.70
2.58
2.53
2.58
0.90
1.72
search
0.70
0.60
0.57
0.70
0.42
0.70
0.63
1.23
100,000
ins/del
2.60
4.68
2.53
4.78
5.13
4.42
1.52
2.43
search
0.63
0.60
0.55
0.62
0.42
0.70
0.58
1.35
delete
0.62
1.65
0.78
1.87
2.15
1.42
0.45
0.55
insert
3.12
4.82
3.38
5.67
5.25
4.72
1.80
3.52
search
1.38
1.30
1.22
1.33
0.90
1.60
1.25
2.70
200,000
ins/del
5.15
10.40
5.35
10.40
10.88
9.48
3.10
5.13
search
1.33
1.33
1.18
1.32
0.90
1.50
1.28
2.72
delete
1.35
3.63
1.68
4.12
4.58
2.98
0.93
1.12
Time Unit : sec

Table 3.19: Run time on random real inputs (version 1 code)
n
operation
BST
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
0.14
0.15
0.21
0.17
0.23
0.36
0.22
0.23
0.30
search
0.09
0.07
0.09
0.10
0.08
0.10
0.13
0.13
0.21
10,000
ins/del
0.24
0.27
0.51
0.32
0.38
0.79
0.62
0.41
0.53
search
0.09
0.08
0.09
0.10
0.08
0.10
0.12
0.12
0.21
delete
0.09
0.09
0.17
0.14
0.14
0.30
0.28
0.11
0.19
insert
0.94
0.97
1.22
0.86
1.29
1.93
1.48
1.19
1.67
search
0.64
0.52
0.50
0.51
0.51
0.52
0.87
0.71
1.44
50,000
ins/del
1.68
1.77
2.74
1.53
2.29
4.22
3.93
2.17
3.15
search
0.66
0.55
0.56
0.54
0.56
0.55
0.86
0.71
1.33
delete
0.63
0.67
1.10
0.72
0.92
1.76
1.80
0.69
1.22
insert
2.06
1.85
2.34
1.90
2.66
4.36
3.05
2.67
3.61
search
1.43
1.13
1.09
1.13
1.14
1.16
1.84
1.66
3.00
100,000
ins/del
3.63
3.93
6.18
3.33
4.96
9.30
8.45
4.84
7.10
search
1.45
1.26
1.27
1.17
1.26
1.22
1.83
1.65
3.01
delete
1.39
1.50
2.51
1.55
2.03
3.95
3.91
1.61
2.75
insert
4.34
3.95
5.20
3.88
5.56
9.33
6.77
5.81
7.90
search
3.19
2.49
2.42
2.50
2.45
2.57
4.14
3.67
6.62
200,000
ins/del
8.01
8.25
13.78
7.29
10.65
20.46
18.88
10.48
15.83
search
3.21
2.83
2.86
2.62
2.74
2.66
4.08
3.73
6.74
delete
3.11
3.27
5.55
3.41
4.43
8.80
8.56
3.54
6.04
Time Unit : sec

96
Table 3.20. Run time on ordered real inputs (version 1 code)
n
operation
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
0.13
0.22
0.15
0.25
0.20
0.25
0.12
0.30
search
0.07
0.08
0.07
0.07
0.07
0.10
0.07
0.15
10,000
ins/del
0.23
0.42
0.27
0.40
0.43
0.47
0.18
0.28
search
0.07
0.05
0.08
0.08
0.05
0.08
0.08
0.12
delete
0.07
0.17
0.08
0.15
0.20
0.20
0.05
0.07
insert
1.15
1.58
1.12
1.85
1.12
1.30
0.67
1.35
search
0.42
0.42
0.43
0.40
0.30
0.53
0.38
0.82
50,000
ins/del
1.28
2.75
1.57
2.57
2.37
3.02
0.92
1.40
search
0.40
0.42
0.42
0.48
0.30
0.53
0.40
0.75
delete
0.38
0.95
0.55
0.93
0.97
1.15
0.33
0.35
insert
1.77
3.23
2.12
3.35
2.77
3.13
1.17
2.42
search
0.90
0.87
0.90
0.88
0.63
1.12
0.92
1.70
100,000
ins/del
3.00
6.00
3.42
5.38
5.13
6.32
1.92
3.22
search
0.97
0.92
0.88
0.98
0.63
1.12
0.82
1.70
delete
0.87
2.08
1.17
2.05
2.10
2.40
0.70
0.67
insert
3.92
6.42
4.27
7.25
4.92
6.03
2.58
4.93
search
1.92
1.87
1.92
1.88
1.32
2.40
1.85
3.87
200,000
ins/del
5.78
13.80
7.33
11.88
10.93
13.72
3.75
6.67
search
1.90
1.93
1.92
2.13
1.33
2.38
1.75
3.97
delete
1.67
4.55
2.48
4.45
4.43
5.10
1.40
1.35
Time Unit : sec

97
Time is sum of time for parts (b) and (d) of the experiment
Figure 3.16. Run time on ordered real inputs (version 1 code)
lists performed best on part (a). Of the deterministic methods, BBSTs slightly out¬
performed the others on part (a). On parts (b) - (e), AVL, RB-T, RB-B, WB, and
BBSTs were quite competitive and outperformed BSTs and the randomized schemes.
BBSTs performed best on parts (b) and (d), RB-Ts did best on part (e) and RB-B
and AVL did best on part (c). In comparing the results of Table 3.21 to those of
Table 3.11 (using version 1 code), we see that the change to version 2 generally in¬
creased the comparison cost of the deterministic tree structures by about 25%. For
the DSL, the change in code had mixed results. Notice that for RB-T and DSLs,
the comparison count for parts (a), (c), and (e) are the same as for the version 1
code. This is because for inserts and deletes, it is necessary to do the equal check
first when using these structures. For SKIPs the count is the same for all five parts
as the version 1 and 2 codes are the same.

98
With ordered data (Table 3.22), treaps required the fewest comparisons for part
(a). Skip lists did best on parts (c) and (e), and AVL trees generally outperformed
the other structures on parts (b) and (d). Once again, the comparison counts were
generally higher using the version 2 code than using the version 1 code.
Run time data using real keys is given in Tables 3.23 and 3.24. The sum of the
run time for parts (b) and (d) of the experiment is graphed in Figure 3.17 for random
data and in Figure 3.18 for ordered data. The graph of Figure 3.17 shows only one
line MIX for AVL, RB-T, RB-B, WB, and BBST while that of Figure 3.18 shows
MIX for AVL, RB-T, RB-B, and WB as the times for these are very close. With
random data, RB-B generally performed best on part (a), on parts (b) and (d), the
front runner varied among AVL, RB-T, and WB, and on parts (c) and (e) RB-Bs
generally did best. On ordered data, TRPs did best on paxts (a), (c), and (e) while
BBSTs did best on parts (b) and (d).
3.8 Conclusion
We have developed a new weight balanced data structure called /9-BBST. This
was developed for the representation of a dictionary. In developing the insert/delete
algorithms, we sought to minimize the search cost of the resulting tree. Our experi¬
mental results show that BBSTs generally have the best search cost of the structures
considered. Furthermore, this translates into reduced search time when the key com¬
parison cost is relatively high (e.g., for real keys). The insert and delete algorithms
for /?-BBSTs are not as efficient as those for other dictionary structures (such as
AVL trees). As a result, we recommend /TBBSTs for environments where searches

Table 3.21: The number of key comparisons on random inputs (version 2 code)
n
operation
BST
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
332753
262198
262838
262726
263177
260896
276247
375698
224757
search
322753
241557
242258
242262
242824
240126
348403
329411
255072
10,000
ins/del
650901
514371
515184
513920
515732
518124
923524
755629
519430
search
318749
241536
247130
243191
242867
240126
335613
320612
256124
delete
271004
206558
200218
206721
207622
207210
526242
300619
231745
insert
1983939
1546988
1550701
1549795
1554520
1539666
1640660
2184066
1357076
search
1933939
1443879
1447870
1446679
1452920
1435927
2043618
1921255
1537547
50,000
ins/del
3892221
3043090
3058045
3040654
3055092
3061443
5351715
4393520
2996512
search
1913068
1443837
1476158
1451163
1452625
1435726
1969926
1909919
1501731
delete
1674128
1267637
1242426
1268881
1275612
1270935
3077266
1815736
1373858
insert
4245062
3297162
3305332
3302792
3314410
3281959
3513401
4637264
2919371
search
4145062
3090057
3098143
3096011
3111095
3074661
4387427
4161175
3188621
100,000
ins/del
8336846
6490752
6564352
6486464
6520729
6528606
11545200
9484761
6399463
search
4102672
3089826
3176862
3105465
3110184
3074305
4270168
4224698
3225343
delete
3623179
2738267
2692672
2740846
2756006
2744369
6561272
4008111
2981173
insert
9045367
6999791
7016676
7012317
7040203
6969465
7483199
9834444
6178596
search
8845367
6584279
6603044
6599643
6633218
6554714
9373163
8752856
6697223
200,000
ins/del
17782478
13790643
13940982
13789492
13862467
13867876
24207106
19825904
13377747
search
8757433
6585758
6747566
6618833
6630334
6554354
8995685
8889053
6680642
delete
7800524
5882302
5800203
5889983
5923552
5893982
13811271
8456931
6149268
SO

Table 3.22: The number of key comparisons on ordered inputs (version 2 code)
n
operation
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
267234
376228
442766
302034
261108
435199
216958
247129
search
237262
239442
247706
237298
243110
372444
332060
256706
10,000
ins/del
493028
718040
727620
562808
558770
983676
482499
354566
search
240910
238834
246330
238320
242736
349730
344982
250538
delete
178076
276136
208216
204930
239028
468244
183080
84392
insert
1568930
2233658
2672450
1791440
1545934
2585557
1375770
1422120
search
1418962
1421560
1459588
1420858
1455936
2159176
1990474
1467217
50,000
ins/del
2881762
4311748
4360200
3301668
3251450
6019215
2742877
1973416
search
1419154
1444824
1444258
1424442
1452494
2131862
1956194
1449810
delete
1064956
1719212
1273918
1226262
1427504
2785792
1131612
486498
insert
3337858
4767564
5744778
3824402
3301096
5521408
2898893
2925618
search
3037892
3043084
3119128
3041676
3121098
4618272
4538718
2970715
100,000
ins/del
6113530
9223606
9320376
7027676
6930932
12788447
5492066
4406427
search
3038276
3089612
3088470
3048844
3114012
4563600
4158994
3277089
delete
2279908
3737982
2747792
2635270
3056908
5971196
2303622
961283
insert
7075714
10135418
12289426
8131870
7006166
11743159
5756575
6403207
search
6475750
6486128
6638204
6483310
6646168
9836456
9102954
6448304
200,000
ins/del
12927066
19647336
19840728
14904040
14671602
27076911
11290926
9062233
search
6476518
6579184
6576890
6497646
6634260
9727066
8918638
6458321
delete
4859812
8075474
5895538
5636096
6529928
12741948
4894044
1995215

Table 3.23: Run time on random real inputs (version 2 code)
n
operation
BST
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
0.15
0.14
0.20
0.18
0.25
0.36
0.23
0.25
0.31
search
0.10
0.08
0.10
0.11
0.09
0.11
0.13
0.16
0.21
10,000
ins/del
0.27
0.27
0.52
0.34
0.47
0.80
0.64
0.50
0.54
search
0.10
0.08
0.10
0.11
0.09
0.11
0.13
0.14
0.21
delete
0.10
0.10
0.20
0.14
0.18
0.32
0.29
0.14
0.19
insert
1.02
0.98
1.15
0.89
1.46
1.88
1.44
1.34
1.65
search
0.69
0.55
0.57
0.55
0.57
0.55
0.89
0.83
1.42
50,000
ins/del
1.79
1.80
2.99
1.59
2.93
3.97
3.82
2.44
3.16
search
0.71
0.60
0.63
0.55
0.57
0.56
0.87
0.79
1.32
delete
0.67
0.67
1.22
0.66
1.19
1.63
1.80
0.75
1.21
insert
2.15
2.00
2.58
1.90
3.18
4.01
3.11
2.95
3.69
search
1.52
1.21
1.24
1.18
1.23
1.23
1.97
1.84
3.04
100,000
ins/del
3.88
3.92
6.74
3.46
6.28
8.73
8.50
5.39
7.18
search
1.55
1.32
1.45
1.25
1.29
1.27
1.95
1.82
2.98
delete
1.51
1.49
2.75
1.45
2.57
3.64
3.93
1.73
2.77
insert
5.04
4.45
5.79
4.28
6.92
9.20
7.05
6.81
8.01
search
3.43
2.63
2.70
2.64
2.73
2.69
4.43
4.00
6.60
200,000
ins/del
8.92
8.87
15.36
7.88
13.85
19.53
19.55
12.17
16.11
search
3.43
2.98
3.13
2.73
2.83
2.77
4.37
4.02
6.70
delete
3.33
3.32
6.08
3.20
5.65
8.24
8.91
3.88
6.04
Time Unit : sec

102
Time is sum of time for parts (b) and (d) of the experiment
Figure 3.17. Run time on random real inputs (version 2 code)
Time is sum of time for parts (b) and (d) of the experiment
Figure 3.18. Run time on ordered real inputs (version 2 code)

103
Table 3.24. Run time on ordered real inputs (version 2 code)
n
operation
AVL
RB-T
RB-B
WB
BBST
DSL
TRP
SKIP
insert
0.17
0.23
0.28
0.27
0.30
0.23
0.15
0.30
search
0.08
0.08
0.12
0.08
0.08
0.12
0.12
0.13
10,000
ins/del
0.23
0.43
0.40
0.47
0.60
0.48
0.17
0.27
search
0.08
0.08
0.07
0.08
0.08
0.08
0.10
0.13
delete
0.08
0.J5
0.12
0.17
0.20
0.20
0.08
0.05
insert
0.83
1.45
1.43
1.57
1.37
1.35
0.82
1.18
search
0.45
0.48
0.48
0.47
0.38
0.60
0.50
0.83
50,000
ins/del
1.35
2.65
1.95
2.75
2.47
3.05
1.05
1.42
search
0.45
0.47
0.45
0.47
0.37
0.63
0.58
0.77
delete
0.45
1.05
0.50
1.00
1.03
1.17
0.43
0.33
insert
1.78
2.75
2.73
3.43
2.63
3.23
1.33
2.18
search
0.97
0.98
1.00
1.03
0.77
1.30
1.15
1.55
100,000
ins/del
2.85
6.22
3.98
6.00
5.33
6.37
2.02
3.33
search
0.97
1.10
0.98
1.02
0.77
1.32
1.03
1.70
delete
0.97
2.18
1.05
2.15
2.22
2.43
0.63
0.67
insert
3.78
6.08
5.43
7.18
5.37
6.07
2.87
5.23
search
2.08
2.13
2.13
2.17
1.63
3.10
2.27
3.47
200,000
ins/del
6.13
13.93
8.48
13.42
11.33
13.60
4.10
7.02
search
2.12
2.15
2.13
2.17
1.63
2.80
2.18
4.27
delete
2.03
4.75
2.27
4.77
4.72
5.18
1.35
1.35
Time Unit : sec

104
are done with much greater frequency than inserts and/or deletes. Based on our
experiments, we conclude that AVL trees remain the best dictionary structure for
general applications.
We have also proposed two simplified versions of the BBST called SBBST and
BBSTD. The SBBST seeks only to provide logarithmic run time per operation and
unlike the general BBST, does not reduce search cost at every opportunity. The
SBBST provides slightly better balance than provided by WB(a) trees. The BBSTD
does not attempt to maintain /2-balance. However it performs rotations to reduce
search cost whenever possible. Both versions are very competitive with BBSTs. The
SBBST exhibited much better run time performance than BBSTs on random data
and the BBSTD slightly outperformed the BBST on ordered data. However, BBSTs
generated trees with the lowest search cost (though not by much).

CHAPTER 4
WEIGHT BIASED LEFTIST TREES AND MODIFIED SKIP LISTS
4.1 Introduction
Several data structures (e.g., heaps, leftist trees [9], binomial heaps [10]) have
been proposed for the representation of a (single ended) priority queue. Heaps permit
one to delete the min element and insert an arbitrary element into an n element
priority queue in O(logn) time. Leftist trees support both these operations and the
merging of pairs of priority queues in logarithmic time. Using binomial heaps, inserts
and combines take 0(1) time and a delete-min takes O(logn) amortized time. In
this chapter, we begin in Section 4.2, by developing the weight biased leftist tree.
This is similar to a leftist tree. However biasing of left and right subtrees is done
by number of nodes rather than by length of paths. Experimental results presented
in Section 4.5 show that weight biased leftist trees provide better performance than
provided by leftist trees. The experimental comparisons of Section 4.5 also include a
comparison with heaps and binomial heaps as well as with unbalanced binary search
trees and the probabilistic structures treap [1] and skip lists [26].
In Section 4.3, we propose a fixed node size representation for skip lists. The
new structure is called modified skip lists and is experimentally compared with the
105

106
variable node size structure skip lists. Our experiments indicate that modified skip
lists are faster than skip lists when used to represent dictionaries.
Modified skip lists are augmented by a thread in Section 4.4 to obtain a structure
suitable for use as a priority queue. For completeness, we include, in Section 4.5, a
comparison of data structures for double ended priority queues.
4.2 Weight Biased Leftist Trees
Let T be an extended binary tree. For any internal node x of T, let LeftChild(x)
and RightChild(x), respectively, denote the left and right children of x. The weight,
w(x), of any node x is the number of internal nodes in the subtree with root x.
The length, shortest(x), of a shortest path from x to an external node satisfies the
recurrence
shortest^ x) — <
0 if x is an external node
1 + min{shortest(LeftChild(x)),shortest(RightChild(x))} otherwise.
Definition [9] A leftist tree (LT) is a binary tree such that if it is not empty, then
shortest(LeftChild(x)) > shortest(RightChild(x))
for every internal node x.
A weight biased leftist tree (WBLT) is defined by using the weight measure in
place of the measure shortest.

107
Definition A weight biased, leftist tree (WBLT) is a binary tree such that if it is
not empty, then
weight(LeftChild(x)) > weight(RightChild(x))
for every internal node x.
It is known [9] that the length, rightmost(x), of the rightmost root to external
node path of any subtree, x, of a leftist tree satisfies
rightmost(x ) < log2M*) + !)•
The same is true for weight biased leftist trees.
Theorem 13 Let x be any internal node of a weight biased leftist tree, rightmost(x) <
log2(tt;(x) + 1).
Proof The proof is by induction on w(x). When w(x) = 1, rightmost(x) =
1 and log2(u;(x) + 1) = log2 2 = 1. For the induction hypothesis, assume that
rightmost(x) < log2(u>(x)+l) whenever w(x) < n. When w(x) = n, w(RightChild(x)) <
(n — l)/2 and rightmost(x) — 1 + rightmost(RightChild(x)) < 1 + log2((n — l)/2 +
1) = 1 + log2(n + 1) - 1 = log2(n + 1). â–¡
Definition A min (max)-WBLT is a WBLT that is also a min (max) tree.
Each node of a min-WBLT has the fields: Isize (number of internal nodes
in left subtree), rsize, left (pointer to left subtree), right, and data. While the

108
bottom
(a) Empty min-WBLT (b) Nonempty min-WBLT
Figure 4.1. Example min-WBLTs
number of size fields in a node may be reduced to one, two fields result in a faster
implementation. We assume a head node head with Isize = oo and ¡child = head. In
addition, a bottom node bottom with data.key = oo. All pointers that would normally
be nil are replaced by a pointer to bottom. Figure 4.1(a) shows the representation of
an empty min-WBLT and Figure 4.1(b) shows an example non empty min-WBLT.
Notice that all elements are in the right subtree of the head node.
Min (max)-WBLTs can be used as priority queues in the same way as min (max)-
LTs. For instance, a min-WBLT supports the standard priority queue operations
of insert and delete-min in logarithmic time. In addition, the combine operation
(i.e., join two priority queues together) can also be done in logarithmic time. The
algorithms for these operations have the same flavor as the corresponding ones for
min-LTs. A high level description of the insert and delete-min algorithm for min-
WBLT is given in Figures 4.2 and 4.3, respectively. The algorithm to combine two

109
procedure Insert(d) ;
{insert d into a min-WBLT}
begin
create a node x with x.data = d ;
t = head ; {head node}
while (t.right.data.key < d.key) do
begin
t.rsize = t.rsize + 1 ;
if (t.lsize < t.rsize) then
begin swap Vs children ; t = t.left ; end
else t = t.right ;
end ;
x.left = t.right ; x.right = bottom ;
x.lsize = t.rsize ; x.rsize = 0 ;
if (t.lsize = t.rsize) then {swap children}
begin
t.right = t.left ;
t.left = x ; t.lsize = x.lsize + 1 ;
end
else
begin t.right = x ; t.rsize = t.rsize + 1 ; end ;
end ;
Figure 4.2. min-WBLT Insert
min-WBLTs is similar to the delete-min algorithm. The time required to perform
each of the operations on a min-WBLT T is 0(rightmost(T)).
Notice that while the insert and delete-min operations for min-LTs require a
top-down pass followed by a bottom-up pass, these operations can be performed by
a single top-down pass in min-WBLTs. Hence, we expect min-WBLTs to outperform
min-LTs.
4.3 Modified Skip Lists
Skip lists were proposed in Pugh [26] as a probabilistic solution for the dictionary
problem (i.e., represent a set of keys and support the operations of search, insert,
and delete). The essential idea in skip lists is to maintain upto Imax ordered chains

procedure Delete-min ;
begin
x = head.right ;
if (x = bottom) then return ; {empty tree}
head.right = x.left ; head.rsize = x.lsize ;
a = head;
b = x.right ; 6s¿ze = x.rsize ;
delete x ;
if (i> = bottom) then return ;
r = a.right ;
while (r ^ bottom) do
begin
5 = ¿size + a.rsize ; t = a.rsize ;
if (a.lsize < 5) then {work on a.left)
begin
a.right = a.left ; a.rsize — a.lsize ; a.lsize = s ;
if (r.data.key > b.data.key) then
begin a.left = b ; a = b ; 6 = r ; bsize = t ; end
else
begin a.left = r ; a = r ; end
end
else
do symmetric operations on a.right ;
r = a.right ;
end ;
if (a.lsize < bsize) then
begin
a.right = a.left ; a.left = b ;
a.rsize = a.lsize ; a.lsize = bsize ;
end
else
begin a.right = b ; a.rsize = bsize ; end ;
end ;
Figure 4.3. min-WBLT Delete-min

Ill
level
Figure 4.4. Skip Lists
designated as level 1 chain, level 2 chain, etc. If we currently have ¡current number
of chains, then all n elements of the dictionary are in the level 1 chain and for each
/, 2 < / < /current, approximately a fraction p of the elements on the level / — 1
chain are also on the level / chain. Ideally, if the level / — 1 chain has m elements
then the approximately m x p elements on the level / chain are about 1/p apart in
the level / — 1 chain. Figure 4.4 shows an ideal situation for the case ¡current = 4
and p = 1/2.
While the search, insert, and delete algorithms for skip lists are simple and have
probabilistic complexity O(logn) when the level 1 chain has n elements, skip lists
suffer from the following implementational drawbacks:
1. In programming languages such as Pascal, it isn’t possible to have variable size
nodes. As a result, each node has one data field, and ¡max pointer fields. So,
the n element nodes have a total of n x ¡max pointer fields even though only
about n/( 1 —p) pointers are necessary. Since ¡max is generally much larger than
3 (the recommended value is log^^nMax where nMax is the largest number of
elements expected in the dictionary), skip lists require more space than WBLTs.

112
2. While languages such as C and C++ support variable size nodes and we can
construct variable size nodes using simulated pointers [27] in languages such as
Pascal that do not support variable size nodes, the use of variable size nodes
requires more complex storage management techniques than required by the
use of fixed size nodes. So, greater efficiency can be achieved using simulated
pointers and fixed size nodes.
With these two observations in mind, we propose a modified skip list (MSL)
structure in which each node has one data field and three pointer fields: left, right,
and down. Notice that this means MSLs use four fields per node while WBLTs use
five (as indicated earlier this can be reduced to four at the expense of increased run
time). The left and right fields are used to maintain each level / chain as a doubly
linked list and the down field of a level / node x points to the leftmost node in the level
/ — 1 chain that has key value larger than the key in x. Figure 4.5 shows the modified
skip list that corresponds to the skip list of Figure 4.4. Notice that each element is
in exactly one doubly linked list. We can reduce the number of pointers in each node
to two by eliminating the field left and having down point one node the left of where
it currently points (except for head nodes whose down fields still point to the head
node of the next chain). However, this results in a less time efficient implementation.
H and T, respectively, point to the head and tail of the level /current chain.
A high level description of the algorithms to search, insert, and delete are given in
Figures 4.6, 4.7, and 4.8. The next theorem shows that their probabilistic complexity
is O(log n) where n is the total number of elements in the dictionary.

113
level H
T
Figure 4.5. Modified Skip Lists
procedure Search(key) ;
begin
P-H;
while (p / nil) do
begin
while (p.data.key < key) do
p = p.right ;
if (p.data.key = key) then report and stop
else p = p.left.down ; {1 level down}
end ;
end ;
Figure 4.6. MSL Search

114
procedure Insert(d) ;
begin
randomly generate the level k at which d is to be inserted ;
search the MSL H for d.key saving information useful for insertion ;
if d.key is found then fail ; {duplicate}
get a new node x and set x.data = d ;
if ((k > leurrent) and (/current ^ Imax)) then
begin
leurrent = ¡current + 1 ;
create a new chain with a head node, node x, and a tail and
connect this chain to H ;
update H ;
set x.down to the appropriate node in the level leurrent — 1 chain
(to nil if k = 1) ;
end
else
begin
insert x into the level k chain ;
set x.down to the appropriate node in the level k — 1 chain
(to nil if k = 1) ;
update the down field of nodes on the level k + 1 chain (if any) as needed ;
end ;
end ;
Figure 4.7. MSL Insert
procedure Delete(z) ;
begin
search the MSL H for a node x with data.key = z saving information useful
for deletion;
if not found then fail ;
let k be the level at which z is found ;
for each node p on level k + 1 that has p.down = x, set p.down = x.right ;
delete x from the level k list ;
if the list at level leurrent becomes empty then
delete this and succeeding empty lists until we reach the first non empty list,
update leurrent ;
end ;
Figure 4.8. MSL Delete

115
Theorem li The probabilistic complexity of the MSL operations is Ofiog n).
Proof We establish this by showing that our algorithms do at most a logarithmic
amount of additional work than do those of Pugh [26]. Since the algorithms of Pugh
[26] has probabilistic O(logn) complexity, so also do ours. During a search, the extra
work results from moving back one node on each level and then moving down one level.
When this is done from any level other than /current, we expect to examine upto
c = 1/p— 1 additional nodes on the next lower level. Hence, upto c(lcurrent — 2) addi¬
tional nodes get examined. During an insert, we also need to verify that the element
being inserted isn’t one of the elements already in the MSL. This requires an addi¬
tional comparison at each level. So, MSLs may make upto c(lcurrent — 2) -f leurrent
additional compares during an insert. The number of down pointers that need to
be changed during an insert or delete is expected to be ]C¿^i *p' = (Tl^p- Since c
and p are constants and Imax — logj/p n, the expected additional work is O(log n). □
The relative performance of skip lists and modified skip lists as a data structure
for dictionaries was determined by programming the two in C. Both were implemented
using simulated pointers. The simulated pointer implementation of skip lists used
fixed size nodes. This avoided the use of complex storage management methods and
biased the run time measurements in favor of skip lists. For the case of skip lists, we
used p = 1/4 and for MSLs, p = 1/5. These values of p were found to work best for
each structure. Imax was set to 16 for both structures.

116
We experimented with n = 10,000, 50,000, 100,000, and 200,000. For each n,
the following five part experiment was conducted:
(a) start with an empty structure and perform n inserts;
(b) search for each item in the resulting structure once; items are searched for in the
order they were inserted
(c) perform an alternating sequence of n inserts and n deletes; in this, the n elements
inserted in (a) are deleted in the order they were inserted and n new elements are
inserted
(d) search for each of the remaining n elements in the order they were inserted
(e) delete the n elements in the order they were inserted.
For each n, the above five part experiment was repeated ten times using different
random permutations of distinct elements. For each sequence, we measured the total
number of element comparisons performed and then averaged these over the ten
sequences. The average number of comparisons for each of the five parts of the
experiment are given in Table 4.1.
Also given in this table is the number of comparisons using ordered data. For
this data set, elements were inserted and deleted in the order 1,2,3,.... For the
case of random data, MSLs make 40% to 50% more comparisons on each of the five
parts of the experiment. On ordered inputs, the disparity is even greater with MSLs
making 30% to 140% more comparison. Table 4.2 gives the number of levels in SKIP
and MSL. The first number of each entry is the number of levels following part (a)
of the experiment and the second the number of levels following part (b). As can be

117
Table 4.1. The number of key comparisons
n
operation
random inputs
ordered inputs
SKIP
MSL
SKIP
MSL
insert
224757
322499
247129
318854
search
255072
362865
256706
339019
10,000
ins/del
519430
734161
354566
560219
search
256124
349591
250538
339121
delete
231745
320594
84392
185489
insert
1357076
1950583
1422120
1911818
search
1537547
1965649
1467217
1836713
50,000
ins/del
2996512
4142186
1973416
3204400
search
1501731
2038774
1449810
1989550
delete
1373858
1853671
486498
931975
insert
2919371
4146428
2925618
4275880
search
3188621
4315576
2970715
4082193
100,000
ins/del
6399463
9103135
4406427
6895510
search
3225343
4427979
3277089
4345874
delete
2981173
4161994
961283
2052638
insert
6178596
8927523
6403207
9022631
search
6697223
9273707
6448304
8946474
200,000
ins/del
13377747
19370831
9054078
9062233
search
6680642
9662006
6458321
9197714
delete
6149268
9101721
1995215
4837867

118
Table 4.2. Number of levels
n
random inputs
ordered inputs
SKIP
MSL
SKIP
MSL
10,000
8,8
7,7
8,8
7,7
50,000
9,9
7,7
9,9
7,7
100,000
9,9
7,8
9,9
7,8
200,000
9,9
8,9
9,9
8,9
seen, the number of levels is very comparable for both structures. MSLs generally
had one or two levels fewer than SKIPs had.
Despite the large disparity in number of comparisons, MSLs generally required
less time than required by SKIPs (see Table 4.3 and Figure 4.9). Integer keys were
used for our run time measurements. In many practical situations the observed time
difference will be noticeably greater as one would need to code skip lists using more
complex storage management techniques to allow for variable size nodes.
4.4 MSLs As Priority Queues
At first glance, it might appear that skip lists are clearly a better choice than
modified skip lists for use as a priority queue. The min element in a skip list is
the first element in the level one chain. So, it can be identified in 0(1) time and
then deleted in O(logn) probabilistic time. In the case of MSLs, the min element
is the first one in one of the leurrent chains. This can be identified in logarithmic
time using a loser tree whose elements are the first element from each MSL chain.
By using an additional pointer field in each node, we can thread the elements in an
MSL into a chain. The elements appear in non-decending order on this chain. The

119
Table 4.3. Run time
n
operation
random inputs
ordered inputs
SKIP
MSL
SKIP
MSL
insert
0.24
0.18
0.20
0.17
search
0.18
0.12
0.12
0.07
10,000
ins/del
0.45
0.35
0.20
0.20
search
0.18
0.12
0.13
0.07
delete
0.16
0.12
0.07
0.05
insert
1.36
1.22
0.92
0.80
search
1.25
0.98
0.62
0.38
50,000
ins/del
2.73
2.53
1.07
1.08
search
1.16
1.00
0.62
0.42
delete
1.10
0.83
0.27
0.23
insert
2.84
2.86
1.72
1.60
search
2.63
2.39
1.23
0.85
100,000
ins/del
6.13
5.80
2.43
2.28
search
2.61
2.33
1.35
0.92
delete
2.41
2.02
0.55
0.52
insert
6.25
6.49
3.52
3.47
search
5.85
5.34
2.70
1.87
200,000
ins/del
13.29
13.02
5.13
4.75
search
5.81
5.51
2.72
1.92
delete
5.35
4.85
1.12
1.18

120
Time is sum of time for parts (a)-(e) of the experiment
Figure 4.9. Run time
resulting threaded structure is referred to as TMSL (threaded modified skip lists).
A delete min operation can be done in 0(1) expected time when a TMSL is used.
The expected time for an insert remains O(logn). The algorithms for the insert and
delete min operations for TMSLs are given in Figures 4.10 and 4.11, respectively.
The last step of Figure 4.10 is implemented by first finding the largest element on
level 1 with key < d.key (for this, start at level leurrent — 1) and then follow the
threaded chain.
Theorem 15 The expected complexity of an insert and delete-min operation in a
TMSL is O(\ogn) and 0(1), respectively.
Proof Follows from the notion of a thread, Theorem 14, and Pugh [26]. â–¡

121
procedure Insert(d) ;
begin
randomly generate the level k at which d is to be inserted ;
get a new node x and set x.data = d ;
if ((* > lcurrent) and (/current ^ Imax)) then
begin
leurrent = /current + 1 ;
create a new chain with a head node, node x, and a tail and
connect this chain to H ;
update H ;
set x.down to the appropriate node in the level leurrent — 1 chain
(to nil if k = 1) ;
end
else
begin
insert x into the level k chain ;
set x.down to the appropriate node in the level k— 1 chain (to nil if k = 1) ;
update the down field of nodes on the level k + 1 chain (if any) as needed ;
end ;
find node with largest key < d.key and insert x into threaded list ;
end ;
Figure 4.10. TMSL Insert
procedure Delete-min ;
begin
delete the first node x from the thread list ;
let k be the level x is on ;
delete x from the level k list (note there are no down fields on level k + 1
that need to be updated) ;
if the list at level leurrent becomes empty then
delete this and succeeding empty lists until we reach the first non empty list,
update leurrent ;
end ;
Figure 4.11. TMSL Delete-min

122
procedure Delete-max ;
begin
delete the last node x from the thread list ;
let k be the level x is on ;
delete x from the level k list updating p.down for nodes on level k-\-1 as necessary;
if the list at level ¡current becomes empty then
delete this and succeeding empty lists until we reach the first non empty list,
update /current ;
end ;
Figure 4.12. TMSL Delete-max
TMSLs may be further extended by making the threaded chain a doubly finked
fist. This permits both delete-min and delete-max to be done in 0(1) expected time
and insert in O(logn) expected time. With this extension, TMSLs may be used to
represent double ended priority queues.
4.5 Experimental Results For Priority Queues
The single-ended priority queue structures min heap (Heap), binomial heap
(B-Heap), leftist trees (LT), weight biased leftist trees (WBLT), and TMSLs were
programmed in C. In addition, priority queue versions of unbalanced binary search
trees (BST), AVL trees, treaps (TRP), and skip fists (SKIP) were also programmed.
The priority queue version of these structures differed from their normal dictionaxy
versions in that the delete operation was customized to support only a delete min.
For skip fists and TMSLs, the level allocation probability p was set to 1/4. While
BSTs are normally defined only for the case when the keys are distinct, they are
easily extended to handle multiple elements with the same key. In our extension,
if a node has key x, then its left subtree has values < x and its right values > x.
To minimize the effects of system call overheads, all structures (other than Heap)

123
were programmed using simulated pointers. The min heap was programmed using a
one-dimensional array.
For our experiments, we began with structures initialized with n = 100, 1,000,
100,000, and 100,000 elements and then performed a random sequence of 100,000
operations. This random sequence consists of approximately 50% insert and 50%
delete min operations. The results axe given in Tables 4.4, 4.5, and 4.6. In the data
sets ‘randoml’ and ‘random2’, the elements to be inserted were randomly generated
while in the data set ‘increasing’ an ascending sequence of elements was inserted
and in the data set ‘decreasing’, a descending sequence of elements was used. Since
BST have very poor performance on the last two data sets, we excluded it from this
part of the experiment. In the case of both randoml and random2, ten random
sequences were used and the average of these ten is reported. The randoml and
random2 sequences differed in that for randoml, the keys were integers in the range
0..(106 — 1) while for random2, they were in the range 0..999. So, random2 is expected
to have many more duplicates.
Table 4.4 gives the total number of comparisons made by each of the methods.
On the two random data tests, weight biased leftist trees required the fewest number
of comparisons except when n = 100,000. In this case, AVL trees required the fewest.
With ascending data, treaps did best and with descending data, LTs and WBLTs did
best. For both, each insert could be done with one comparison as both structures
build a left skewed tree.

Table 4.4: The number of key comparisons
inputs
n
BST
Heap
B-Heap
LT
WBLT
TRP
SKIP
TMSL
AVL
100
621307
823685
268224
165771
165525
407326
270584
503567
373683
random 1
1,000
677570
1317728
383285
203550
202274
537418
397181
729274
542789
10,000
726875
1693955
645664
476534
468713
757946
711869
1104805
685422
100,000
1067670
1746376
1153516
1199207
1171181
1119554
1327083
1778579
877131
100
522728
781808
260209
164288
164031
384721
261169
485995
358885
random2
1,000
612630
1273828
389333
199886
198720
518022
386837
707997
534406
10,000
1027346
1576921
642968
448379
439014
768783
710612
1131421
685926
100,000
5641638
1713676
1146214
1032043
978785
1746732
1339410
1800404
890190
100
-
564332
410119
552081
536796
363045
629085
836223
382513
increasing
1,000
-
946659
655107
917664
882490
496223
1018473
1331317
584422
10,000
-
1284712
814412
1234622
1192325
645765
1313298
1685844
747581
100,000
-
1617645
923048
1550741
1560866
723825
1568819
1939835
902437
100
-
836361
202402
50010
50010
362723
194245
425238
334965
decreasing
1,000
-
1394286
313587
50010
50010
515579
300298
637341
512942
10,000
-
1950062
413840
50010
50010
558222
400082
836879
672910
100,000
-
2400032
534821
50010
50010
648730
450090
936986
835044
n = the number of elements in initial data structures
Total number of operations performed = 100,000

125
The structure height initially and following the 100,000 operations is given in
Table 4.5 for BSTs, Heaps, TRPs and AVL trees. For B-Heaps, the height of the
tallest tree is given. For SKIPs and TMSLs, this table gives the number of levels. In
the case of LT and WBLT, this table gives the length of the rightmost path following
initialization and the average of its length following each of the 100,000 operations.
The two leftist structures axe able to maintain their rightmost paths so as to have a
length much less than log2(n + 1).
The measured run times on a Sun Sparc 5 are given in Table 4.6. For this, the
codes were compiled using the cc compiler in optimized mode. The run time for the
data set randoml is graphed in Figure 4.13. The run time for the data set random2
and Heap, LT, WBLT, SKIP, TMSL, and AVL is graphed in Figure 4.14. For the
data sets randoml and random2 with n = 100 and 1,000, WBLTs required least time.
For randoml with n = 10,000, BSTs took least time while when n = 100,000, both
BSTs and Heaps took least time. For random2 with n = 10,000, WBLTs were fastest
while for n — 100,000, Heap was best. On the ordered data sets, BSTs have a very
high complexity and are the poorest performers (times not shown in Table 4.6). For
increasing data, Heap was best for n = 100, 1,000 and 10,000 and both Heap and
TRP best for n = 100,000. For decreasing data, WBLTs were generally best. On
all data sets, WBLTs always did at least as well (and often better) as LTs. Between
SKIP and TMSL, we see that SKIP generally did better for small n and TMSL for
large n.

Table 4.5: Height/level of the structures
inputs
n
BST
Heap
B-Heap
LT
WBLT
TRP
SKIP
TMSL
AVL
100
13,13
7,6
1,6
4,2
4,2
12,11
4,4
4,4
8,7
randoml
1,000
22,22
10,10
1,10
7,2
7,2
22,24
6,6
6,6
12,12
10,000
31,32
14,14
1,14
8,4
8,4
33,30
8,7
8,7
16,16
100,000
40,41
17,17
1,17
10,9
10,9
42,42
9,9
9,9
20,20
100
13,16
7,7
1,7
4,2
4,2
14,17
4,4
4,4
8,7
random2
1,000
23,69
10,10
1,10
5,2
5,2
23,63
6,5
6,5
12,11
10,000
39,93
14,14
1,14
6,4
6,4
36,83
8,7
8,7
16,15
100,000
147,201
17,17
1,17
6,8
6,8
133,183
9,9
9,9
19,19
100
-
7,8
1,8
6,4
6,4
11,15
4,5
4,5
7,8
increasing
1,000
-
10,11
1,11
9,7
9,7
24,24
6,6
6,6
10,11
10,000
-
14,14
1,14
13,9
13,9
33,34
8,8
8,8
14,14
100,000
-
17,17
1,17
16,11
16,11
46,41
9,9
9,9
17,17
100
-
7,7
1,7
U
1,1
11,15
4,4
4,4
7,7
decreasing
1,000
-
10,10
1,10
1,1
1,1
24,24
6,6
6,6
10,10
10,000
-
14,14
1,14
1,1
1,1
33,33
8,8
8,8
14,14
100,000
-
17,17
1,17
1,1
1,1
46,46
9,9
9,9
17,17
n = the number of elements in initial data structures
Total number of operations performed = 100,000

Table 4.6: Run time using integer keys
inputs
n
BST
Heap
B-Heap
LT
WBLT
TRP
SKIP
TMSL
AVL
100
0.32
0.32
0.53
0.30
0.23
0.56
0.36
0.37
0.50
random 1
1,000
0.34
0.44
0.59
0.29
0.25
0.56
0.35
0.39
0.56
10,000
0.38
0.57
0.98
0.62
0.49
0.71
0.62
0.59
0.70
100,000
0.66
0.66
1.90
1.77
1.26
1.26
1.33
1.32
0.96
100
0.29
0.30
0.55
0.27
0.25
0.55
0.32
0.35
0.50
random2
1,000
0.32
0.44
0.58
0.27
0.23
0.53
0.32
0.36
0.53
10,000
0.49
0.51
0.93
0.57
0.44
0.70
0.59
0.57
0.67
100,000
3.83
0.68
1.92
1.44
0.99
1.70
1.32
1.34
1.02
100
-
0.22
0.63
0.50
0.40
0.42
0.43
0.42
0.62
increasing
1,000
-
0.35
0.92
0.95
0.70
0.48
0.78
0.58
0.58
10,000
-
0.47
1.05
1.25
0.95
0.55
0.72
0.60
0.63
100,000
-
0.60
1.30
1.83
1.40
0.60
0.83
0.65
0.78
100
-
0.30
0.38
0.13
0.12
0.45
0.23
0.28
0.40
decreasing
1,000
-
0.45
0.47
0.13
0.10
0.50
0.28
0.33
0.48
10,000
-
0.58
0.55
0.12
0.12
0.52
0.32
0.38
0.67
100,000
-
0.70
0.73
0.12
0.12
0.55
0.35
0.45
0.80
Time Unit : sec
n = the number of elements in initial data structures
Total number of operations performed = 100,000

128
n
Figure 4.13. Run time on random 1
Another way to interpret the time results is in terms of the ratio m/n (m =
number of operations). In the experiments reported in Table 4.6, m = 100,000. As
m/n increases, WBLTs and LTs perform better relative to the remaining structures.
This is because as m increases, the (weight biased) leftist trees constructed are very
highly skewed to the left and the length of the rightmost path is close to one.
Tables 4.7, 4.8, and 4.9 provide an experimental comparison of BSTs, AVL
trees, MMHs (min-max heaps) [2], Deaps [7], TRPs, SKIPs, and TMSLs as a data
structure for double ended priority queues. The experimental setup is similar to that
used for single ended priority queues. However, this time the operation mix was 50%
insert, 25% delete-min, and 25% delete-max. On the comparison measure, treaps
did best on increasing data (except when n = 100) and skip lists did best when
decreasing data was used. On all other data, AVL trees did best. As far as run time

129
n
Figure 4.14. Run time on random2
is concerned, BSTs did best on the random data tests except when n = 100,000 and
the set random2 was used. In this case, deaps and AVL trees took least time. For
increasing data, treaps were best and for decreasing data, skip lists were best. The
run time for the data set randoml is graphed in Figure 4.15. The run time for the
data set random2 and MMH, Deap, SKIP, TMSL, and AVL is graphed in Figure 4.16.
4.6 Conclusion
We have developed two new data structures: weight biased leftist trees and
modified skip lists. Experiments indicate that WBLTs have better performance (i.e.,
run time characteristic and number of comparisons) than LTs as a data structure for
single ended priority queues and MSLs have a better performance than skip lists as

Table 4.7: The number of key comparisons
inputs
n
BST
MMH
Deap
TRP
SKIP
TMSL
AVL
100
534197
994374
581071
402987
462690
666845
371435
random 1
1,000
676634
1677964
912282
550363
698132
996774
545600
10,000
738100
2328247
1122599
759755
1034669
1437925
693935
100,000
1068250
2795369
1123423
1127616
1439387
1862864
878709
100
514680
941925
557549
396910
447909
650533
357556
random2
1,000
908437
1651830
909747
564414
658284
949230
537616
10,000
1339444
2239519
1145724
881889
1003300
1379690
689415
100,000
5956407
2754760
1125210
1873922
1454138
1875706
891332
100
-
926017
592503
364026
624430
803024
363399
increasing
1,000
-
1660945
1052894
507812
999886
1304223
562873
10,000
-
2392627
1411640
614534
1373332
1766792
726087
100,000
-
3015425
1849318
679120
1487465
1866498
878978
100
-
926041
615944
360030
193021
413284
355999
decreasing
1,000
-
1711076
1062010
490022
287887
603089
563085
10,000
-
2400014
1480876
676190
352780
732673
725079
100,000
-
3035292
1851845
698740
450090
927278
878244
n = the number of elements in initial data structures
Total number of operations performed = 100,000

131
Table 4.8. Height/level of the structures
inputs
n
BST
MMH
Deap
TRP
SKIP
TMSL
AVL
100
13,12
7,6
7,6
13,11
4,4
4,4
8,7
randoml
1,000
22,22
10,10
10,10
23,22
6,6
6,6
12,12
10,000
32,31
14,14
14,14
33,32
8,8
8,8
16,16
100,000
41,41
17,17
17,17
41,42
9,9
9,9
20,20
100
13,13
7,7
7,7
14,12
4,4
4,4
8,7
random2
1,000
23,69
10,10
10,10
22,60
6,6
6,6
12,11
10,000
38,93
14,14
14,14
35,82
8,7
8,7
16,15
100,000
147,199
17,17
17,17
135,186
9,9
9,9
19,19
100
—
7,8
7,8
11,16
4,5
4,5
7,8
increasing
1,000
-
10,11
10,11
24,27
6,7
6,7
10,11
10,000
-
14,14
14,14
33,33
8,8
8,8
14,14
100,000
-
17,17
17,17
46,43
9,9
9,9
17,17
100
-
7,7
7,7
11,15
4,5
4,5
7,8
decreasing
1,000
-
10,10
10,10
24,21
6,7
6,7
10,11
10,000
-
14,14
14,14
33,36
8,8
8,8
14,14
100,000
-
17,17
17,17
46,43
9,9
9,9
17,17
n = the number of elements in initial data structures
Total number of operations performed = 100,000
n
Figure 4.15. Run time on randoml

132
Table 4.9. Run time using integer keys
inputs
n
BST
MMH
Deap
TRP
SKIP
TMSL
AVL
100
0.29
0.42
0.39
0.44
0.45
0.42
0.52
random 1
1,000
0.32
0.62
0.57
0.49
0.56
0.51
0.57
10,000
0.34
0.83
0.81
0.65
0.87
0.74
0.67
100,000
0.64
1.18
1.05
1.17
1.51
1.45
0.99
100
0.27
0.42
0.39
0.47
0.51
0.46
0.56
random2
1,000
0.47
0.64
0.59
0.54
0.53
0.48
0.55
10,000
0.59
0.85
0.78
0.72
0.89
0.71
0.69
100,000
4.22
1.07
1.01
1.91
1.50
1.47
1.01
100
-
0.38
0.38
0.35
0.48
0.38
0.60
increasing
1,000
-
0.60
0.63
0.42
0.65
0.53
0.53
10,000
-
0.88
0.82
0.47
0.80
0.63
0.62
100,000
-
1.12
1.15
0.57
0.92
0.85
0.77
100
-
0.37
0.40
0.35
0.35
0.42
0.53
decreasing
1,000
-
0.63
0.62
0.43
0.33
0.38
0.55
10,000
-
0.83
0.83
0.50
0.35
0.40
0.62
100,000
-
1.05
1.10
0.63
0.40
0.45
0.83
Time Unit : sec
n = the number of elements in initial data structures
Total number of operations performed = 100,000

133
n
Figure 4.16. Run time on random2
a data structure for dictionaries. MSLs have the added advantage of using fixed size
nodes.
Our experiments show that binary search trees (modified to handle equal keys)
perform best of the tested double ended priority queue structures using random data.
Of course, these are unsuitable for general application as they have very poor per¬
formance on ordered data. Min-max heaps, deaps and AVL trees guarantee O(logn)
behavior per operation. Of these three, AVL trees generally do best for large n. It
is possible that other balanced search structures such as bottom-up red-black trees
might do even better. Treaps and skip lists are randomized structures with O(logn)
expected complexity. Treaps were generally faster than skip lists (except for decreas¬
ing data) as double ended priority queues.

134
For single ended priority queues, if we exclude BSTs because of their very poor
performance on ordered data, WBLTs did best on the data sets randoml and random2
(except when n = 100,000), and decreasing. Heaps did best on the remaining data
sets. The probabilistic structures TRP, SKIP and TMSL were generally slower than
WBLTs. When the ratio m/n (m = number of operations, n = average queue size)
is large, WBLTs (and LTs) outperform heaps (and all other tested structures) as the
binary trees constructed tend to be highly skewed to the left and the length of the
rightmost path is close to one.
Our experimental results for single ended priority queues are in marked con¬
trast to those reported in Gonnet and Baeza-Yates [11, p228] where leftist trees are
reported to take approximately four times as much time as heaps. We suspect this
difference in results is because of different programming techniques (recursion vs. it¬
eration, dynamic vs. static memory allocation, etc.) used in Gonnet and Baeza-Yates
[11] for the different structures. In our experiments, all structures were coded using
similar programming techniques.

CHAPTER 5
CONCLUSIONS
We have considered the problem of joining a row of compacted cells and de¬
veloped heuristics to stretch cells and river-route the nets so that the layout area is
minimized. Our proposed heuristics were compared, experimentally, with Fang [8]
and found to produce layouts with less area.
We have developed a new weight balanced search structure called /3-BBST. Our
experimental results show that BBSTs generally have the best search cost of the
structures considered. We recommend /3-BBSTs for environments where searches
are done with much greater frequency than inserts and/or deletes. Based on our
experiments, we conclude that AVL trees remain the best dictionary structure for
general applications.
We have also proposed two simplified versions of the BBST called SBBST and
BBSTD. Both versions are very competitive with BBSTs. The BBST and SBBST
provide slightly better balance than provided by WB(a) trees.
Another new data structures, weight biased leftist trees, developed by us, has
better performance than leftist trees as a data structure for single ended priority
queues. We have also proposed a fixed node size representation for skip lists. As
a data structure for dictionaries, this has a better performance than skip lists with
variable node size.
135

APPENDIX A
ABBREVIATIONS
136

Table A.l. Abbreviations used in tables
AVL
Adelson-Velskii and Landis’ trees
B-Heap
binomial heaps
BBST
/^-balanced binary search trees
BBSTD
BBSTs without deletion
BST
unbalanced binary search trees
Deap
deaps
DSL
deterministic skip lists
Heap
heaps
LT
leftist trees
MMH
min-max heaps
MSL
modified skip lists
RB-B
bottom-up red-black trees
RB-T
top-down red-black trees
SBBST
simple /2-balanced binary search trees
SKIP
skip lists
TMSL
threaded modified skip lists
TRP
treaps
WB
weight balanced trees
WBLT
weight biased leftist trees

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BIOGRAPHICAL SKETCH
Seonghun Cho was brought up in the Republic of Korea. He completed his
undergraduate degree in mathematics from Seoul National University. He obtained
an M.S degree in computer science from University of Missouri-Columbia, Columbia,
MO. He started to work towards his Ph.D from Fall 1991 in the Computer and
Information Science and Engineering Department at the University of Florida. His
research interests include VLSI CAD algorithms and data structures. He will be
graduating in May 1996.
141

I certify that I have read this study and that in my opinion it conforms
to acceptable standards of scholarly presentation and is fully adequate, in
scope and quality, as a dissertation for the degree of Doctor of Philosophy.
Professor of Computer and
Information Science and Engineering
I certify that I have read this study and that in my opinion it conforms
to acceptable standards of scholarly presentation and is fully adequate, in
scope and quality, as a dissertation for th Li-Min Fu
Associate Professor of Computer and
Information Science and Engineering
I certify that I have read this study and that in my opinion it conforms
to acceptable standards of scholarly presentation and is fully adequate, in
scope and quality, as a dissertation for the degree of Doctor of Philosophy.
r" â– '
Theodore Johnson
Associate Professor of Computer and
Information Science and Engineering
I certify that I have read this study and that in my opinion it conforms
to acceptable standards of scholarly presentation and is fully adequate, in
scope and quality, as a dissertation for the degree of Doctor of Philosophy.
Sangilthevar Rajasekaran
Associate Professor of Computer and
Information Science and Engineering

I certify that I have read this study and that in my opinion it conforms
to acceptable standards of scholarly presentation and is fully adequate, in
scope and quality, as a dissertation for the degree of Doctor of Philosophy.
Pyi W.Omn
P/ofessor of Biochemistry and
Molecular Biology
This dissertation was submitted to the Graduate Faculty of the College
of Engineering and to the Graduate School and was accepted as partial ful¬
fillment of the requirements for the degree of Doctor of Philosophy.
May 1996
/ £>•.
Winfred M. Phillips
Dean, College of Engineering
Karen A. Holbrook
Dean, Graduate School