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Query processing in time-dependent spatial networks

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Query processing in time-dependent spatial networks

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Content QUERY PROCESSING IN TIME-DEPENDENT SPATIAL NETWORKS by Ugur Demiryurek A Dissertation Presented to the FACULTY OF THE GRADUATE SCHOOL UNIVERSITY OF SOUTHERN CALIFORNIA In Partial Fulfillment of the Requirements for the Degree DOCTOR OF PHILOSOPHY (COMPUTER SCIENCE) December 2012 Copyright 2012 Ugur Demiryurek Dedication This thesis is dedicated to my parents and to my wife ii Acknowledgments I would like to thank my advisor, Professor Cyrus Shahabi for his endless support throughout my PhD studies at the University of Southern California. He has greatly motivated my interests in this topic and guided me thoughtfully in my pursuit of knowl- edge and ideas. I have been always inspired by his vision and his dedication to work. I would also like to thank my proposal and dissertation committee members Profes- sor Gaurav S. Sukhatme, Professor Ulrich Neumann, Professor Aiichiro Nakano, and Professor Genevieve Giuliano for all their valuable advice and guidance from various perspectives. I thank Professor David Kempe for all the stimulating and invaluable con- versations. My deep appreciation to Professor Micheal Arbib who has thought me the fundamentals of research and supported me in my first years at USC. I am also thankful to all my colleagues at Infolab for their friendship and help. I very much enjoyed the time in the past several years with them: Songhua Xing, Kiyoung Yang, Hyunjin Yoon, Colin Gu, Jeff Khoshgozaran-Haghighi, Afsin Akdogan, Mehrdad Jahangiri, Mehdi Sharifzadeh, Leyla Kazemi, Houtan Shirani-Mehr, Ali Khodaei, Bei Pan, Ling Hu, Lian Liu, and Shireesh Asthana. I wish them all personal and professional successes in life. As always my family have been a constant source of encouragement and support dur- ing my studies at USC. My love and gratitude to my parents Hacer and Ihsan Demiryurek who have been my inspiration and motivation for continuing my dreams, my sisters iii Emine and Emel Demiryurek, and my brother Ufuk Demiryurek. I thank my dear wife Onur Demiryurek for all of the love and patience she has has given through all of these years, and for standing beside me throughout the though times. I also thank my won- derful children: Emre and Arda for always making me laugh and for understanding on those weekends when I was working on this thesis instead of playing with them. I hope that one day they can look into this thesis and smile why I spent so much time in front of my computer. This landmark in my life’s journey would not have been possible without love, inspi- ration, and sacrifices of my family. To them, I dedicate this thesis. iv Table of Contents Dedication ii Acknowledgments iii List of Tables vii List of Figures viii Abstract x Chapter 1 : Introduction 1 1.1 Motivation . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 1 1.2 Challenges . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 4 1.3 Thesis Statement . . . . . . . . . . . . . . . . . . . . . . . . . . . 7 1.4 Overview of Proposed Approaches . . . . . . . . . . . . . . . . . . 7 1.5 Dissertation Outline . . . . . . . . . . . . . . . . . . . . . . . . . . 11 Chapter 2 : Related Work 13 2.1 Nearest Neighbor Search . . . . . . . . . . . . . . . . . . . . . . . 13 2.1.1 kNN Search in Euclidean Space . . . . . . . . . . . . . . . 13 2.1.2 kNN Search in Spatial Networks . . . . . . . . . . . . . . . 15 2.2 Time-dependent Shortest Path . . . . . . . . . . . . . . . . . . . . 18 Chapter 3 : Problem Definition 22 3.1 Time-dependent Spatial Network . . . . . . . . . . . . . . . . . . . 22 3.2 Time-dependent Query Processing . . . . . . . . . . . . . . . . . . 24 Chapter 4 : Baseline kNN Search Solutions in Time-dependent Road Networks 26 4.1 Time-dependent kNN Search using Time-Expanded Networks . . . 27 4.2 Time-dependent kNN Search using Network Expansion . . . . . . . 29 4.3 Performance Evaluation . . . . . . . . . . . . . . . . . . . . . . . . 32 4.4 Summary of Results . . . . . . . . . . . . . . . . . . . . . . . . . . 36 v Chapter 5 : Efficient kNN Query Processing in Time-Dependent Spa- tial Networks 38 5.1 Index Construction . . . . . . . . . . . . . . . . . . . . . . . . . . 39 5.1.1 Tight Index Construction . . . . . . . . . . . . . . . . . . . 40 5.1.2 Loose Index Construction . . . . . . . . . . . . . . . . . . . 42 5.1.3 Tight and Loose R-Tree . . . . . . . . . . . . . . . . . . . . 44 5.2 Time-dependent kNN Query Processing . . . . . . . . . . . . . . . 47 5.2.1 Nearest Neighbor Query . . . . . . . . . . . . . . . . . . . 48 5.2.2 kNN Query . . . . . . . . . . . . . . . . . . . . . . . . . . 49 5.3 Time-dependent Shortest Path Computation . . . . . . . . . . . . . 52 5.4 Performance Evaluation . . . . . . . . . . . . . . . . . . . . . . . . 55 Chapter 6 : Online Fastest Path Computation in Time-dependent Spa- tial Networks 61 6.1 Feasibility of Time-dependent Path Planning . . . . . . . . . . . . . 65 6.2 Time-dependent Fastest Path Computation . . . . . . . . . . . . . . 68 6.2.1 Precomputation Phase . . . . . . . . . . . . . . . . . . . . 70 6.2.2 Online B-TDFP Computation . . . . . . . . . . . . . . . . . 74 6.3 Performance Evaluation . . . . . . . . . . . . . . . . . . . . . . . . 80 Chapter 7 : Conclusion and Future Work 87 7.1 Conclusion . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 87 7.2 Discussion of Future Work . . . . . . . . . . . . . . . . . . . . . . 88 Bibliography 92 Appendix A: Indexing Network Voronoi Diagrams 98 A.1 Preliminaries . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 102 A.2 Indexing Network V oronoi Cells . . . . . . . . . . . . . . . . . . . 104 A.2.1 Network V oronoi Diagram Construction . . . . . . . . . . . 105 A.2.2 Index Generation on Network V oronoi Diagram . . . . . . . 106 A.3 Performance Evaluation . . . . . . . . . . . . . . . . . . . . . . . . 113 vi List of Tables 4.1 Experimental parameters . . . . . . . . . . . . . . . . . . . . . . . 33 5.1 Experimental parameters for TD-kNN . . . . . . . . . . . . . . . . 56 6.1 Experimental Results . . . . . . . . . . . . . . . . . . . . . . . . . 84 6.2 Lower-bound Quality . . . . . . . . . . . . . . . . . . . . . . . . . 86 A.1 Experimental parameters . . . . . . . . . . . . . . . . . . . . . . . 114 vii List of Figures 1.1 Real-world travel time for a weekday on a segment of I-10 in Los Angeles County . . . . . . . . . . . . . . . . . . . . . . . . . . . . 1 1.2 Time-dependent graph . . . . . . . . . . . . . . . . . . . . . . . . 3 1.3 Arrival-time cost functions . . . . . . . . . . . . . . . . . . . . . . 5 1.4 Blind network expansion . . . . . . . . . . . . . . . . . . . . . . . 6 4.1 Time-dependent 1-NN search . . . . . . . . . . . . . . . . . . . . . 27 4.2 A Time-dependent GraphG T (V;E) . . . . . . . . . . . . . . . . . 28 4.3 A Time-expanded graph . . . . . . . . . . . . . . . . . . . . . . . . 28 4.4 Correctness and impact of k . . . . . . . . . . . . . . . . . . . . . . 34 4.5 Response time versus distribution and network size . . . . . . . . . 35 4.6 Response time versus query/object cardinality and agility . . . . . . 36 5.1 Tight cell construction ofP 1 . . . . . . . . . . . . . . . . . . . . . 41 5.2 Tight Cells . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 41 5.3 Loose Cells . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 43 5.4 LN R-Tree . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 45 5.5 Second NN example . . . . . . . . . . . . . . . . . . . . . . . . . . 50 5.6 TDSP localization . . . . . . . . . . . . . . . . . . . . . . . . . . . 53 5.7 Coverage ratio . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 57 5.8 Average number of neighbors . . . . . . . . . . . . . . . . . . . . . 57 5.9 Impact of k . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 58 5.10 Network node access . . . . . . . . . . . . . . . . . . . . . . . . . 58 5.11 Object cardinality . . . . . . . . . . . . . . . . . . . . . . . . . . . 59 5.12 Query cardinality . . . . . . . . . . . . . . . . . . . . . . . . . . . 59 5.13 Object and Query distribution . . . . . . . . . . . . . . . . . . . . . 60 6.1 Static vs Time-dependent path planning . . . . . . . . . . . . . . . 63 6.2 TDFP and FP Comparison . . . . . . . . . . . . . . . . . . . . . . 67 6.3 Road network partitioning . . . . . . . . . . . . . . . . . . . . . . 72 6.4 Lower-bound distance computation. . . . . . . . . . . . . . . . . . 73 6.5 Bidirectional search . . . . . . . . . . . . . . . . . . . . . . . . . . 77 6.6 TD-ALT Comparison . . . . . . . . . . . . . . . . . . . . . . . . . 82 6.7 Speed-up Ratio Analysis . . . . . . . . . . . . . . . . . . . . . . . 86 A.1 Network V oronoi Diagram . . . . . . . . . . . . . . . . . . . . . . 100 viii A.2 V oronoi diagram in Euclidean space . . . . . . . . . . . . . . . . . 102 A.3 A Road network and network V oronoi diagram . . . . . . . . . . . 105 A.4 Network V oronoi diagram with P =fp 1 ;:::;p 50 g in Los Angeles road network. . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 106 A.5 Network V oronoi cell construction in VR-tree . . . . . . . . . . . . 107 A.6 False-negative edges of a NVC in Los Angeles road network . . . . 109 A.7 Minimum bounding rectangles on network V oronoi cells . . . . . . 110 A.8 VQ-tree on Los Angeles road network . . . . . . . . . . . . . . . . 111 A.9 VQ-tree . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 112 A.10 Impact of object cardinality and distribution . . . . . . . . . . . . . 115 A.11 Impact of network size . . . . . . . . . . . . . . . . . . . . . . . . 116 A.12 Response time vs object cardinality and Index reconstruction . . . . 117 ix Abstract Recent advances in online map services and their wide deployment in hand-held devices and car-navigation systems have led to extensive use of location-based services. The most popular class of such services are route planning and k-nearest neighbor(kNN) queries where users search for geographical points of interests (e.g., restaurants, gas stations) with corresponding travel-times to these locations. Accordingly, many recent studies focused on developing efficient techniques to answer point-to-point fastest path and k-nearest neighbor search queries in the spatial network space. However, most of the existing approaches in spatial networks make the simplify- ing assumption that the cost of traveling each edge of the spatial network is constant (e.g., corresponding to the length of the edge). Whereas in real world, the actual travel cost of a network edge is time-dependent i.e., the cost of a network edge depends on the arrival-time to that edge. Unfortunately, once we consider the time-dependent edge weights in road networks, all proposed kNN and shortest path query solutions assum- ing constant edge weights fail. With time-dependent edge costs, the network distance between two nodes is not unique and varies based on the departure time from the source. This dynamism of the distance introduces great challenges in developing precomputa- tion techniques to expedite spatial query processing in time-dependent spatial networks. In this thesis, for the first time we study the problem of k-nearest neighbor search in time-dependent road networks where the weight of each edge is a function of time. We x propose series of index structures that efficiently and accurately answer the k nearest neighbor queries in time-dependent road networks and effectively handle the database updates where points of interests are added or removed. In addition, we study the prob- lem of point-to-point shortest path computation in time-dependent spatial networks and present a technique which speeds-up the path computation using a bidirectional time- dependent A* search based on a novel heuristic function. In this thesis, the efficacy of all proposed techniques for both kNN search and point-to-point shortest path computation have been verified with extensive experiments using real data-sets including a variety of large spatial networks with real traffic data. xi Chapter 1 Introduction 1.1 Motivation With the ever growing popularity of online map services (e.g., Google-Maps) and their wide deployment in hand-held devices (e.g., iPhone) and car-navigation systems, more and more users search for geographical points of interests (e.g., restaurants) and the corresponding directions to these locations. Consequently, many recent research studies (e.g., [3, 28, 29, 36, 39, 49, 57]) focused on developing techniques to accurately and efficiently compute the distance and route between the objects in large road networks. The majority of existing studies make the simplifying assumption that the cost of traveling each edge of the road network is constant and rely on pre-computation of distances in the network. However, the actual travel-time on road networks heavily depends on the traffic congestion on the edges and hence is a function of the time of Figure 1.1: Real-world travel time for a weekday on a segment of I-10 in Los Angeles County 1 the day. For example, Figure 1.1 shows the measured variation of real-world travel- time for a particular segment of I-10 freeway in Los Angeles between 6AM and 8PM on a weekday. Two main observations can be made from this figure. First, the arrival- time to the segment entry determines the travel-time on that segment i.e., travel-time is time-dependent. Second, the change in travel-time is significant and continuous (not abrupt), for example from 8:30AM to 9:00AM, the travel-time of this segment changes from 30 minutes to 18 minutes (40% decrease). These observations have major com- putation implications: the actual optimal path from a source to a destination may vary significantly depending on the departure-time from the source, and hence, the result of spatial queries (e.g., such KNN or point-to-point fastest path) on such time-dependent network heavily depend on the time at which the query is issued. Unfortunately, once we consider the time-dependent road networks, all the techniques assuming constant edge-weights would fail to address the spatial-queries. To illustrate, we show a simple example in Figure 1.2 where a spatial network is modeled as a graph and edge travel- times are time-dependent. Consider the snapshot of the network (i.e., a static network) with edge weights correspond to travel-time values att=0. With classic fastest path com- putation approaches that disregard travel-time time-dependency, the fastest path froms tod goes throughv 1 ;v 2 ;v 4 with a cost of 13 time units. However, by the time whenv 2 is reached (i.e., att=5), the cost of edgee(v 2 ;v 4 ) changes from 8 to 12 time units, and hence reachingd throughv 2 takes 17 time units instead of 13 (as it was anticipated at t=0). In contrast, if the time-dependency of edge travel-times are considered and hence path going throughv 1 ;v 3 ;v 4 was taken, the total travel-cost would have been 15 units which is the actual optimal fastest path. Although time-dependent fastest path computation is the most accurate and realistic path computation method in road networks, we observe (at the time of this thesis is being written) that the existing state-of-the art online map path planning applications (e.g., 2 Figure 1.2: Time-dependent graph Google Maps, Bing Maps) and car navigation devices do not employ time-dependency in their path computations, and hence, their fastest path recommendation remains the same throughout the day regardless of the departure-time from the source (i.e., query time). While some of these applications provide alternative paths under traffic condi- tions (which may seem similar to time-dependent planning at first), we note that the recommended alternative paths and their corresponding travel-times still remain unique during the day, and hence no time-dependent planning. To the best of our knowledge, these applications computetop-k fastest paths (i.e.,k alternative paths) and their corre- sponding travel-times with and without taking into account the traffic conditions. The travel-times which take into account the traffic conditions are simply computed by con- sidering increased edge weights (that corresponds to traffic congestion) for each path. Meanwhile, an increasing number of navigation companies have started releasing their time-dependent travel-time information for road networks. For example, Navteq [42], the leading provider of navigation services, offer traffic flow services that provide time-dependent travel-time (at the temporal granularity of as low as five minutes) of road network edges up to one year. The time-dependent travel-times are usually extracted 3 from the historical traffic data and local information like weather, school schedules, and events. Considering the availability of time-dependent travel-time information for road net- works on the one hand and the importance of time-dependency for accurate and realistic route planning on the other hand, this thesis extends existing literature on spatial query processing and planning in road networks to a new family of time-dependent query pro- cessing solutions. 1.2 Challenges The efficient computation of point-to-point fastest path and more complex k nearest neighbor search on time-dependent spatial networks is very challenging due to following reasons: 1) Input Size for Precomputation: Given the nature of online map applications, the fastest path and kNN queries require very fast response time. Towards that end one can consider several precomputation techniques in time-dependent road networks similar to those proposed for static road networks (e.g., [28, 36, 57, 68]). However, precom- putation in time-dependent road networks is very challenging due to huge input size (i.e., the number of shortest paths). Specifically, unlike static road networks where the shortest path between any pair of nodes is unique, the shortest path in time-dependent networks depends on the departure time from the source, and hence is not unique. For instance, consider our previous example in Figure 1.2 where a spatial network is mod- eled as a time-dependent graph and edge-weights are piece-wise linear function of time. There are three possible paths from v 1 to v 4 ; p 1 =fv 1 ;v 2 ;v 4 g, p 2 =fv 1 ;v 2 ;v 3 ;v 4 g, and p 3 = fv 1 ;v 3 ;v 4 g. Since edge weights in the network are function of time, the arrival-time (and the total travel-cost) to destination is also function of time and changes 4 based on the departure-time fromv 1 , i.e.,f p 1 =f 24 (f 12 (t));f p 2 =f 34 (f 23 (f 12 (t))), and f p 3 = f 34 (f 13 (t)), wheref p i represent the total travel-cost function of pathp i . Figure 1.3 plots the total travel-time function to destination for different departure-times from v 1 . The blue-line in this figure represents the lower-envelope of the cost functions for all paths. Each piece of the lower-envelope points to the shortest path for the corresponding time-interval in time domain. For instance, one should takep 3 if (s)he leavesv 1 between t = 3 andt = 7. On the other hand,p 1 would be the optimal path if the departure-time is betweent=7 andt=10. It has been conjectured that a lower-envelope between any nodes in time-dependent spatial networks contains super-polynomial number of linear pieces (i.e., path selections) [5]. Clearly, an algorithm which precomputes the every possible path for any pair of nodes in large time-dependent networks and stores the cor- responding path selections would suffer from exponential time and storage complexity. Figure 1.3: Arrival-time cost functions 2) Object Distribution: In practice, it is possible to answer kNN queries in time- dependent road networks by implementing an incremental network expansion approach with which starting from q all network nodes reachable from q in every direction are visited in order of their shortest time proximity toq until allk nearest data objects are located. However, the overhead of executing network expansion is prohibitively high 5 particularly in large networks with a sparse (but perhaps large) set of data objects. This is because such a blind search approach has to redundantly visit many network nodes. For example, Figure A.3 depicts a real spatial network (i.e., San Joaquin, CA) and illustrates the set of nodes that network expansion would have to visit (marked by the shaded area) to locate the first nearest data object (1-NN) for the query objectq. In this case,47:2% of the entire set of network nodes must be visited to find the first nearest neighbor. Figure 1.4: Blind network expansion 3) Lack of Efficient Index Structures: Classic spatial index structures (e.g., Quad- tree [18], R-tree [25]) have been used to expedite query processing in Euclidean spaces. These indexing schemes have also been used in static road networks (e.g., [57, 68]) to expedite the process of computing shortest paths from that of spatial query processing by decoupling the domain of the participating objects from the domain of the vertices of the spatial network. These index structures are created assuming there exists only one (unique) shortest path between any pair of the nodes in the network and can be used as long as the spatial network is unchanged. It is not feasible, if not impossible, to simply extend these index structures to time-dependent road networks where the shortest path between any pair of nodes is dynamic, i.e., changes based on the departure time from the 6 source. To the best of our knowledge, there is no index structure for efficient processing of spatial queries in time-dependent spatial networks. 1.3 Thesis Statement Due to recent sensor instrumentation of road networks in major cities as well as the advances in crowd-sourcing techniques that collect large amount of traffic data from GPS enabled devices such as in-car navigation systems and smart phones, it is now becoming possible to forecast and model time-dependent traffic flows in road networks. Therefore, traffic-sensitive time-dependent spatial queries (e.g., fastest path, k nearest neighbor search, and etc.) on road networks will be a common practice in the near future. Existing spatial query processing techniques in road networks make the sim- plifying assumption that travel-times of the network edges are constant. However, in real-world road networks, the edge travel-times are time-dependent, where the arrival- time to an edge determines the actual travel-time. This dynamism of the distance posses great challenges in developing efficient algorithms to evaluate spatial queries in time- dependent road networks. Our goal in this thesis is to develop index structures for efficient and exact processing of k nearest neighbor and fastest path queries in time- dependent road networks. 1.4 Overview of Proposed Approaches In this thesis, we propose solutions for two fundamental spatial queries on time- dependent road networks. First, we study the problem of time-dependent k Nearest Neighbor(TD-kNN) query which finds the k points of interest that are in the shortest network distance (in travel-time) to a given query point. Second, we propose a bidi- rectional time-dependent A* algorithm based on a novel heuristic function to efficiently 7 answer point-to-point fastest path queries. The objectives of our proposed algorithms for both kNN and fastest path queries are as follows: 1) find exact (not approximate) solutions in time-dependent road networks, 2) efficiently answer the queries to support location based service applications that require immediate response time, 3) be scalable in order to be applicable to large-scale road networks, 4) enable fast pre-computation time and low space overhead, and 5) efficiently cope with database updates where edge weights and data objects are updated. kNN Query Processing in Time-dependent Spatial Networks The kNN query searches for the k closest points of interest (e.g., gas station, restau- rant) with minimum distance to a query point q, where the distance is defined by a domain-specific metric such as Minkowski distance (e.g., Euclidean distance) or net- work distance. Unlike the existing studies on kNN queries that assume the cost of trav- eling each edge of the spatial network is constant (e.g., corresponding to the length of the edge), we consider the cost of a network edge is time-dependent (e.g., travel-time). As a result the network distance between any nodes in the spatial network becomes time-dependent (not unique), hence the term time-dependent spatial network. We intro- duce two baseline solutions [11] to kNN query processing in time-dependent spatial networks. Firstly, we use an incremental network expansion algorithm with which all network nodes reachable fromq are visited in order of their time-dependent travel-time proximity until all k nearest objects are located. While this is a practical approach, it cannot scale to large networks with sparse set of data objects due to high overhead of executing network expansion. Secondly, we exploit time-expanded graphs [33] to model the time-dependent networks. With this approach, we discretize the time domain 8 and at each discrete time instant we use a snapshot of the network to represent the time- dependent network. Although this approach allows exploiting the existing precompu- tation and index structures developed for static networks, it fails to provide the correct results because the time-expanded model misses the state of the network between any two discrete time instants. We address the disadvantages of both baseline approaches by developing a novel technique [10] that efficiently and accurately finds k nearest neighbors of a query object in time-dependent road networks. The main idea behind our approach is to employ a filter and refinement strategy by utilizing a distance ranking method, in which we use the time-independent upper-bound and lower-bound distances between data objects to filter-out the obviously wrong answers and filter-in only a small number of objects that are potential candidates when the number of required objects is larger than 1. In particular, we partition the spatial network into neighborhoods around the data objects by creating two subnetworks for each data object called Tight Cell (TC) and Loose Cell (LC). The tight cellTC(p i ) is a subnetwork around data objectp i in which any query object is guaranteed to have p i as its nearest neighbor in a time-dependent network. On the other hand, the loose cellLC(pi) is a subnetwork aroundp i outside which any point is guaranteed not to have p i as its nearest neighbor. In other words, data object p i is guaranteed not to be the nearest neighbor of q if q is outside of the loose cell of p i . The complementary index structures built on the tight and loose cells enable us to localize the search space in time-dependent road networks by immediately finding the first nearest neighbor and then expand the search area (in a much smaller network) to find the remaining k-1 neighbors. 9 Online Computation of Fastest Path in Time-dependent Spatial Networks Given a sources and destinationd, and a departure-timet s from the source, the fastest path query in time-dependent spatial networks finds the path with the minimum travel- time among all paths froms tod with the departure-timet s . The time-dependent fastest path problem was shown first by Dreyfus [16] to be polynomially solvable by a triv- ial modification to a label-setting (e.g., Dijkstra) algorithm where, analogous to shortest path distances, the arrival-time to the nodes is used as the labels that form the basis of the greedy algorithm. However, Dreyfus’s algorithm is far too slow for online map appli- cations which are usually deployed on very large networks and require almost instant response times. There are many efficient precomputation approaches that answer fastest path queries in near real-time (e.g., [58], [67]) in static road networks. However, it is infeasible to extend these approaches to time-dependent networks. This is because the input size (i.e., the number of fastest paths) increases drastically in time-dependent networks. Specif- ically, since the length of as-d path changes depending on the departure-time froms, the fastest path is not unique for any pair of nodes in time-dependent networks. It has been shown in [5, 20] that the number of fastest paths between any pair of nodes in time-dependent road networks can be super-polynomial. Hence, an algorithm which considers the every possible path (corresponding to every possible departure-time from the source) for any pair of nodes in large time-dependent networks would suffer from exponential time and prohibitively large storage requirements. Given these challenges, we propose a bidirectional time-dependent fastest path algo- rithm [12] based on A* search [27]. There are two main challenges to employ bidirec- tional A* search in time-dependent networks. First, finding an admissible heuristic func- tion (i.e., lower-bound distance) between an intermediatev i node and the destinationd is challenging as the distance between v i and d changes based on the departure-time 10 from v i . Second, it is not possible to implement a backward search without knowing the arrival-time at the destination. We address the former challenge by partitioning the road network to non-overlapping partitions (an off-line operation) and precompute the intra (node-to-border) and inter (border-to-border) partition distance labels with respect to Lower-bound GraphG which is generated by substituting the edge travel-times inG with minimum possible travel-times. We use the combination of intra and inter distance labels ( found by simple table lookups) as a heuristic function in the online computation. To address the latter challenge, we run the backward search on the lower-bound graph (G) which enables us to identify the set of the nodes that needs to be explored by the forward search. 1.5 Dissertation Outline The remainder of this thesis is organized as follows. In Chapter 2, we review the related work on kNN queries on both Euclidean and spatial network spaces as wells as the time-dependent fastest path algorithms. In Chapter 3, we formally define the problem ofkNN and fastest path queries in time-dependent spatial networks and introduce the terminology we use throughout this thesis. In Chapter 4, we introduce two baseline approaches for time-dependentkNN queries based on a) time-expanded networks, and b) network expansion framework that exploits time-dependent label setting algorithm [16]. We analyze the performance of these solutions on real datasets and discuss their inefficiencies and drawbacks. In Chapter 5, we addresses the disadvantages of both baseline approaches by introducing a novel algorithm that efficiently and accurately answerskNN queries in time-dependent road networks. We evaluate the performance of our proposed algorithms with a variety of real-world time-dependent spatial networks with large number of data and query objects. In Chapter 6, we propose an algorithm to 11 answer fastest path computation in time-dependent spatial networks based on a bidi- rectional A* search with a novel heuristic function. Finally, we conclude and propose extensions to our research. 12 Chapter 2 Related Work In this section we review the previous studies onkNN query processing in both Euclid- ian space and road networks as well as the time-dependent shortest path computation. 2.1 Nearest Neighbor Search ThekNN query searches for the k closest points of interest (e.g., gas station, restaurant) with minimum distance to query points, where the distance is defined by a domain- specific metric such as Minkowski distance (e.g., Euclidean distance) or network dis- tance. The majority of the existing work on kNN queries in the past literature assumes Euclidean space (with Euclidean distance) as the native metric space. These techniques focus on developing spatial access methods (e.g., R-tree, Quad-tree) for efficient pro- cessing of kNN queries. However, the access methods that assume Euclidean space are not readily applicable with spatial networks, where the distance (i.e., network dis- tance) between the objects depend on the connectivity of the network. Accordingly, many recent studies focus on developing various pre-computation based techniques to speed-up distance computation and, hence, efficiently answer kNN queries in the spatial network space. Next we review both categories in turn. 2.1.1 kNN Search in Euclidean Space Given a set of n data objects P =fp 1 ;p 2 ;:::;p n g in Euclidean space, a query point q and a distance function d, the kNN query with respect to q finds a subset P 0 P 13 ofk objects with minimum distance toq, i.e., for any objectp 0 2 P 0 andp2 PP 0 , d(q;p 0 )d(q;p). In this case, the distance function is the Minkowski metric (Euclidean metric) and the data object are usually indexed by a spatial index (e.g., R-Tree [25]). The first nearest neighbor algorithm proposed by Roussopoulos et al. [53] exploits the R-Tree in a depth-first manner, recursively visiting the node with the minimum dis- tance fromq. Zheng et al. [76] proposed akNN algorithm where the data objects are static and the query objects are mobile. In their work, they pre-compute and store in an R-Tree the V oronoi [46] diagram of data objects. When a NN query submitted, the server computes the nearest neighbor efficiently using the V oronoi diagram as well as the validity time T. While the result remain the same during the T, the mobile query object requests server to re-evaluate the NN after the T. This method is only efficient for 1st NN queries since computingkNN would require constructing order-k V oronoi diagrams that is very complex and requires extensive space. Furthermore, estimating the validity time assuming the maximum speed is not very realistic in real-word applications. Zhang et al. [75] improved this method by finding the validity region that the NN set remains the same. In [62], Song and Roussopoulos introduced an algorithm that computes and caches m-NNs (m> k) for a moving query object so that if the client moves out of the kNN range, the newkNNs can be computed among the cached m objects. The major drawback of this approach is to define the right value of the m. In [65], Tao et al. pre- sented an algorithm that pre-calculates thekNN of a query object along the trajectory by utilizing split points. In their work, they divide the trajectory into sub-segments by assuming the query object moves in a steady speed and computed thekNN for each sub- segment bounded by split points. Later, Tao et al. proposed time-parameterized (TP) queries [63]. With TP when the central server receives akNN query it computes the current result set (R), validity time (T) and set of data objects that will cause a change (C) after the T and returns the client (R,T,C). Both [65] and [63] assumes that the 14 query objects move in a steady speed and their movement is constant during the query. Although tree based (e.g., R-tree) data structures are efficient to handle stationary spatial data, they suffer from the node reconstruction caused by location updates with mobile objects. Therefore, some researchers have focused on the simple yet efficient grid-based structures to index and query the moving objects. Kalashnikov et al. and Yu et al. intro- duced grid-based in-memory algorithms (referred to as Q-Index [31] and YPK-CNN [74] respectively) for efficient and scalable processing of continuous range and kNN queries over moving objects. With both studies query indexing technique is used to avoid constant updates to index structures. The shared execution method has been used in SINA [37] for continuous spatio-temporal range queries, and in SEA-CNN [70] for continuous spatio-temporal kNN queries. The goal of shared execution algorithm is to abstract numerous spatial queries as spatial joins between the set of moving objects and queries. Finally, Mouratidis et al. [38] proposed an approach termed as CPM that defines a conceptual partitioning of the space by arranging grid cells into rectangles. All of the above approaches are based on either spatial (e.g., R-Tree) or grid index structures that are applicable to the spaces where the distance between objects is only a function of their spatial attributes (e.g., Euclidean distance). In real-world scenarios, however, the objects move on spatial networks. We review the existing work on kNN query processing in spatial networks next. 2.1.2 kNN Search in Spatial Networks UnlikekNN query in Euclidean space where the distance can be obtained immediately in constant time, the distance function in road networks is the length of the shortest path between q and the data objects. In this case, a road network is usually modeled as a connected planar graph. Since the distance function in road networks is the shortest 15 path, the majority ofkNN query techniques in road networks are based on incremental network expansion [2, 39, 49] which relies on the Dijkstra Algorithm [14]. In [49], Papadias et al. studied two different techniques, namely Incremental Net- work Expansion INE and Euclidean Restriction ER method to answer kNN queries in road networks. With INE, given a query pointq and connected planar graphG (i.e., static road network), starting fromq all network nodes reachable fromq in every direc- tion are visited in order of their proximity to q (hence, a one-to-many search) until allk nearest data objects are located. Clearly, the overhead of executing network expansion can be very high in large networks with a sparse set of data objects, because such a blind search approach has to redundantly visit many network nodes which are away from the shortest paths to the nearest data objects. On the other hand,ER exploits the fact that Euclidean distance is a lower-bound for the network distance to answerkNN queries in static road networks. Specifically,ER uses a filtering mechanism to rapidly identify a set of candidate data objects based on their Euclidean distance fromq , which is then refined by computing their actual shortest path fromq to identify the exact set of nearest neighbors. SinceER relies on lower-bound restriction, it yields better results when net- work distance and Euclidean distance betweenq and data objects are correlated (i.e., the network distance is close to the Euclidean distance). In [61], Shahabi et al. proposed an road network embedding technique to transform a spatial network to a constraint-free high dimensional Euclidean space to fast but approximately retrieve nearest objects by applying traditional Euclidean based algorithms. Kolahdouzan and Shahabi utilized the first degree network Voronoi diagrams [35, 36] to partition the spatial network to net- work V oronoi polygons (NVP ), one for each data object. They indexed theNVP s with a spatial access method to reduce the problem to a point location problem in Euclidean space and minimize the on-line network distance computation by precomputing the 16 NVPs. Cho et al. [3] presented a system UNICONS where the main idea is to inte- grate the precomputedk nearest neighbors into the Dijkstra algorithm. In [69], Huang et al. addressedkNN problem using Island approach [29] where each vertex is associ- ated (and network distance precomputed) to all the data points that are centers of given radius r (so called islands) covering the vertex. With their approach, they utilized a restricted network expansion from the query point while using the precomputed islands. Aside from their specific drawbacks, these algorithms rely on data object dependent precomputations (i.e., the network distance to the data objects are precomputed) and subdivide the spatial network based on the location of the data objects. Therefore, they assume that data objects are static and/or trajectory of query objects is known. This assumption is undesirable in applications where the query and data objects change their positions frequently. Recently, Huang et al. [28] and Samet et al. [57] proposed two different algo- rithms that address the drawbacks of data object dependent precomputation. Huang et al. introduced S-GRID where they partition (using grid) the spatial network to disjoint sub-networks and precompute the shortest path for each pair of connected border points. To find thek nearest neighbors, they first perform a network expansion within the sub- networks and then proceed to outer expansion between the border points by utilizing the precomputed information. Samet et al. proposed a method where they associate a label to each edge that represents all nodes to which a shortest path starts with this particular edge. They use these labels to traverse shortest path quadtrees that enables geometric pruning to find the network distance between the objects. All these studies simplistically assume the network edge weights are constant, and hence they are invalidated with time-dependent edge weights. With our study, we make a fundamentally different assumption that the cost of the network edges are function of 17 time rather than constant. Our assumption yields a much more realistic scenario and versatile approach forkNN query processing in spatial networks. 2.2 Time-dependent Shortest Path In the last decade, numerous efficient fastest path algorithms with precomputation meth- ods have been proposed (see [58, 59] for an overview). However, there are limited num- ber of studies that focus on efficient computation of time-dependent fastest path (TDFP) problem. Cooke and Halsey [4] first studied TDFP computation where they solved the problem using Dynamic Programming in discrete time. Another discrete-time solution to TDFP problem is to use time-expanded networks [33]. In general, time-expanded network (TEN) and discrete-time approaches assume that the edge weight functions are defined over a finite discrete window of time t2 t 0 ;t 1 ;::;t n , where t n is determined by the total duration of time interval under consideration. Therefore, the problem is reduced to the problem of computing minimum-weight paths over a static static network per time window. Hence, one can apply any static fastest path algorithms to compute TDFP. Although these algorithms are easy to design and implement, they have numerous short- comings. First, TEN models create a separate instance of network for each time instance hence yielding a substantial amount of storage overhead. Second, such approaches can only provide approximate results because the model misses the state of the network between any two discrete-time instants. Moreover, the difference between the shortest path obtained using TEN approach and the optimal shortest path is unbounded. This is because the query time can be always between any two of the intervals which are not captured by the model, and hence the error is is accumulated on each edge along the path. In [21], George and Shekhar proposed a time-aggregated graph approach where 18 they aggregate the travel-times of each edge over the time instants into a time series. Their model requires less space than that of the TEN and the results are still approxi- mate with no bounds. In [16], Dreyfus showed that TDFP problem can be solved by a generalization of Dijkstra’s method as efficiently as for static fastest path problems. However, Halpern [26] proved that the generalization of Dijkstra’s algorithm is only true for FIFO net- works. If the FIFO property does not hold in a time-dependent network, then the prob- lem is NP-Hard. In [47], Orda and Rom introduced Bellman-Ford based algorithm where they determine the path toward destination by refining the arrival-time functions on each node in the whole time interval T . In [32], Kanoulas et al. proposed Time- Interval All Fastest Path (allFP) approach in which they maintain a priority queue of all paths to be expanded instead of sorting the priority queue by scalar values. There- fore, they enumerate all the paths from the source to a destination node which incurs exponential running time in the worst case. In [15], Ding et al. used a variation of Dijk- stra’s algorithm to solve the TDFP problem. With their TDFP algorithm, using Dijkstra like expansion, they decouple the path-selection and time-refinement (computing earli- est arrival-time functions for nodes) for a given starting time intervalT . Their algorithm is also shown to run in exponential time for special cases (see [6]). The focus of both [32] and [15] is to find the fastest path in time-dependent road networks for a given start time-interval (e.g., between 7:30AM and 8:30AM). The ALT algorithm [22] was originally proposed to accelerate fastest path computa- tion in static road networks. With ALT, a set of nodes called landmarks are chosen and then the shortest distances between all the nodes in the network and all the landmarks are computed and stored. ALT employs triangle inequality based on distances to the landmarks to obtain a heuristic function to be used in A* search. The time-dependent variant of this technique is studied in [8] (unidirectional) and [41] (bidirectional A* 19 search) where heuristic function is computed w.r.t lower-bound graph. However, the landmark selection is very difficult (relies on heuristics) and the size of the search space is severely affected by the choice of landmarks. So far no optimal strategy with respect to landmark selection and random queries has been found. Specifically, landmark selec- tion is NP-hard [52] and ALT does not guarantee to yield the smallest search spaces with respect to fastest path computations where source and destination nodes are chosen at random. Our experiments with real-world time-dependent travel-times show that our approach proposes in Section 2 consumes much less storage as compared to ALT based approaches and yields faster response times. In two different studies, The Contraction Hierarchies (CH) and SHARC methods (also developed for static networks) were aug- mented to time-dependent road networks in [1] and [7], respectively. The main idea of these techniques is to remove unimportant nodes from the graph without changing the fastest path distances between the remaining (more important) nodes. However, unlike the static networks, the importance of a node can change throughout the time under consideration in time-dependent networks, hence the importance of the nodes are time varying. Considering the super-polynomial input size (as discussed in Section 1.2), and hence the super-polynomial number of important nodes with time-dependent net- works, the main shortcomings of these approaches are impractical preprocessing times and extensive space consumption. For example, the precomputation time for SHARC in time-dependent road networks takes more than 11 hours for relatively small road net- works (e.g. LA with 304,162 nodes) [7]. Moreover, due to the significant use of arc flags [7], SHARC does not work in a dynamic scenario: whenever an edge cost function changes, arc flags should be recomputed, even though the graph partition need not be updated. While CH also suffers from slow preprocessing times, the space consumption for CH is at least 1000 bytes per node for less varied edge-weights where the storage cost increases with real-world time-dependent edge weights. Therefore, it may not be 20 feasible to apply SHARC and CH to continental size road networks which can consist of more than 45 million road segments (e.g., North America road network) with possibly large varied edge-weights. 21 Chapter 3 Problem Definition In this chapter, we will formally define spatial network, time-dependent travel-time, time-dependent shortest path, time-dependentkNN search in spatial networks, and sev- eral important concepts that we use throughout this proposal. There are various criteria to define the cost of a path in road networks. Throughout this thesis, the cost of a path is defined as its travel-time and the term shortest path is interchangeably used to denote minimum-travel-time (or fastest) paths. 3.1 Time-dependent Spatial Network To make the discussion more general, we introduce the concept of a spatial network which is an extension of a network model. The networks are modeled as a weighted graphG(V;E), where V denotes the set of vertices and E denotes the set of edges. A spatial network is an extension of a network such that additional spatial components (e.g., the spatial position of each vertex with respect to a reference coordinate system) are associated with the vertices and/or edges)of the graph. In this thesis, we assume a spatial network (e.g. the Los Angles road network) con- taining a set of static data objects (i.e., points of interest such as restaurants, hospitals) as well as static or moving query objects searching for their destination or kNN. We model the spatial network as a time-dependent weighted graph (directed) where the non- negative weights are time-dependent travel-times between the nodes. We assume both data and query objects lie on the network edges and all relevant information about the 22 objects is maintained by a central server. As a query object moves, the central server is updated with the new location of the object. Below, we formally define our terminology. Definition 1 Time-dependent Graph A Time-dependent Graph (G T ) is defined as G T (V;E) where V = fv i g is a set of nodes representing the intersections and ter- minal points, andE (E VV ) is a set of edges representing the network segments each connecting two nodes. Each edgee is represented bye(v i ;v j ) wherev i andv j are starting and ending nodes, respectively, and v i 6= v j . For every edge e(v i ;v j )2 E, there is an edge travel-time functionc i;j (t), wheret is the time variable in time domain T . An edge travel-time functionc i;j (t) specifies how much time it takes to travel fromv i tov j starting at timet. For example, Figure 1.2 depicts a road network modeled as a time-dependent graphG T with four nodes and five edges. Definition 2 Lower-bound Graph. Given a G T (V;E), the corresponding Lower- bound Graph G(V;E) is a graph with the same topology (i.e, nodes and edges) as graphG, where the weight of each edgec v i ;v j is fixed (not time-dependent) and is equal to the minimum possible weightc min v i ;v j where8e(v i ;v j )2E;t2T c min v i ;v j c v i ;v j (t). Definition 3 Upper-bound Graph. Given a G T (V;E), the corresponding Upper- bound GraphG(V;E) is a graph with the same topology as graphG, where the weight of each edgec v i ;v j is fixed (not time-dependent) and is equal to the maximum possible weightc max v i ;v j where8e(v i ;v j )2E;t2T c max v i ;v j c v i ;v j (t). Definition 4 Time-dependent Travel Cost. Let fs = v 1 ;v 2 ;:::;v k = dg denotes a path which contains a sequence of nodes where e(v i ;v i+1 )2 E andi = 1;:::;k1. Given aG T (V;E), a path (s d) from source s to destination d, and a departure-time at the sourcet s , the time-dependent travel cost TT(s d;t s ) is the travel-time it takes to travel the path. Since the travel-time of 23 an edge varies depending on the arrival-time to that edge, the travel-time of a path is computed as follows: TT(s d;t s )= k1 X i=1 c v i ;v i+1 (t i ) wheret 1 =t s ,t i+1 =t i +c (v i ;v i+1 ) (t i );i=1;::;k. Definition 5 Lower-bound Travel Cost. The lower-bound travel-timeLTT(s d) of a path is less than the actual travel-time along that path and computed with respect to G(V;E) by considering the minimum possible travel-times of all edges along the path s d, i.e., LTT(s d)= k1 X i=1 c min v i ;v i+1 ,i=1;::;k, Definition 6 Upper-bound Travel Cost. The upper-bound travel-timeUTT(s d) of a path is greater than the actual travel-time of the path and computed with respect to G(V;E) by considering the maximum possible travel-times of all edges along the path s d, i.e., UTT(s d)= k1 X i=1 c max v i ;v i+1 ,i=1;::;k. Note that we do not need to consider arrival-dependency when computingUTT and LTT hence; t is not included in their definitions. Given the definitions of TT , UTT andLTT , the following property holds for any path inG T : LTT(s d) TT(s d;t s ) UTT(s d). We will use this property in subsequent sections to establish some properties of our algorithm. 3.2 Time-dependent Query Processing Definition 7 Time-dependent Shortest Path Given a time-dependent graphG T with a sources2V and a destinationd2V , and a starting timet s 2T , the time-dependent shortest path (also referred as time-dependent fastest path in this thesis)TDSP(s;d;t) 24 is a path with the minimum travel-time among all paths froms tod with the departure- timet s . Definition 8 Time-dependent k Nearest Neighbor Query (TD-kNN) A time- dependentk nearest neighbor query in spatial networks is defined as a query that finds thek nearest neighbors of a query object moving on a time-dependent networkG T . Con- sidering a set ofn data objectsP =fp 1 ;p 2 ;:::;p n g, the TD-kNN query with respect to a query pointq finds a subsetP 0 P ofk objects with minimum time-dependent travel- time toq, i.e., for any objectp 0 2P 0 andp2PP 0 ,TDSP(q;p 0 ;t)TDSP(q;p;t). In the rest of this paper, we assume that the edge travel-time functions are given as positive piece-wise linear functions of time and all piece-wise functions have a finite number of pieces. This is consistent with how traffic trends are reported for a given edge in real-world road networks. We also assume that the spatial networkG T satisfies the First-In-First-Out(FIFO) property [5]. This property suggests that moving objects exit from an edge in the same order they enter the edge. Finally, with our algorithm, we do not allow objects to wait at a node, because, in most real-world applications, waiting at a node is not realistic as it requires the moving object to get out of the current road (e.g., the exit freeway) and find a place to park and wait. 25 Chapter 4 Baseline kNN Search Solutions in Time-dependent Road Networks In this section, we explain two different baseline algorithms to evaluatek nearest neigh- bor queries in time-dependent spatial networks. With the first approach, we model the time-dependent road network as a time-expanded graph [33, 48] that approximates the time-dependent network with a snapshot of the network in each time interval. With the second approach, we generalize the incremental network expansion [49] method (pro- posed for static spatial networks) to time-dependent road networks. Our incremental network expansion method uses the time-dependent arrival times as the labels of the nodes to form the greedy search [16]. Figure 4.1 illustrates an example of time-dependentk nearest neighbor search. With this example, an ambulance is looking for the nearest hospital at 8:30AM and 2:00PM on the same day on a particular road network. Note that the travel times on the edges (the number shown on the edges) change with time in Figures 4.1(a) and 4.1(b). Therefore, the queries launched by the ambulance at 2PM and 5PM return different results. 26 (a) 1-NN Query at 8:30 AM (b) 1-NN Query at 2 PM Figure 4.1: Time-dependent 1-NN search 4.1 Time-dependent kNN Search using Time-Expanded Networks Given a time-dependent graph G T (V;E), a time-expanded model discretizes the time domainT =[t 0 ;t n ] inton points of time, and constructs a static graphG(V;E) by mak- ingn copies of each node and each edge, respectively. Specifically, time-expanded net- work replicates the original network for each discrete time unitt=0;1;:::;t n , wheret n is determined by the total duration of the time interval under consideration. This model connects a node and its copy at the next instant in addition to the edges in the origi- nal network, replicated for every time instant. The weight of an edge in time-expanded network is the time difference between the time events associated with its endpoints. Therefore, a time-dependent edge costs can be interpreted as a static flow in the corre- sponding time-expanded network. Figure 4.3 shows four consecutive snapshots (taken every 10 minutes) and the corresponding time-expanded model of the time-dependent network in Figure 4.2. In this figure, for example, the weight (i.e., travel-time) of edge (v 1 ;v 2 ) att = 0 is represented by connecting the copy of nodev 1 att = 0 to the copy of nodev 2 att=20 (see Figure 4.3(e)). 27 (a) GraphG T (b) c 1;2 (t) (c) c 2;3 (t) (d) c 2;4 (t) (e) c 4;5 (t) (f) c 3;5 (t) change Figure 4.2: A Time-dependent GraphG T (V;E) (a) t 0 =0 (b) t 1 =10 (c) t 2 =20 (d) t 3 =30 (e) Time-expanded model Figure 4.3: A Time-expanded graph The time-expanded network approach enables time-dependent k nearest neighbor problem to be solved by applying any precomputation techniques developed for static networks (e.g., [57]). However, there are two drawbacks with any solution based on 28 time-expanded networks. First, since the original network is replicated across time instants, the size of the network increases hence, resulting in high storage overhead and slower response time. The storage requirement for a time expanded-network is O(jVjT) +O(jVj +jEjT), where T is the total number of snapshots. Second, the difference between the shortest path obtained using time-expanded model and the opti- mal shortest path is very sensitive to the parametern, and is unbounded. Because, the query time (or the arrival-time to an edge) can always be between any two time points (e.g., between t 0 and t 1 ), but the edge weights are only captured in either of the time points. For example, consider a shortest path query executed att = 12 in Figure 4.3, and an error between the optimal path and the path found using time-expanded net- work model. In this case, the network snapshot att=10 is used to compute the shortest path for t = 12 and is accumulated on each edge along the path. Our experiments show that the error rate is especially high during rush hours (see 4.3), hence causing time-expanded models to generate inaccurate results. 4.2 Time-dependent kNN Search using Network Expan- sion In this section, we propose an algorithm that generalizes the incremental network expan- sion method [49] (originally proposed for static road networks) to answer k nearest neighbor queries in time-dependent road networks. With this algorithm, starting from the query objectq all network nodes reachable fromq in every direction are visited in order of their proximity (i.e., time-dependent travel-time) to q until all k nearest data objects are located. We use four main data structures to enable the network expan- sion solution. The adjacency component captures the network connectivity. The edge component includes the poly-line representation of each network edge(u;v), length of 29 the edge, and a pair of pointers to the disk pages containing the adjacency lists of its endpoints u, v. The travel-time component includes the travel-time functions of each network edge. We use hash table to associate travel-time functions to network edges. The last component is R-tree [25] that indexes the edges’ MBRs. Each leaf entry of R-tree contains a pointer to the disk page storing the corresponding edge. With this approach, we adopt a variant of Dijkstra algorithm (proposed in [16]) to expand the network based on the time-dependent travel-time to each node around q until k objects are found. We outline our proposed approach in Algorithm 6. The algorithm takes three parameters as the input, i.e., query locationq, number of desired nearest neighbors k, and query time t q . We first execute a findEdge(q) operation to find the edge that containsq by performing a point location query on R-tree index. Let t(v) denote the time taken to travel fromq to nodev along the shortest path in a time- dependent network, i.e., time to travel forTDSP(q;v;t q ). Analogous to shortest path distances, for every edgee(u;v), we havet(v) f e (t(u)) and this forms the basis of our greedy algorithm. Note that f e (t(u)) is the time taken to travel from q to u plus the time travel from u to v. We use a priority queue S to keep track of the nodes to be examined. WithS, we maintain a) the set explored nodes (which includes the nodes for which we have calculated the actual t(v)), and b) the label l(v) for the nodes not in S, where the label l(v) is our current estimate for the least time it takes to reach v fromq. As shown in Algorithm 6, we updatel(v) based onmin(l(v);f e (t(u))) (Line 8) only considering the edges (u;v) whereu2 S. Finally, we pick the nodew = 2 S with smallestl(:) value and add it to the setS and repeat. If the recently added node is a data object (data objects can easily be modeled as network nodes), we add that data object to nearest neighbor array NN and accordingly compute its travel time (Line 11-13). This process is repeated until the algorithm findsk data objects. It is important to note that Algorithm 6 holds for FIFO networks in which greedy property is maintained. 30 Algorithm 1: kNN(q,k,t q ) 1: //S: set of nodes,q: query location,dt: departure-time fromq 2: //tt: travel-time of the fastest path,v i : last node added toS 3: //NN: array of current NNs 4: InitializeS =fqg;t(q)=0;l(v)=1 for allv = 2S 5: v i =q 6: whileS6=V do 7: foreache(v i ;v j )2Ewherev j = 2S 8: l(v j )=min(l(v j );f e (t(v i ))) 9: Letw = 2Ssuchthatl(w)=min v j = 2S l(v j ) 10: S =S S fwg;t(w)=l(w);v i =w 11: Ifv i isadataObject Then 12: addv i toNN; 13: tt=t(v i )t q ; //compute travel-time tov i 14: End If; 15: IfNN:size()=k Then break; 16: end while 17: returnNN and travel times Lemma 1 Algorithm 6 is correct Proof 1 The proof of correctness for Algorithm 6 follows that of the Dijkstra algorithm. Let us consider any node w, and a set S just before w is added to S. Let P w and t w represent the pathq w and the time it takes to travel alongs w, respectively. Note thatt(w) = l(w) = min u 0 2S:(u 0 ;w)2E f u 0 ;w (t(u 0 )), so ifu 0 2 S is the node that attains the minimum in l(w), then P w is obtained by adding the edge (u 0 ;w) to the path P 0 u (which is obtained recursively). Now we show thatt(w) is indeed the least time it takes to reach w along any path q w. Consider any q w path P . Let v be the first node onP that is not inS, andu2 S be the node just beforev. LetP 0 be the portion of the path P from q to v and t 1 and t 2 represent the times at which we reach u and v, respectively, by traveling along the pathP . Lete 1 = (v;v 1 );e 2 ;:::;e k = (v k1 ;w) be the portion of the path fromv tow. Then, the time it takes to reachw by traveling along P is TT(P) = f e k (f e k1 (:::f e 2 (f e 1 (t 2 )):::)) t 2 , because8e f e (t) t. Also, 31 t 2 = f u;v (t 1 ) f u;v (t(u)) sincet 1 t(u) and thef e function is monotonic. Thus, we haveTT(P) f u;v (t(u)) l(v) l(w), where the last two inequalities follow from the definition ofl(:), and since we choosew to add to the setS. This holds for any path P , and concludes the proof. 4.3 Performance Evaluation We conducted several experiments with different spatial networks and various parame- ters (see Table 1) to evaluate the performance of both TD-kNN algorithms. As our road network dataset, we use Los Angeles (LA) and San Joaquin (SJ) road networks with 304,162 and 24,123 road segments, respectively. We evaluate our proposed techniques using both syntectic and actual time-dependent travel-times gathered from real-world traffic sensor data. To generate time-dependent edge costs (travel-time) we use real- world traffic sensor dataset that we have been collecting (past 2 years) and archiving from a collection of approximately 7000 sensors located on the road network of Los Angeles. We collect speed, occupancy, volume information from these sensors and the sampling rate of the data is 1 reading/sensor/min. We spatially and temporally aggregate (average) historical sensor data based on 7 days (Monday to Sunday) of each month by assigning interpolation points for each 5 minutes. The interpolation points represent the travel-times at different times of a particular day. For example, an edge is assigned 180 travel-time attributes to represent how traffic tends to change between 6:00AM and 9:00PM for a particular date in a particular month, e.g., Monday traffic pattern in September. We assume all roads are un-congested between 9:00PM and 6:00AM, and hence consider static edge weights during this interval. However, unfortunately not every edge has a sensor in road networks. In order to generate time-dependent edge weights onSJ and for the edges that does not contain any sensor inLA, we developed a 32 Table 4.1: Experimental parameters Parameters Default Range Number of objects 10 (K) 1,5,10,15,20(K) Number of queries 3 (K) 1,2,3,4,5 (K) Number of k 20 1,10,20,30,40,50 Object Distribution Uniform Uniform, Gaussian Query Distribution Uniform Uniform, Gaussian traffic modeling approach that creates edge travel-time profiles [13]. Our approach uses spatial (e.g., locality, connectivity) and temporal (e.g., rush hour, weekday) characteris- tics to generate travel-time of network edges that does not have readily available sensor data. We generated the parameters represented in Table 1 using a simulator prototype developed in Java. We conducted our experiments on a workstation with 2.7 GHz Pen- tium Core Duo processor and 12GB RAM memory. We computed the time-expanded network model of both LA and SJ networks by discritizing the networks for each 5 minutes corresponding to our interpolation points. Similar to Algorithm 6 (denoted by TD-NE), we implemented a network expansion method to find k nearest neighbors in time-expanded networks (denoted by TE). We continuously monitored each query for 50 timestamps in both of the implementations. For each set of experiments, we only vary one parameter and fix the remaining to the default values in Table 1. Since the experimental results with both LA and SJ networks differ insignificantly, we only present the results from LA dataset. 33 Correctness and Impact of k With this experiment, we compare the correctness of the two algorithms (i.e., percentage of correctly identified nearest neighbors). Figure 4.4(a) plots the correctness versus time ranging from 6 am to 6 pm, while using default settings in Table 1 for all other param- eters. As shown, while TD-NE returns correct results all the time, TE’s correctness is substantially low around rush hours (i.e., 7-9 am, 4-6 pm). This is because time- dependent weights of each network segment change rapidly especially at the boundaries of the traffic peak periods, resulting the error accumulating along the path. (a) Correctness versus time (b) Impact of k on response time Figure 4.4: Correctness and impact of k Next, we compare the performance of the two algorithms with regard tok. Figure 4.4(b) shows the average query efficiency versus k ranging from 1 to 50. The results indicate that TD-NE outperforms TE with all values ofk and the response time of both algorithms increases with the large values ofk. Note that the slower response time of TE in this and the following experiments is due to increased size of the network because of replication. It is important to note that one can adopt a precomputation technique (proposed for static networks) to accelerate the response time of TE. 34 Impact of Object/Query Distribution and Network Size With this experiment, we study the impact of object and query distribution as well as net- work size. Figure 5.13 shows the response time of both algorithms where the objects and queries follow either uniform or Gaussian distributions. As illustrated, TD-NE yields better performance for queries with Gaussian distribution. This is because as queries are clustered in the spatial network with Gaussian distribution, their nearest neighbor would overlap; hence, allowing TD-NE to save computation. In addition, we measured the performance of both algorithms with respect to the net- work size. In order to evaluate the impact of network size, we conducted experiments with the sub-networks of LA dataset ranging from 50K to 250K segments. Figure 4.5(b) illustrates the response time of both algorithms with different network sizes. In general, with the default parameters in the Table 1, the response time increases for both algo- rithms as the network size increases. (a) Impact of object distribution (b) Impact of network size Figure 4.5: Response time versus distribution and network size 35 Impact of Object and Query Cardinality With this set of experiments, we compare the performance of the two algorithms by varying the cardinality of the data objects (P) from 1K to 20K while using default set- tings in Table 1 for all other parameters. Figure 5.11 illustrates the impact of the growing object cardinality on response time. The results indicate that the response time linearly increases with the number of data objects in both methods, where TD-NE outperforms TE for all cases. From P=1K to 5K, the response time is slower. Because, since the objects are sparsely distributed when P is small, network expansion visits more redun- dant network nodes causing extra processing time. Figure 5.12 shows the impact of the query cardinality (Q), ranging from 1K to 5K, on response time. As shown, TD-NE scales better with large number of Q and the performance gap between the approaches increases as Q grows. (a) Impact of object cardinality (b) Impact of query cardinality Figure 4.6: Response time versus query/object cardinality and agility 4.4 Summary of Results In this chapter, we studied the problem ofk nearest neighbor queries in time-dependent spatial networks. We formulated a generalized type ofk nearest neighbor query where 36 we, unlike the existing studies, assume the edge weights of the network are time vary- ing rather than fixed. We proposed two baseline solutions by exploiting time-expanded network and network expansion framework. We evaluated and compared the efficiency of these solutions with real-world data-sets, including a variety of large spatial networks with real traffic-data. Although time-expanded network framework provides a mecha- nism to use existingkNN algorithms (and their precomputation techniques) developed for static networks, the experimental results suggest that the error rate (incorrectly iden- tified nearest neighbors) of this approach is very high especially during traffic peak hours. On the other hand, while network expansion yields correct results at all times, the overhead of executing network expansion is very high in large networks with a sparse set of data objects. Hence, the network expansion approach is not suitable for onlinek nearest neighbor applications. In the next chapter, we propose an algorithm that address the shortcomings of the both baseline solutions. 37 Chapter 5 Efficient kNN Query Processing in Time-Dependent Spatial Networks In the previous chapter, we introduced two baseline solutions for the time-dependent kNN problem based on the time-expanded networks and incremental network expan- sion. With time-expanded graphs the time domain is discretized and at each discrete time instant a snapshot of the network is used to represent the network. Hence, the time-dependent kNN problem is reduced to the problem of computing the minimum- weight paths through a series of static networks. Although this approach allows for exploiting the existing algorithms and precomputations forkNN computation on static networks, it often fails to provide the correct results because the model misses the state of the network between any two discrete time instants. Secondly, we developed a solu- tion based on the incremental network expansion approach where Dreyfus’s modified Dijkstra algorithm is used for time-dependent distance calculation. With this approach, starting from a query objectq all network nodes reachable fromq are visited in order of their time-dependent travel-time proximity toq until allk nearest objects are located (i.e., blind network expansion). However, considering the prohibitively high overhead of executing blind network expansion particularly in large networks with a sparse set of data objects, this approach is far too slow to scale for real-timekNN query processing. In this chapter, we address the disadvantages of both baseline approaches by devel- oping a novel technique that findskNN of a query object in time-dependent road net- works. Our solution a) efficiently and accurately answers the queries in order to support 38 moving object kNN search on time-dependent networks, b) is independent of density and distribution of the data objects, and c) effectively handles the database updates where nodes, links, and data objects are added or removed. Towards that end we develop two types of complementary index structures. The core idea behind these index struc- tures is to localize the search space and minimize the costly time-dependent shortest path computation between the objects hence incurring low computation costs. Our proposed algorithm involves two phases: an off-line time-dependent network indexing phase and an on-line query processing phase. In the off-line phase, we parti- tion the time-dependent spatial network into Tight Cells (TC) and Loose Cells (LC) for each data object p and construct two complementary indexing schemes, namely Tight Network Index (TNI) and Loose Network Index (LNI). In the on-line phase, we useTNI andLNI structures to efficiently answerk nearest neighbor queries. Specifically, with TNI, we can find the nearest objects without performing any shortest path computation. Our experiments show that in 70% of the cases the nearest neighbor can be found with this index. For those cases that the nearest objects cannot be identified by TNI, LNI allows us to filter in only a small number of objects that are potential candidates (and filter out the rest of the objects). Subsequently, we only need to perform the shortest path computation only for these candidates. In the following sections, we will first mention the construction of our spatial index structures and then describe our query processing algorithm that utilizes these index structures. 5.1 Index Construction In this section, we present the construction of our tight and loose network index struc- tures. 39 5.1.1 Tight Index Construction Before we explain the tight index structure, we introduce the concept of tight cell and it’s properties. A Tight CellTC(p i ) is a sub-network around each data objectp i in which any query object is guaranteed to have p i as its nearest neighbor in a time-dependent spatial network. We compute the tight cell of a data object by utilizing the parallel Dijkstra algorithm. Specifically, we expand fromp i (i.e., the generator of the tight cell) assuming upper-bound travel-time (i.e., UTT), while in parallel we expand from each and every other data object assuming the lower-bound travel-time (i.e., LTT). We stop the expansions when the shortest path trees meet. We repeat the same process for each data object to compute its tight cell. The main idea behind tight cells is that if the upper-bound travel time between the query objectq and a particular data objectp i is less than any of the lower-bound travel time fromq to any other data objectp j , then we can conclude thatp i is guaranteed to be the nearest neighbor of q. Figure 5.1 depicts the network expansion from the data objects during the tight cell construction forp 1 . For the sake of clarity, we represent the tight cell of each data object with a polygon as shown in Figure 5.2. We generate the edges of the polygons by connecting the adjacent border nodes (i.e., nodes where the shortest path trees meet) of a generator to each other. We refer the readers to [17] for the implementation details of parallel Dijkstra algo- rithm that we used to construct tight cells. Note that the complexity of computing the tight cells with parallel Dijkstra is asymptotically not worse than the complexity of Dijk- stra’s algorithm (i.e.,O(jEj+jVjlogjVj). The formal lemma and the proof of tight cells is as follows. Lemma 2 LetP be a set of data objectsP =fp 1 ;p 2 ;:::;p n g inG T andTC(p i ) be the tight cell of a data objectp i . For any query pointq2TC(p i ), the nearest neighbor ofq isp i , i.e.,f8q2TC(p i ),8p j 2P;p j 6=p i ,TDSP(q;p i ;t)<TDSP(q;p j ;t)g. 40 Figure 5.1: Tight cell construction ofP 1 Figure 5.2: Tight Cells Proof 2 We prove the lemma by contradiction. Assume thatp i is not the nearest neigh- bor of the query objectq. Then there exists a data objectp j (p i 6= p j ) which is closer toq; i.e.,TDSP(q;p j ;t) < TDSP(q;p i ;t). Let us now consider a pointb (where the shortest path trees ofp i andp j meet) on the boundary of the tight cellTC(p i ). We denote shortest upper-bound path fromp i tob (i.e., the shortest path among allUTT(p i b) paths) asD UTT (p i ;b), and similarly, we denote shortest lower-bound path fromp j to b (i.e., the shortest path among allLTT(p j b) paths) asD LTT (p j ;b). Then, we have TDSP(q;p i ;t) < D UTT (p i ;b) = D LTT (p j ;b) < TDSP(q;p j ;t). This is a contradic- tion; hence,TDFP(q;p i ;t)<TDFP(q;p j ;t). With our proposed algorithm, we utilize TCs in the following way. If a query point q is inside a specific TC(p i ), one can immediately identify the generator of that tight cell (i.e., p i ) as the nearest neighbor for q. This stage can be expedited by using a spatial index structure generated on the TCs. Although TCs are constructed based on the network distance metric, each TC is actually a polygon in Euclidean space. Therefore, TCs can be indexed using spatial index structures (e.g., R-tree [25]). This way a function (e.g., contain(q)) invoked on the spatial index structure would efficiently return the TC 41 whose generator has the minimum time-dependent network distance toq. We formally define Tight Network Index as follows. Definition 9 Tight Network Index (TNI). Let P be the set of data objects P = fp 1 ;p 2 ;:::;p n g, the Tight Network Index is a spatial index structure generated on tight cells ofP ,fTC(p 1 );TC(p 2 );:::;TC(p n )g. As illustrated in Figure 5.2, the set of tight cells often does not cover the entire network. For the cases whereq is located in an area which is not covered by any tight cell, we utilize the Loose Network Index (LNI) to identify the candidate nearest data objects. Next, we describe loose network index. 5.1.2 Loose Index Construction The loose cellLC(p i ) is a sub-network aroundp i outside which any point is guaranteed not to have p i as its nearest neighbor in a time-dependent spatial network. In other words, the data objectp i is guaranteed not to be the nearest neighbor ofq ifq is outside of the loose cell ofp i . Similar to the construction process forTC(p i ), we use the parallel shortest path tree expansion to constructLC(p i ). However, this time, we use minimum travel-time between the nodes of the network (i.e., LTT ) to expand from p i (i.e., the generator of the loose cell) and maximum travel-time (i.e.,UTT ) to expand from every other data object. Lemma 3 proves the property ofLC: Lemma 3 Let P be a set of data objects P =fp 1 ;p 2 ;:::;p n g in G T and LC(p i ) be the loose cell of a data objectp i . Ifq is outside ofLC(p i ), p i is guaranteed not to be the nearest neighbor of q, i.e.,f8q 62 LC(p i );9p j 2 P;p j 6= p i ;TDSP(q;p i ;t) > TDSP(q;p j ;t)g. Proof 3 We prove the lemma by contradiction. Assume thatp i is the nearest neighbor of aq, even though theq is outside ofLC(p i ); i.e.,TDSP(q;p i ;t) < TDSP(q;p j ;t). 42 Figure 5.3: Loose Cells Suppose there exists a data object p j whose loose cell LC(p j ) covers q (such a data object must exist, because as we will next prove that set of loose cells cover the entire network). Letb be a point on the boundary ofLC(p i ). Then, we have,TDSP(q;p j ;t)< D UTT (p j ;b)=D LTT (p i ;b)<TDSP(q;p i ;t). This is a contradiction; hence,p i cannot be the nearest neighbor ofq. As illustrated in Figure 5.3, loose cells, unlike TCs, collectively cover the entire network and have some overlapping regions among each other. Lemma 4 Loose cells collectively cover the network and they may overlap. Proof 4 As we mentioned, during loose cell construction, LTT is used for expansion from the generator of the loose cell. Since the parallel Dijkstra algorithm traverses every node until the priority queue is empty as described in [17], every node in the network is visited; hence, the network is covered. Since the process of expansion withLTT is repeated for each data object, in the overall process some nodes are visited more than once; hence, the overlapping areas. Therefore, loose cells cover the entire network and may have overlapping areas. Note that if the edge weights are constant, the LCs would not overlap, and TCs cells and LCs would be the same. 43 Based on the properties of tight and loose cells, we know that loose cells and tight cells have common edges (i.e., all the tight cell edges are also the edges of loose cells). We refer to data objects that share common edges as direct neighbors and remark that loose cells of the direct neighbors always overlap. For example, consider Figure 5.3 where the direct neighbors ofp 2 arep 1 , p 3 , andp 6 . This property is especially useful for processing k-1 neighbors (see Section 5.2.2) after finding the first nearest neighbor. We determine the direct neighbors during the generation of the loose cells and store the neighborhood information in a data component. Therefore, finding the neighboring cells does not require any complex operation. Similar toTNI, we can use spatial index structures to access loose cells efficiently. We formally define the Loose Network Index (LNI) as follows. Definition 10 Loose Network Index (LNI). Let P be the set of data objects P = fp 1 ;p 2 ;:::;p n g, the Loose Network Index is a spatial index structure generated on loose cells ofP ,fTC(p 1 );TC(p 2 );:::;TC(p n )g. Note that LNI and TNI are complementary index structures. In particular, if a q cannot be located with TNI (i.e., q falls outside of any TC), then we use LNI to identify theLCs that containq; based on Lemma 3, the generators of suchLCs are the only NN candidates forq. 5.1.3 Tight and Loose R-Tree With our approach, we adopt R-Tree [25] data structure to implement TNI and LNI, termed TN R-tree and LN R-tree, respectively. Figure 5.4 depicts LN R-tree (TN R-tree is a similar data structure without extra pointers at the leaf nodes, hence not discussed). As shown, LN R-tree has the basic structure of an R-tree generated on minimum bound- ing rectangles of loose cells. The difference is that we modify R-tree by linking its 44 leaf nodes to the the pointers of additional components that facilitate TD-kNN query processing. These components are the direct neighbors (N(p i )) of p i and the list of nodes (VL p i ) that are insideLC(p i ). WhileN(p i ) is used to filter the set of candidate nearest neighbors wherek > 1, we useVL p i to prune the search space during TDSP computation (see Section 5.3). Figure 5.4: LN R-Tree Our proposed index structures need to be updated when the set of data objects and/or the travel-time profiles change. Fortunately, due to local precomputation nature of TD- kNN, the affect of the updates with both cases are local, hence requiring minimal change in tight and loose cell index structures. Below, we explain each update type. Data Object Updates: We consider two types of object update; insertion and deletion (object relocation is performed by a deletion following by insertion at the new location). With a location update of a data objectp i , only the tight and loose cells ofp i ’s neighbors are updated locally. In particular, when a newp i is inserted, first we find the loose cell(s) LC(p j ) containingp i . Clearly, we need to shrinkLC(p j ) and since the loose cells and tight cells share common edges, the region that contains LC(p j ) and LC(p j )’s direct neighbors needs to be adjusted. Towards that end, we find the neighbors of LC(p j ); the tight and loose cells of these direct neighbors are the only ones affected by the 45 insertion. Finally, we compute the new TCs and LCs forp i ,p j andp j ’s direct neighbors by updating our index structures. Deletion of ap i is similar and hence not discussed. Edge Travel-time Updates: With travel-time updates, we do not need to update our index structures. This is because the tight and loose cells are generated based on the minimum (LTT) and maximum (UTT) travel-times of the edges in the network that are time-independent. The only case we need to update our index structures is when minimum and/or maximum travel-time of an edge changes, which is not that frequent. Moreover, similar to the data object updates, the affect of the travel-time profile update is local. When the maximum and/or minimum travel-time of an edge e i changes in the network, we first find the loose cell(s)LC(p j ) that overlaps withe i and thereafter recompute the tight and loose cells ofLC(p j ) and its direct neighbors. As mentioned, given a query pointq, tight and loose cells in a time-dependent road network, the first step in answering a kNN query is to locate the network tight or loose cell that contains q. Considering the large size of the underlying space (e.g., a conti- nental size road network) with numerous data objects as well as the online nature of the queries that requires fast response-time, an index structure is necessary to efficiently access the portion of the network associated withq. Without loss of generality, we used R-tree index structures to index tight and loose cells. We note that although tight and loose cells are created based on the network distance, they are treated as regular poly- gons (by connecting border points) to be able to use R-tree. This approach in some cases may cause misclassification of the network edges (i.e., false-negative edges) in the cells due to network topology. For example, a network edge not belonging to the tight cell ofp i may be classified as a member ofTC(p i ) (see an example of such misclassi- fication in Appendix A). One solution to avoid inaccurate results due to false-negative edges is to perform a refinement step. In particular, one can maintain false-negative edges (along with their corresponding tight or loose cell generators) in a separate data 46 structure. This structure can be checked before each index scan of TNI and LNI. Ifq is located in any of the false-negative edges, the corresponding tight (or loose) cell gener- ator is returned as the result. Otherwise, the search is continued based on TNI and LNI as explained above. It is possible to overcome extra refinement step and hence improve the performance of kNN search by using quad-tree index on tight and loose cells. In Appendix A, we discuss a quad-tree index structure to index network V oronoi diagram (NVD) which partitions the road network into V oronoi cells similar to tight and loose cells. This approach can easily be extend to index tight and loose cells. 5.2 Time-dependent kNN Query Processing So far, we have defined the properties of TNI and LNI. We now explain how we use these index structures to processkNN queries in time-dependent spatial networks. As we discussed, the tight cells do not cover the entire area thus causing unclassified areas in between them. On the other hand, loose cells cover the entire area but they overlap with each other. We proved that if the query objectq is inside one of the tight cells, the generator of that tight cell is the first nearest neighbor ofq. Ifq falls in to an unclassified area between the tight cells, then we investigate the loose cells that containq, the gener- ator of these loose cells are the only candidates for the first nearest neighbor. Hence, we will only compute the (point-to-point) time-dependent shortest path distance between q and those candidates. Clearly, this is much more efficient as compared to the naive way of computing time-dependent network distance fromq to all the sites. After finding the first nearest neighbor we need to find the k-1 nearest neighbors. In order to find the k-1 nearest neighbor efficiently, we use the direct neighbor property. Below, we first describe our algorithm to find the nearest neighbor (i.e.,k=1), and then we extend it to address thekNN case (i.e.,k1). 47 5.2.1 Nearest Neighbor Query Our proposed approach to determine the nearest neighbor of a query object is based on TNI and LNI discussed above. We present the algorithm to process the nearest neighbor query in Algorithm 2. Given the location of a query objectq, first we carry out a depth-first search from theTNI root to the node that containsq (Line 5 of Algorithm 2). If a tight cell that containsq is located, we return the generator of that tight cell as the first nearest neighbor. Our experiments show that, in most cases (7 out of 10), we can findq withTNI search (see Section 5.4). If we cannot locateq inTNI (i.e., when q falls outside all tight cells), we proceed to searchLNI (Line 7). At this step, we may find one or more loose cells that containq. Based on Lemma 3, the generators of these loose cells are the only possible candidates to be the NN forq. Therefore, we compute time-dependent shortest path (TDSP) to find the distance betweenq and each candidate in order to determine the first NN (Line 8-12). We store the candidates in a minimum heap based on their travel-time toq (Line 10) and retrieve the nearest neighbor from the heap (Line 12). Algorithm 2: NN-Algorithm(q,t,TNI,LNI) 1: //q: location of the query object,t: query time 2: //S: an array containing the candidate set, 3: //H: a minimum heap,p: the first NN 4: InitializeS andH; 5: p contain TNI (q); 6: ifp is null then 7: S contain LNI (q); 8: for each data objects i inS do 9: computeTDSP(q;s i ;t); 10: inserts i toH; 11: end for 12: p deHeapH; 13: end if 14: returnp; 48 The time complexity of Algorithm 2 is as follows. The major time consuming steps are traversing TNI or (if necessary) LNI, and point-to-point TDSP computation. There- fore, the complexity of the index search isO(2log(P)) in the worst case whereP is the total number of data objects. The complexity of the point-to-point TDFP is same as in Dijkstra (see Section 5.3). Thus, the total time complexity isO(log(P)+S(E+VlogV)) where S is the number of overlapping cells. 5.2.2 kNN Query Our proposed algorithm for finding the remaining k-1 nearest neighbor is based on the direct neighbor property discussed above. We argue that the second nearest neighbor must be among the direct neighbors of the first NN. Once we identify the second NN, we continue by including the neighbors of the second NN to find the third NN and so on. This search algorithm is based on the following Lemma which is derived from the properties ofTNI andLNI. Lemma 5 The i-th nearest neighbor ofq is always among the neighbors of the i-1 near- est neighbors ofq. Proof 5 We prove this lemma by induction. We prove the base case (i.e., the second NN is a direct neighbor of the first NN ofq) by contradiction. Consider Figure 5.5 where p 2 is the first NN of q. Lets assume that p 5 (which is not a direct neighbor of p 2 ) is the second NN of q. Since p 2 and p 5 are not direct neighbors, a point w on the time- dependent shortest path betweenq andp 5 can be found that is outside bothLC(p 2 ) and LC(p 5 ). However, due to Property 2,p 5 cannot be a candidate NN forw, becausew is not inLC(p 5 ). Thus, there exists another object such asp 1 for instance which is closer tow as compared top 5 . Therefore,TDSP(w;p 5 ;t) > TDSP(w;p 1 ;t). However, as shown in Figure 5.5, we haveTDSP(q;p 5 ;t) = TDSP(q;w;t)+TDSP(w;p 5 ;t) > 49 TDFP(q;w;t)+TDFP(w;p 1 ;t) = TDFP(q;p 1 ;t). Thus,p 5 is farther fromq than bothp 2 andp 1 , which contradicts the assumption thatp 5 is the second NN ofq. Let us now prove the inductive step. The proof of inductive step is straight forward and similar to the above proof by contradiction. Suppose the inductive hypothesis holds for k-1, we prove that it also holds fork. LetS =fp 1 ;p 2 ;:::;p k1 g be the k-1 nearest neighbors of the query objectq, we prove that the k-th neighborp k is among the neighbor cells of S. Consider ak-th nearest neighbor pointp k which is not neighbor of S. Then, there exists a point w on the time-dependent shortest path from q to p k where w does not belong to any of the LC(p k ) or LC(p 1 ),...,LC(p k1 ). Thus, p k is not the nearest neighbor ofw. Suppose the nearest neighbor ofw isp i wherep i 6=p 1 ;p 2 ;:::;p k . Hence TDFP(q;p k ) = TDFP(q;w)+TDFP(w;p k ) TDFP(q;w)+TDFP(w;p i ) = TDFP(q;p i ). Therefore, since p k is farther from q than p 1 ,p 2 ,...,p k1 and p i which contradicts the assumption thatp k is the k-th nearest neighbor. Figure 5.5: Second NN example The complete TD-kNN query process is given in Algorithm 3. Algorithm 3 calls Algorithm 2 to find the first NN and add it to an arrayN, which maintains the current set of nearest neighbors (Lines 4-5). To find the remaining k 1 NNs, we expand 50 the search area by including the direct neighbors of the first NN. Specifically, we add all direct neighbors of the current NN set to a candidate set to find the next NN. We compute the TDSP for each candidate by inserting them to a minimum heap (Lines 7-9 ) based on its time-dependent travel-time to q. Thereafter, we select the one with minimum distance as the second NN (Line 11). Once we identify the second NN, we continue by investigating the neighbor loose cells of the second NN to find the third NN and so on. Our experiments show that the average number of neighbors for a data object is a relatively small number less than 9 (see Section 5.4). Note that, with Algorithm 3, we use a min heap data structure to store the candidate nearest neighbors. This allows us to report the results incrementally even without a pre-specified value ofk. Algorithm 3:kNN-Algorithm(q,k,t,TNI,LNI) 1: //q:location of the query object, k: number of NN 2: //t: query time,N: an array of NN set, H: Min Heap 3: Initialize H, N 4: p NN-Algorithm(q;t;TNI;LNI); 5: addptoN 6: whileN:sizek do 7: for each direct neighborp i of N do 8: computeTDSP(q;p i ;t) 9: addp i toH 10: end for 11: p deheapH; //find next NN 12: addptoN 13: end while 14: ReturnN The time complexity of Algorithm 3 is as follows. The major time consuming step of Algorithm 3 is the point-to-point TDSP computation (Line 7-9) to find k-1 neighbors among the direct neighbors of the current kNN set. Let C be the average number of neighbors for each site. Then, in the worst case, there are kC candidates. Since the complexity of the point-to-point TDSP is same as in Dijkstra (see Section 5.3), the total complexity isO(kCE +kCVlogV) plus the complexity of the Algorithm 3. Our 51 experiments with real-world datasets show thatC is a relatively small number less than 9 (see Section 5.4). 5.3 Time-dependent Shortest Path Computation As we explained, once the nearest neighbor of q is found and the candidate set for the second NN is determined based on the direct neighbors, we need to compute time- dependent shortest path fromq to all candidates in order to find the second NN (and so on). Before we explain our TDSP computation, we note a very useful property of loose cells. That is, given p i is the nearest neighbor of q, the time-dependent shortest path fromq top i is guaranteed to be inLC(p i ) (see Lemma 6). This property indicates that we only need to consider the edges contained in the loose cell of p i when computing TDSP fromq top i . Obviously, this property allows us to localize the time-dependent shortest path search by extensively pruning the search space. Since the localized area of a loose cell is substantially smaller as compared to the complete graph, the computation cost of TDSP is significantly reduced. Note that the subnetwork bounded by a loose cell is on average1=n of the original network wheren is the total number of sites. Lemma 6 Ifp i is the nearest neighbor ofq, then the time-dependent shortest path from q top i is guaranteed to be inside the loose cell ofp i Proof 6 We prove by the lemma contradiction. Assume that p i is the NN of q but a portion of TDFP from q to p i passes outside of LC(p i ). Suppose a point l on that outside portion of the path. Since l is outside LC(p i ), then9p j 2 P , p j 6= p i that 52 satisfies D LTT (p i ;l) > D UTT (p j ;l) and hence TDFP(p i ;l;t) > TDFP(p j ;l;t). Then, TDFP(p i ;q;t) = TDFP(p i ;l;t) + TDFP(l;q;t) > TDFP(p j ;l;t) + TDFP(l;q;t) = TDFP(p j ;q;t), which contradicts the fact that p i is the NN of q. We note that for TD-kNN withk > 1, the TDFP fromq to thekth nearest neighbor will lie in the combined area of neighboring loose cells. Figure 5.6 shows an example query with k > 1 where p 2 is assumed to be the nearest neighbor (and the candidate neighbors ofp 2 are,p 1 ,p 6 andp 3 ). To compute the TDFP fromq to data objectp 1 , we only need to consider the edges contained inLC(p 1 )[LC(p 2 ). Below, we explain how we compute the TDSP fromq to each candidate. Figure 5.6: TDSP localization As initially showed by Dreyfus [16], the TDSP problem in FIFO networks can be solved by modifying any label-setting or label-correcting static shortest path algorithm. The asymptotic running times of these modified algorithms are same as those of their static counterparts. With our approach, we implement a time-dependent A* search (a 53 label-setting algorithm) to compute TDFP betweenq and the candidate set. The main idea with A* algorithm is to employ a heuristic function h(v) (i.e., lower-bound esti- mator between the intermediate nodev i and the targett) that directs the search towards the target and significantly reduces the number of nodes that have to be traversed. With static road networks where the length of an edge is considered as the cost, the Euclidean distance betweenv i andt is the lower-bound estimator. However, with time-dependent road networks, we need to come up with an estimator that never overestimates the travel- time between v i and t for all possible departure-times (from v i ). One simple lower- bound isd euc (v i ;t)=max(speed), i.e., the Euclidean distance betweenv i andt divided by the maximum speed among the edges in the entire network. Although this estima- tor is guaranteed to be a lower-bound betweenv i andt, it is a very loose bound, hence yields insignificant pruning. Fortunately, our approach can use Lemma 6 to obtain a much tighter lower-bound. Since the shortest path from q to p i is guaranteed to be insideLC(p i ), we can use the maximum speed inLC(p i ) to compute the lower-bound. In Chapter6, we extend our time-dependent A* algorithm to a more generic framework to answer point-to-point fastest path queries in time-dependent road networks. We outline our time-dependent A* algorithm in Algorithm 4 where essential mod- ifications (as compared to [16]) are in Lines 3, 10 and 14. As mentioned, to compute TDSP fromq to candidatep i , we only consider the nodes in the loose cell that contains q andLC(p i ) (Line 3). To compute the labels for each node, we use arrival time and the estimator to each node, i.e., cost(v i )+h LC (v i ) whereh LC (v i ) is the lower-bound estimator calculated based on the maximum speed in the loose cell (Line 10). In Lines 10 and 14,TT(v i ;v j ;t v i ) finds the time-dependent travel-time fromv i tov j as described in Section 3. 54 Algorithm 4: TDSP(q,d,t) 1: //q:source,d:target,t v :departure-time from nodev, 2: //cost(v):cost froms tov,pre(v):previous node in optimal path 3: Q set of nodes inLC(q) andLC(d) 4:8v2Qcost(v)=1,cost(q)=0 5: whileQisnotemtpy do 6: v i node in Q with smallest cost 7: removev i fromQ 8: IFv i =d THEN return path 9: foreach neighborv j ofv i 10: l(v j )=cost(v i )+h LC (v i )+TT(v i ;v j ;t v i ) 11: IFl(v j )<cost(v j ) THEN 12: cost(v j )=l(v j ) 13: pre(v j )=v i 14: t v j =d t (v i )+TT(v i ;v j ;t v i ) 15: end while 5.4 Performance Evaluation We conducted several experiments with different spatial networks and various parame- ters (see Table 1) to evaluate the performance of both TD-kNN algorithms. As our road network dataset, we use Los Angeles (LA) and San Joaquin (SJ) road networks with 304,162 and 24,123 road segments, respectively. We evaluate our proposed techniques using both syntectic and actual time-dependent travel-times gathered from real-world traffic sensor data. To generate time-dependent edge costs (travel-time) we use real- world traffic sensor dataset that we have been collecting (past 2 years) and archiving from a collection of approximately 7000 sensors located on the road network of Los Angeles. We collect speed, occupancy, volume information from these sensors and the sampling rate of the data is 1 reading/sensor/min. We spatially and temporally aggregate (average) historical sensor data based on 7 days (Monday to Sunday) of each month by assigning interpolation points for each 5 minutes. The interpolation points represent the travel-times at different times of a particular day. For example, an edge is assigned 180 55 Table 5.1: Experimental parameters for TD-kNN Parameters Default Range Number of objects 10 (K) 1,5,10,15,20(K) Number of queries 3 (K) 1,2,3,4,5 (K) Number of k 20 1,10,20,30,40,50 Object Distribution Uniform Uniform, Gaussian Query Distribution Uniform Uniform, Gaussian cost attributes to represent how traffic tends to change between 6:00AM and 9:00PM for a particular date in a particular month, e.g., Monday traffic pattern in September. We assume all roads are un-congested between 9:00PM and 6:00AM, and hence con- sider static edge weights during this interval. However, unfortunately not every edge has a sensor in road networks. In order to generate time-dependent edge weights on SJ and for the edges that does not contain any sensor in LA, we developed a traffic modeling approach that creates edge travel-time profiles [13]. Our approach uses spatial (e.g., locality, connectivity) and temporal (e.g., rush hour, weekday) characteristics to generate travel-time of network edges that does not have readily available sensor data. We run our experiments on a workstation with 2.7 GHz Pentium Duo Processor and 12GB RAM memory by monitoring each query for 100 timestamps where we we pick query location and query time uniformly at random. For each set of experiments, we only vary one parameter and fix the remaining to the default values in Table 5.1. With our experiments, we measured the tight cell hit ratio and the impact ofk, data and query object cardinality as well as the distribution. Impact of Tight Cell Hit Ratio and Direct Neighbors As we explained, if a q is located in a certain tight cell TC(p i ), our algorithm imme- diately reportsp i as the first NN. Therefore, it is essential to asses the coverage area of 56 the tight cells over the entire network. Figure 5.7 illustrates the coverage ratio of the tight cells with varying data object cardinality (ranging from 1K to 20K) on two data sets. As shown, the average tight cell coverage is about %68 of the entire network for bothLA andSJ. This implies that the first NN of a query can be answered immedi- ately with a ratio of 7=10 with no further computation. Another important parameter affecting the TD-kNN algorithm is the average number of direct neighbors for each data object. Figure 5.8 depicts the average number of neighbor cells with varying data object cardinality. As shown, the average number of neighbors is less than 9 for bothLA and SJ. Figure 5.7: Coverage ratio Figure 5.8: Average number of neighbors As mentioned in Chapter 4, we developed an incremental network expansion algo- rithm (based on [16]) to evaluatekNN queries in time-dependent networks. Below we compare our results with this baseline approach. For the rest of the experiments, since the experimental results with bothLA andSJ networks differ insignificantly, we only present the results fromLA dataset. 57 Impact ofk In this experiment, we compare the performance of both algorithms by varying the value of k. Figure 5.9 plots the average response time versus k ranging from 1 to 50 while using default settings in Table 5.1 for other parameters. The results show that TD-kNN outperforms naive approach for all values of k and scales better with the large values ofk. As illustrated, when k=1, TD-kNN generates the result set almost instantly. This is because a simplecontain() function is enough to find the first NN. As the value of k increases, the response time of TD-kNN increases at linear rate. Because, TD-kNN, rather than expanding the search blindly, benefits from localized computation. In addi- tion, we compared the average number of network node access with both algorithms. As shown in Figure 5.10, the number of nodes accessed by TD-kNN is less than the naive approach for all values ofk. Figure 5.9: Impact of k Figure 5.10: Network node access Impact of Object and Query Cardinality Next, we compare the algorithms with respect to cardinality of the data objects (P). Fig- ure 5.11 shows the impact of P on response time. The response time linearly increases with the number of data objects in both methods where TD-kNN outperforms the naive 58 approach for all cases. From P=1K to 5K, the performance gap is more significant. This is because, for lower densities where data objects are possibly distributed sparsely, naive approach requires larger portion of the network to be retrieved. Figure 5.12 shows the impact of the query cardinality (Q) ranging from 1K to 5K on response time. As shown, TD-kNN scales better with larger Q and the performance gap between the approaches increases as Q grows. Figure 5.11: Object cardinality Figure 5.12: Query cardinality Impact of Object/Query Distribution Finally, we study the impact of object, query distribution. Figure 5.13 shows the response time of both algorithms where the objects and queries follow either uniform or Gaussian distributions. TD-kNN outperforms the naive approach significantly in all cases. TD-kNN yields better performance for queries with Gaussian distribution. This is because as queries with Gaussian distribution are clustered in the network, their nearest neighbors would overlap hence allowing TD-kNN to reuse the path computations. 59 Figure 5.13: Object and Query distribution 60 Chapter 6 Online Fastest Path Computation in Time-dependent Spatial Networks The fastest path problem in time-dependent road networks was first shown by Dreyfus [16] to be polynomially solvable in FIFO networks by a trivial modification to Dijk- stra algorithm where, analogous to shortest path distances, the arrival-time to the nodes is used as the labels that form the basis of the greedy algorithm. The FIFO property, which typically holds for many networks including road networks, suggests that moving objects exit from an edge in the same order they entered the edge 1 . However, the mod- ified Dijkstra algorithm [16] is far too slow for online map applications which are usu- ally deployed on very large networks and require almost instant response times. On the other hand, there are many efficient precomputation approaches that answer fastest path queries in near real-time (e.g., [58]) in static road networks. However, it is infeasible to extend these approaches to time-dependent networks. This is because the input size (i.e., the number of fastest paths) increases drastically in time-dependent networks. Specifi- cally, since the length of as-d path changes depending on the departure-time froms, the fastest path is not unique for any pair of nodes in time-dependent networks. It has been conjectured in [5] and settled in [20] that the number of fastest paths between any pair of nodes in time-dependent road networks can be super-polynomial. Hence, an algorithm which considers the every possible path (corresponding to every possible departure-time 1 The fastest path computation is shown to be NP-hard in non-FIFO networks where waiting at nodes is not allowed [47]. Violation of the FIFO property rarely happens in real-world and hence is not the focus of this study. 61 from the source) for any pair of nodes in large time-dependent networks would suffer from exponential time and prohibitively large storage requirements. For example, the time-dependent extension of Contraction Hierarchies (CH) [1] and SHARC [7] speed- up techniques (which are proved to be very efficient for static networks) suffer from the impractical precomputation times and intolerable storage complexity. Although time-dependent fastest path computation is the most accurate and realis- tic path computation method in road networks, we observe (at the time of this thesis is being written) that most of the existing state of the art path planning applications (e.g., Google Maps, Bing Maps) do not employ time-dependency in their path compu- tations, and hence, their fastest path recommendation remains the same throughout the day regardless of the departure-time from the source (i.e., query time). While some of these applications provide alternative paths under traffic conditions (which may seem similar to time-dependent planning at first), we note that the recommended alternative paths and their corresponding travel-times still remain unique during the day, and hence no time-dependent planning. To the best of our knowledge, these applications com- pute top-k fastest paths (i.e., k alternative paths) and their corresponding travel-times with and without taking into account the traffic conditions. The travel-times which take into account the traffic conditions are simply computed by considering increased edge weights (that corresponds to traffic congestion) for each path. However, our time- dependent path planning results in different optimum paths for different departure-times from the source. For example, consider Figure 6.1(a) where Google Maps offer two alternative paths (and their travel-times under no-traffic and traffic conditions) for an origin and destination pair in Los Angeles road network. Note that the path recom- mendation and the travel-times remain the same regardless of when the user submits the query. On the other hand, Figure 6.1(b) depicts the time-dependent path recom- mendations (in different colors for different departure times) for the same origin and 62 destination pair where we computed the time-dependent fastest paths for 38 consecutive departure-times between 8AM and 5:30PM, spaced 15 minutes apart 2 . As shown, the optimal paths change frequently during the course of the day. (a) Static path planning (b) Time-dependent path planning Figure 6.1: Static vs Time-dependent path planning One may argue against the feasibility of time-dependent path planning algorithms due to a) unavailability of the time-dependent edge travel-times, or b) negligible gain of time-dependent path planning (i.e., how much time-dependent planning can improve the travel-time) over static path planning. To address the first argument, note that recent advances in sensor networks enabled instrumentation of road networks in major cities for collecting real-time traffic data, and hence it is now feasible to accurately model the time-dependent travel-times based on the vast amounts of historical data. For instance, at our research center, we maintain a very large traffic sensor dataset of Los Angeles County that we have been collecting and archiving the data for past two years (see Sec- tion 4.3 for the details of this dataset). As another example, PeMS [50] project devel- oped by UC Berkeley generates time-varying edge travel-times using historical traffic 2 The paths are computed using the algorithm presented in Section 6.2 where time-dependent edge travel-times are generated based on the two-years of historical traffic sensor data collected from Los Angeles road network. 63 sensor data throughout California. Meanwhile, we also witness that the leading naviga- tion service providers (such as Navteq [42] and TeleAtlas [66]) started releasing their time-dependent travel-time data for road networks at high temporal resolution. With regards to the second argument, several recent studies showed the importance of time- dependent path planning in road networks where real-world traffic datasets have been used for the assessment. For example, in we show in the following section that the fastest path computation that considers time-dependent edge travel-times in Los Angeles road network decreases the travel-time by on average %36 over the fastest path computation that assumes constant edge travel-times. A similar observation has been done in another study [24] under IBM’s Smart Traffic Project where the time-dependent fastest path computation in Stockholm road network can improve the travel-time accuracy signifi- cantly. Considering the availability of high-resolution time-dependent travel-time data for road networks, and the importance of time-dependency for accurate and useful path planning, the need for efficient algorithms to enable next-generation time-dependent path planning applications becomes apparent and immediate. In this section, we propose a bidirectional time-dependent fastest path algorithm (B- TDFP) based on A* search [27]. There are two main challenges to employ bidirectional A* search in time-dependent networks. First, finding an admissible heuristic function (i.e., lower-bound distance) between an intermediate v i node and the destination d is challenging as the distance betweenv i andd changes based on the departure-time from v i . Second, it is not possible to implement a backward search without knowing the arrival-time at the destination. We address the former challenge by partitioning the road network to non-overlapping partitions (an off-line operation) and precompute the intra (node-to-border) and inter (border-to-border) partition distance labels with respect to Lower-bound Graph G which is generated by substituting the edge travel-times in G with minimum possible travel-times. We use the combination of intra and inter distance 64 labels as a heuristic function in the online computation. To address the latter challenge, we run the backward search on the lower-bound graph (G) which enables us to filter-in the set of the nodes that needs to be explored by the forward search. The remainder of this section is organized as follows. In Section 3, we formally define the time-dependent fastest path problem in spatial networks. In Section 6.2, we establish the theoretical foundation of our proposed bidirectional algorithm and explain our approach. In Section 6.3, we present the results of our experiments for both approaches with a variety of spatial networks with real-world time-dependent edge weights. 6.1 Feasibility of Time-dependent Path Planning As we discussed there are handful of studies that focus on efficient computation of time- dependent fastest path. However, none of these studies investigate the practicality of time-dependent planning in real-world road networks. In this section, for the first time we assess the importance and the practical usefulness of time-dependent planning by comparing the results of time-independent fastest computation on a real-world spatial network with real time-varying edge travel-times. We focus on answering two specific questions: i) how much does time-dependent path planing reduce the travel-time as compared to static path planning, and ii) how different are the time-dependent fastest path and the static fastest path for a given source and destination. We answer these question in the following section based on our experimental evaluation with real-world datasets. Towards this end we conducted extensive experiments to evaluate the practical use- fulness of TDFP. As of our dataset, wee used California (CA) and Los Angeles (LA) road network data [42] with approximately 1,965,300 and 304,162, respectively. In our 65 lab, we maintain large-scale and high resolution (both spatially and temporally) traffic sensor (i.e., loop detector) dataset collected from the entire Los Angeles County high- ways and arterial streets. This dataset includes both inventory and real-time data (with update rate as high as every 1 minute) for 6300 traffic sensors covering approximately 3200 miles. The sampling rate of the streaming data is 1 reading/sensor/min. We have been continuously collecting and archiving the traffic sensor data past two years. We use this real-world dataset create time varying edge weights, we spatially and temporally aggregate sensor data by assigning interpolation points (for each 5 minutes) that depict the travel-times on the network segments. Based on our observation, all roads are un- congested between 9PM and 6AM, and hence we assume static edge weights between those times. In order to create time-dependent edge weights for the local streets in LA as well as the entire CA road network, we developed a traffic modeling approach (based on our observations from LA dataset) that synthetically generates edge travel-time profiles [13]. Our approach uses spatial (e.g., locality, connectivity) and temporal (e.g., rush hour, weekday) characteristics to generate travel-time of network edges that does not have readily available sensor data. In this section we report our experimental results from our fastest path queries in which we determine the source s and destination d nodes uniformly at random. We also pick our departure time randomly and uniformly distributed in time domainT . The average results are derived from 1000 randoms-d queries. With our experiment we investigate how much TDFP improves the total travel-time as compared to static fastest path (FP). We use Dreyfus’s algorithm [16] to compute time-dependent fastest path for a given s and d. To compute static fastest path (with Dijkstra’s algorithm), we use the maximum attainable speed (hence minimum travel- time) on the network edges. We conduct our experiments with the following settings. 66 (a) TDFP vs FP (b) Standard deviation Figure 6.2: TDFP and FP Comparison Givens-d path and for each 5 minutes from 6AM to 9PM, we first determine the time- dependentP =(fn 1 ;:::;n l g;t) and time-independentP =fn 0 1 ;:::;n 0 k g optimum paths as well as their corresponding travel-times to d, i.e., A t P (t) and A P (t), respectively. Next, we compute the actual time-dependent travel-timeA t P (t) of time-independent path P . Specifically, we takeP =fn 0 1 ;:::;n 0 k g and determine actual time-dependent cost of travel along P departing from n 0 1 for a given t. Figure 6.2(a) plots the improvement gained by the TDFP over its static counter part for which we measure the relative per- centage increase of FP’s travel-time over that of TDFP computed asA t P (t)=A t P (t)1. As shown,the cost of the path found by the TDFP is on average 36% better than that of FP and the difference is more significant (i.e., 68%, 43% ) during rush hours (i.e., 7-9:30AM, 4-6PM). The reason for significant difference during rush hours is that the edge weights change rapidly especially at the boundaries of the traffic peak periods and hence causing an overall increase in the cost of FP. However, TDFP avoids the conges- tion by selecting alternative segments and hence yields better travel-times. As expected, the paths found by TDFP and FP is often same before 6AM and after 9PM. Figure 6.2(b) depicts the standard deviation (in minutes) ofA t P (t)A t P (t). As illustrated, the standard deviation is also more significant during rush hours. 67 We also compared the path similarity (number of identical edges in P and P ) of TDFP and FP. Our results showed that the path found by the TDFP deviates on average 28% with maximum recorded deviation of 87% (where TDFP finds almost completely different path than that of FP) from FP. One interesting observation from this experiment is that although different departure times return different optimum paths, there exists only a limited number of different paths during a day for a givens andd. In particular, we used 180 different departure times, and on average the number of distinct optimum- time path computed by time-dependent fastest path algorithm was on average 7, and at most 12. In sum, we observe that the use of time-dependent information can significantly reduce the travel-times especially during peak hours when the faster travel-time routes needed the most. In addition, one interesting observation is that although time- dependent fastest path computation returns different optimal-paths for different depar- ture times, there are only a limited number of distinct paths (i.e., 8 on average) for a given source and destination. 6.2 Time-dependent Fastest Path Computation In this section, we explain our bidirectional time-dependent fastest path approach that we generalize bidirectional A* algorithm proposed for static spatial networks [51] to time-dependent road networks. Our proposed solution involves two phases. At the pre- computation phase, we partition the road network into non-overlapping partitions and precompute lower-bound distance labels within and across the partitions with respect to G(V;E). Successively, at the online phase, we use the precomputed distance labels as a 68 heuristic function in our bidirectional time-dependent A* search that performs simulta- neous searches from source and destination. Below we first remind the definitions that we will use in this section and then elaborate on both phases. Definition 11 Time-dependent Graph. A Time-dependent Graph is defined as G(V;E;T) whereV =fv i g is a set of nodes andE VV is a set of edges repre- senting the network segments each connecting two nodes. For every edgee(v i ;v j )2E, andv i 6=v j , there is a cost functionc v i ;v j (t), wheret is the time variable in time domain T . An edge cost functionc v i ;v j (t) specifies the travel-time fromv i tov j starting at time t. Definition 12 Time-dependent Travel Cost. Letfs = v 1 ;v 2 ;:::;v k = dg denotes a path which contains a sequence of nodes where e(v i ;v i+1 )2 E and i = 1;:::;k1. Given aG(V;E;T), a path (s d) from source s to destination d, and a departure- time at the sourcet s , the time-dependent travel costTT(s d;t s ) is the time it takes to travel the path. Since the travel-time of an edge varies depending on the arrival-time to that edge, the travel-time of a path is computed as follows: TT(s d;t s )= k1 X i=1 c v i ;v i+1 (t i ) wheret 1 =t s ,t i+1 =t i +c (v i ;v i+1 ) (t i );i=1;::;k. Definition 13 Lower-bound Graph. Given a G(V;E;T), the corresponding Lower- bound Graph G(V;E) is a graph with the same topology (i.e, nodes and edges) as graphG, where the weight of each edgec v i ;v j is fixed (not time-dependent) and is equal to the minimum possible weightc min v i ;v j where8e(v i ;v j )2E;t2T c min v i ;v j c v i ;v j (t). Definition 14 Lower-bound Travel Cost. The lower-bound travel-timeLTT(s d) of a path is less than the actual travel-time along that path and computed w.r.tG(V;E) as LTT(s d)= k1 X i=1 c min v i ;v i+1 ,i=1;::;k. 69 It is important to note that for each source and destination pair(s;d),LTT(s d) is time-independent constant value and hencet is not included in its definition. Given the definitions of TT and LTT , the following property always holds for any path in G(V;E;T): LTT(s d) TT(s d;t s ) where t s is an arbitrary departure-time froms. We will use this property in subsequent sections to establish some properties of our proposed solution. Definition 15 Time-dependent Fastest Path (TDFP). Given aG(V;E;T), s, d, and t s , the time-dependent fastest pathTDFP(s;d;t s ) is a path with the minimum travel- time among all paths froms tod for starting timet s . 6.2.1 Precomputation Phase The precomputation phase of our proposed algorithm includes two main steps in which we partition the road network into non-overlapping partitions and precompute lower- bound border-to-border, node-to-border, and border-to-node distance labels. Road Network Partitioning Real-world road networks are built on a well-defined hierarchy. For example, in United States, highways connect large regions such as states, interstate roads connect cities within a state, and multi-lane roads connect locations within a city. Almost all of the road network data providers (e.g., Navteq [42]) include road hierarchy information in their datasets. In this paper, we partition the graph to non-overlapping partitions by exploiting the predefined edge class information in road networks. Specifically, we first use higher level roads (e.g., interstate) to divide the road network into large regions. Then, we subdivide each large region using the next level roads and so on. We adopt this technique from [23] and note that our proposed algorithm is independent of the 70 partitioning method, i.e., it yields correct results with all non-overlapping partitioning methods. With our approach, we assume that the class of each edge class(e) is predefined and we denote the class of a nodeclass(v) by the lowest class number of any incoming or outgoing edge to/from v. For instance, a node at the intersection of two freeway segments and an arterial road (i.e., the entry node to the freeway) is labeled with class of the freeway rather than the class of the arterial road. The input to our hierarchical partitioning method is the road network and the level of partitioning l. For example, if we like to partition a particular road network based on the interstates, freeways, and arterial roads in sequence, we setl =2 where interstate edges represent the class 0. The road network partitions can be conceptually visualized as the areas after removal the nodes withclass(v)l fromG(E;V). Definition 16 Given a graphG(V;E), the partition ofG(V;E) is a set of subgraphs fS 1 ;S 2 ;:::;S k g whereS i =(V i ;E i ) includes node setV i whereV i \V j =; and[ k i=1 V i = V ,i6=j. Given aG(E;V) and level of partitioningl, we first assign to each node an empty set of partitions. Then, we choose a nodev i that is connected to edges other than the ones used for partitioning (i.e., a node withclass(v i )>l) and add partition number (e.g.,S 1 ) to v i ’s partition set. For instance, continuing with our example above, a node v i with class(v i ) > 2 represent a particular node that belongs a less important road segment than an arterial road. Subsequently, we expand a shortest path tree from v i to all it’s neighbor nodes reachable through the edges of the classes greater thanl, and addS 1 to their partition sets. Intuitively, we expand fromv i until we reach the roads that are used for partitioning. At this point we determine all the nodes that belong toS 1 . Then, we select another nodev j with an empty partition set by adding the next partition number (e.g., S 2 ) to v j ’s partition set and repeat the process. We terminate the process when 71 Figure 6.3: Road network partitioning all nodes are assigned to at least one partition. With this method we can easily find the border nodes for each partition, i.e., those nodes which include multiple partitions in their partition sets. Specifically, a node v, with class(v) l belongs to all partitions such that there is an edgee (withclass(e) > l) connectingv tov 0 wherev 0 2 S i and i=1;:::;k, is the border node of the partitions that it connects to. Note thatl is a tuning parameter in our partitioning method. Hence, one can arrange the size of the partitions by increasing or decreasingl. Figure 6.3 shows the partitioning of San Joaquin (California) network based on the road classes. As shown, higher level edges are depicted with different (thicker) colors. Each partition is numbered starting from the north-west corner of the road network. The border nodes between partitions S 1 and S 4 are shown in the circled area. We remark that the number of border nodes (which can be potentially large depending on the den- sity of the network) in the actual partitions have a negligible influence on the storage 72 complexity. We explain the effect of the border nodes on the storage cost in the next section. Distance Label Computation In this step, for each pair of partitions (S i ,S j ) we compute the lower-bound fastest path cost w.r.t G between each border in S i to each border node in S j . However, we only store the minimum of all border-to-border fastest path distances. As an example, con- sider Figure 6.4 where the lower-bound fastest path cost betweenb 1 andb 3 (shown with straight line) is the minimum among all border-to-border distances (i.e., b 1 -b 4 , b 2 -b 4 , b 2 -b 3 ) betweenS 1 andS 2 . In addition, for each nodev i in a partitionS i , we compute the lower-bound fastest path cost from v i to all border nodes in S i w.r.t. G and store the minimum among them. We repeat the same process from border nodes in S i to v i . For example, border nodesb 1 andb 4 in Figure 6.4 are the nearest border nodes to s and d, respectively. We will use the precomputed node-to-border, border-to-border, and border-to-node lower-bound travel-times (referred to as distance labels) to construct our heuristic function for online time-dependent A* search. We used a similar distance label precomputation technique to expedite shortest path computation between network V oronoi polygons in static road networks [34]. Figure 6.4: Lower-bound distance computation. 73 We maintain the distance labels by attaching three attributes to each node represent- ing a) the partitionS i that contains the node, b) minimum of the lower-bound distances from the node to border nodes, and c) minimum of the lower-bound distances from border nodes to the node (this is necessary for directed graphs). We keep border-to- border distance information in a hash table. Since we only store one distance value for each partition pair, the storage cost of the border-to-border distance labels is negligible. Another benefit of our proposed lower-bound computation is that the lower-bounds need to be updated when it is necessary. Specifically, we update the intra and inter distance labels only when the minimum travel-time of an edge changes, otherwise, the travel- time updates are discarded. Note that intra distance label computation is local, i.e., we only update the intra distance labels for the partitions in which the minimum travel-time of an edge changes. 6.2.2 Online B-TDFP Computation As showed in [16], the time-dependent fastest path problem (in FIFO networks) can be solved by modifying Dijkstra algorithm. We refer to modified Dijkstra algorithm as time-dependent Dijkstra (TD-Dijkstra). TD-Dijkstra visits all network nodes reachable froms in every direction until destination noded is reached. On the other hand, a time- dependent A* algorithm can significantly reduce the number of nodes that have to be traversed in TD-Dijkstra algorithm by employing a heuristic functionh(v) that directs the search towards destination. To guarantee optimal results, h(v) must be admissible and consistent (a.k.a, monotonic). The admissibility implies that h(v) must be less than or equal to the actual distance betweenv andd. With static road networks where the length of an edge is constant, Euclidian distance betweenv andd is used ash(v). However, this simple heuristic function cannot be directly applied to time-dependent road networks, because, the optimal travel-time betweenv andd changes based on the 74 departure-timet v fromv. Therefore, in time-dependent road networks, we need to use an estimator that never overestimates the travel-time betweenv andd for any possible t v . One simple lower-bound estimator is d euc (v;d)=max(speed), i.e., the Euclidean distance betweenv andd divided by the maximum speed among the edges in the entire network. Although this estimator is guaranteed to be a lower-bound, it is a very loose bound, and hence yields insignificant pruning. With our approach, we obtain a much tighter bound by utilizing the precomputed distance labels. Assuming that an on-line time-dependent fastest path query requests a path from sources in partitionS i to destinationd in partitionS j , the fastest path must pass through from one border nodeb i inS i and another border nodeb j inS j . We know that the time-dependent fastest path distance passing from b i and b j is greater than or equal to the precomputed lower-bound border-to-border (e.g.,LTT(b l ;b t )) distance for S i andS j pair. We also know that a time-dependent fastest path distance froms tob i is always greater than or equal to the precomputed lower-bound fastest path distance ofs to its nearest border nodeb s . Analogously, same is true from the border nodeb d (i.e., nearest border node) tod inS j . Thus, we can compute a lower-bound estimator ofs by h(s)=LTT(s;b s )+LTT(b l ;b t )+LTT(b d ;d). Lemma 7 Given an intermediate nodev i inS i and destination noded inS j , the esti- mator h(v i ) is admissible, i.e., a lower-bound of time-dependent fastest path distance fromv i tod passing from border nodesb i andb j inS i andS j ,respectively. Proof 7 AssumeLTT(b l ;b t ) is the minimum border-to-border distance betweenS i and S j , andb 0 i ,b 0 j are the nearest border nodes tov i andd inG, respectively. By definition of G(V;E), LTT(v i ;b 0 i ) TDFP(v i ;b i ;t v i ), LTT(b l ;b t ) TDFP(b i ;b j ;t b i ), and LTT(b 0 j ;d) TDFP(b j ;d;t b j ) Then, we haveh(v i ) = LTT(v i ;b 0 i )+LTT(b l ;b t )+ LTT(b 0 j ;d)TDFP(v i ;b i ;t v i )+TDFP(b i ;b j ;t b i )+TDFP(b j ;d;t b j ) 75 We can use ourh(v) heuristic with unidirectional time-dependent A* search in road networks. The time-dependent A* algorithm is a best-first search algorithm which scans nodes based on their time-dependent cost label (maintained in a priority queue) to source similar to [16]. The only difference to [16] is that the label within the priority queue is not determined only by the time-dependent distance to source but also by a lower-bound of the distance tod, i.e.,h(v) introduced above. To further speed-up the computation, we propose a bidirectional search that simulta- neously searches forward from the source and backwards from the destination until the search frontiers meet. However, bidirectional search is challenging in time-dependent road networks for two following reasons. First, it is essential to start the backward search from the arrival-time at the destinationt d and exactt d cannot be evaluated in advance at the query time (recall that arrival-time to destination depends on the departure-time from the source in time-dependent road networks). We address this problem by running a backward A* search that is based on the reverse lower-bound graph G (the lower- bound graph with every edge reversed). The main idea with running backward search in G is to determine the set of nodes that will be explored by the forward A* search. Sec- ond, it is not straightforward to satisfy the consistency (the second optimality condition of A* search) ofh(v) as the forward and reverse searches use different distance func- tions. Next, we explain bidirectional time-dependent A* search algorithm (Algorithm 1) and how we satisfy the consistency. Given G = (V;E;T), s and d, and departure-time t s from s, let Q f and Q b rep- resent the two priority queues that maintain the labels of nodes to be processed with forward and backward A* search, respectively. LetF represent the set of nodes scanned by the forward search andN f is the corresponding set of labeled vertices (those in its priority queue). We denote the label of a node inN f byd fv . Analogously, we define B, N b , and d fv for the backward search. Note that during the bidirectional search F 76 and B are disjoint but N f and N b may intersect. We simultaneously run the forward and backward A* searches on G(V;E;T) and G, respectively (Line 4 in Algorithm 1). We keep all the nodes visited by backward search in a set H (Line 5). When the search frontiers meet, i.e., as soon as N f and N b have a node u in common (Line 6), the cost of the time-dependent fastest path (TDFP(s;u;t s )) from s to u is deter- mined. At this point, we know thatTDFP(u;d;t u ) > LTT(u;d) for the path found by the backward search. Hence, the time-dependent cost of the paths (found so far) passing fromu is the upper-bound of the time-dependent fastest path froms tod, i.e., TDFP(s;u;t s )+TDFP(u;d;t u )TDFP(s;d;t s ). Figure 6.5: Bidirectional search If we stop the searches as soon as a nodeu is scanned by both forward and back- ward searches, we cannot guarantee finding the time-dependent fastest path fromu to d within the set of nodes in H. This is due to inconsistent potential function used in bidirectional search that relies on two independent potential functions for two inner A* algorithms. Specifically, let h f (v) (estimated distance from node v to target) and h b (v) (estimated distance from nodev to source) be the potential functions used in the forward and backward searches, respectively. With the backward search, each origi- nal edgee(i;j) considered ase(j;i) in the reverse graph whereh b used as the potential 77 function, and hence the reduced cost 3 ofe(j;i) w.r.t.h b is computed byc h b (j;i)=c(i;j)- h b (j)+h b (i) where c(i;j) is the cost in the original graph. Note that h f and h b are consistent if, for all edges (i;j), c h f (i;j) in the original graph is equal to c h b (j;i) in the reverse graph. If h f and h b are not consistent, there is no guarantee that the shortest path can be found when the search frontiers meet. For instance, consider Fig- ure 6.5 where the forward and backward searches meet at node u. As shown, if v is scanned beforeu by the forward search, thenTDFP(s;u;t s )>TDFP(s;v;t s ). Sim- ilarly if w is scanned before u by the backward search, the LTT(u;d) > LTT(w;d) and hence TDFP(u;d;t u ) > TDFP(w;d;t w ). Consequently, it is possible that TDFP(s;u;t s )+TDFP(u;d;t u )TDFP(s;v;t s )+TDFP(w;d;t w ). To address this challenge, one needs to find a) a consistent heuristic function and stop the search when the forward and backward searches meet or b) a new termination condition. In this study, we develop a new termination condition (the proof of correctness is given below) in which we continue both searches until the Q b only contains nodes whose labels exceed TDFP(s;u;t s ) + TDFP(u;d;t u ) by adding all visited nodes to H (Line 9-11). Recall that the label (denoted by d bv ) of node v in the backward search priority queue Q b is computed by the time-dependent distance from the destination to v plus the lower-bound distance from v to s, i.e., d bv = TDFP(v;d;t v ) +h(v). Hence, we stop the search when d bv > TDFP(s;u;t s ) + TDFP(u;d;t u ). As we explained, TDFP(s;u;t s ) + TDFP(u;d;t u ) is the length of the fastest path seen so far (not necessarily the actual fastest path) and is updated during the search when a new common node u 0 found with TDFP(s;u 0 ;t s ) + TDFP(u 0 ;d;t u 0) < TDFP(s;u;t s )+TDFP(u;d;t u ). Once both searches stop,H will include all the can- didate nodes that can possibly be part of the time-dependent fastest path tod. Finally, 3 A* search is equivalent to Dijkstra’s algorithm on a transformed network in which the cost of each edgec(i;j) is equal toc(i;j)-h(i)+h(j). 78 we continue the forward search considering only the nodes inH until we reachd (Line 12). Algorithm 5: B-TDFP Algorithm 1: //Input:G T , G, s:source, d:destination,t s :departure time 2: //Output: a(s;d;t s ) fastest path 3: //FS():forward search,BS():backward search,N f /N b : nodes scanned by FS()/BS(),d bv :label of the minimum element in BS queue 4: FS(G T ) andBS( G) //start searches simultaneously 5: N f FS(G T ) andN b BS( G) 6: If N f \N b 6=;thenu N f \N b 7: M =TDFP(s;u;t s )+TDFP(u;d;t u ) 8: endIf 9: Whiled bv >M 10: N b BS( G) 11: EndWhile 12: FS(N b ) 13: return(s;d;t s ) Lemma 8 Algorithm 1 finds the correct time-dependent fastest path from source to des- tination for a given departure-timet s . Proof 8 We prove Lemma 2 by contradiction. The forward search in Algorithm 1 is the same as the unidirectional A* algorithm and our heuristic functionh(v) is a lower- bound of time-dependent distance fromu tov. Therefore, the forward search is correct. Now, let P(s;(u);d;t s ) represent the path from s to d passing from u where forward and backward searches meet and ! denotes the cost of this path. As we showed ! is the upper-bound of actual time-dependent fastest path froms tod. Let be the smallest label of the backward search in priority queue Q b when both forward and backward searches stopped. Recall that we stop searches when > !. Suppose that Algorithm 1 is not correct and yields a suboptimal path, i.e., the fastest path passes from a node outside of the corridor generated by the forward and backward searches. LetP be the 79 fastest path froms tod for departure-timet s and cost of this path is. Letv be the first node onP which is going to be explored by the forward search and not explored by the backward search and h b (v) is the heuristic function for the backward search. Hence, we have h b (v)+LTT(v;d), ! < andh b (v)+LTT(v;d) LTT(s;v)+ LTT(v;d)TDFP(s;v;t s )+TDFP(v;t;t v )=, which is a contradiction. Hence, the fastest path will be found in the corridor of the nodes labeled by the backward search. 6.3 Performance Evaluation We conducted extensive experiments with different spatial networks to evaluate the per- formance of our proposed bidirectional time-dependent fastest path (B-TDFP) approach. As of our dataset, we used California (CA), Los Angeles (LA) and San Joaquin County (SJ) road network data (obtained from Navteq [42]) with approximately 1,965,300, 304,162 and 24,123 nodes, respectively. We evaluate our proposed techniques using both syntectic and actual time-dependent travel-times gathered from real-world traffic sensor data. To generate time-dependent edge costs (travel-time) we use real-world traffic sensor dataset that we have been collecting (past 2 years) and archiving from a collection of approximately 7000 sensors located on the road network of Los Angeles. We collect speed, occupancy, volume information from these sensors and the sampling rate of the data is 1 reading/sensor/min. We spatially and temporally aggregate (average) historical sensor data based on 7 days (Monday to Sunday) of each month by assigning interpolation points for each 5 minutes. The interpolation points represent the travel- times at different times of a particular day. For example, an edge is assigned 180 cost attributes to represent how traffic tends to change between 6:00AM and 9:00PM for a particular date in a particular month, e.g., Monday traffic pattern in September. We 80 assume all roads are un-congested between 9:00PM and 6:00AM, and hence consider static edge weights during this interval. However, unfortunately not every edge has a sensor in road networks. In order to generate time-dependent edge weights onSJ and CA and for the edges that does not contain any sensor in LA, we developed a traffic modeling approach that creates edge travel-time profiles [13]. Our approach uses spatial (e.g., locality, connectivity) and temporal (e.g., rush hour, weekday) characteristics to generate travel-time of network edges that does not have readily available sensor data. In this section, we report the experimental results from our fastest path queries in which we determine thes andd nodes uniformly at random. We also pick our departure- time randomly and uniformly distributed in time domain T . The average results are derived from 1000 random s-d queries. We only present the results for LA and CA, the experimental results for bothSJ andLA are very similar. We conducted our experi- ments on a server with 2.7 GHz Pentium Core Duo processor with 12GB RAM memory. Comparison with ALT In this set of experiments we compare our algorithm with time-dependent ALT (TD- ALT) approaches [8, 41] with respect to storage and response time. We run our pro- posed algorithm both unidirectionally and bidirectionally (in CA network) and compare with [8] and [41], respectively. As we mentioned, selecting good landmarks that lead to good performance is very difficult and hence several heuristics have been proposed for landmark selection. Among these heuristics, we use the best known technique; max- Cover (see [8]) with 64 landmarks. We computed travel-times between each node and the landmarks with respect to G. Under this setting, to store the precomputed distances, TD-ALT attaches to each node an array of 64 elements corresponding to the number of landmarks. Assuming that each array element takes 2 bytes of space, the additional stor- age requirement of TD-ALT is 63 Megabytes. On the other hand, with our algorithm, we 81 divide CA network to 60 partitions and store the intra and inter distance labels. The total storage requirement of our proposed solution is 8.5 Megabytes where we consume, for each node, an array of 2 elements (corresponding to from and to distances to the closest border node) plus the border-to-border distance labels. Since the experimental results for both unidirectional and bidirectional searches differ insignificantly and due to space limitations, we only present the results from unidirectional search below. As shown in Figure 6.6 the response time of our unidirectional time-dependent A* search (U-TDFP) is approximately three times better than that of TD-ALT for all times. This is because the search space of TD-ALT is severely affected by the quality of the landmarks which are selected based on a heuristic. Specifically, TD-ALT may yield very loose bounds based on the randomly selecteds andd, and hence the large search space. In addition, with each iteration, TD-ALT needs to find the best landmark (among 64 landmarks) which yields largest triangular inequality distance for better pruning; it seems that the overhead of this operation is not negligible. On the other hand, U-TDFP yields a more directional search with the help of intra and inter distance labels with no additional computation. Figure 6.6: TD-ALT Comparison 82 Performance of B-TDFP In this set of experiments, we compare the performance of our proposed approach to other existing TDFP methods w.r.t to a) preprocessing time, b) storage (byte per node), c) the average number of relaxed edges, and d) average query time. Table 1 shows the preprocessing time (Pre Processing), storage (Storage), number of scanned nodes (#Nodes), and response time (Res. Time) of time-dependent Dijkstra (TD- Dijkstra) implemented based on [16], unidirectional (U-TDFP) and bidirectional (B- TDFP) time-dependent A* search implemented using our proposed heuristic function, time-dependent Contraction Hierarchies (TD-CH) [1], and time-dependent SHARC (TD-SHARC) [7]. To implement U-TDFP and B-TDFP, we divide CA and LA net- work to 60 (which roughly correspond to counties in CA) and 25 partitions, respectively. Comparing TD-Dijkstra with our approach, we observe a very high trade-off between the query results and precomputation in both LA and CA networks. Our proposed B- TDFP performs 23 times better than TD-Dijkstra depending on the network while pre- processing and storage overhead is relatively small. As shown, the preprocessing time and storage complexity is directly proportional to network size. Comparing the time-dependent variant of SHARC (TD-SHARC) and CH (TD-CH) with our approach, we observe B-TDFP outperforms TD-SHARC and TD-CH in pre- processing and response time. We also observe that as the graph gets bigger or more edges are time-dependent, the preprocessing time of TD-SHARC increases drastically. The preprocessing of TD-SHARC takes very long for both road networks, i.e., up to 20 times more than B-TDFP. The reason for the performance gap is that TD-SHARC’s contraction routine cannot bypass the majority of the nodes in time-dependent road net- works as in the static road networks. Recall that the importance of a node can change throughout the time under consideration in time-dependent road networks. In addition, 83 Table 6.1: Experimental Results Algorithm PreProcessing Storage #Nodes Res. Time [h:m] [B/node] [ms] CA TD-Dijkstra 0:00 0 1162323 4104.11 U-TDFP 1:13 6.82 90575 310.17 B-TDFP 1:13 6.82 67172 182.06 TD-SHARC 19:41 154.10 75104 227.26 TD-CH 3:55 1018.33 70011 209.12 LA TD-Dijkstra 0:00 0 210384 2590.07 U-TDFP 0:27 3.51 11115 197.23 B-TDFP 0:27 3.51 6681 101.22 TD-SHARC 11:12 68.47 9566 168.11 TD-CH 1:58 740.88 7922 140.25 TD-SHARC is very sensitive to edge cost function changes, i.e. whenever cost func- tion of an edge changes, the preprocessing phase needs to be repeated to determine the by-pass nodes. While TD-CH tend to have better response times than TD-SHARC, the space consumption of TD-CH is significantly high (approximately 1000 bytes per node in CA network). For this reason, TD-CH is not feasible for very large road networks such as North America and Europe. We note that, to improve the response and prepro- cessing time, several variations of TD-SHARC and TD-CH algorithms are implemented in the literature. These variations trade-off between the optimality of the solution and the response time. For example, the response time of Heuristic TD-SHARC [7] is shown much better than that of original TD-SHARC algorithm. However, the path found by the Heuristic TD-SHARC is not optimal and the error rate is not bounded. As another example, the performance of TD-SHARC can be improved by combining with another technique called Arc-Flags [7]. Similar performance improvements can be applied to 84 our proposed approach. For instance, we can terminate the search when the search fron- tiers meet and report the combination of path found by the forward and backward search as the result. However, as mentioned in Section 6.2.2, we cannot guarantee the optimal solution in this setting. Moreover, based on our initial observation and implementation, we can also integrate our algorithm with Arc-Flags. However, the focus of our study is to develop a technique that yields exact solutions. Hence, for the sake of simplicity and fair comparison, we only compare the original algorithms that yields exact results and do not consider integrating different methods. Quality of Lower-bounds As discussed, the performance of time-dependent A* search depends on the lower- bound distance. In this set of experiments, we analyze the quality of our proposed lower-bound computed based on the Distance Labels explained in Section 6.2.1. We define the lower-bound quality bylg = (u;v) d(u;v) , where(u;v) andd(u;v) represent the estimated and actual travel-times between nodesu andv, respectively. Table 6.2 reports lg based on three different heuristic function, namely Naive, ALT, and DL (i.e., our heuristic function computed based on Distance Labels). Similar to other experiments, the values in Table 6.2 are obtained by selectings,d andt s uniformly at random between 6AM and 9PM. We compute the naive lower-bound estimator by deuc(u;v) max(speed) , i.e., the Euclidean distance betweenu andv is divided by the maximum speed among the edges in the entire network. We obtain the ALT lower-bounds based on G and the maxCover ([8]) technique with 64 landmarks. As shown, DL provides better heuristic function in both LA and CA. The reason is that the ALT’slg relies on the distribution of the land- marks, and hence depending on the location ofs andd it is possible to get very loose bounds. On the other hand, the lower-bounds computed based on Distance Labels are more directional. Specifically, with our approach thes andd nodes must reside in one 85 of the partitions and the (border-to-border) distance between these partitions is always considered for the lower-bound computation. Table 6.2: Lower-bound Quality Network Naive ALT DL (%) (%) (%) CA 21 42 63 LA 33 46 66 Bidirectional vs Unidirectional Search In another set of experiments, we study the impact of path length (i.e., distance froms to d) on the speed-up of bidirectional search. Hence, we measure the performance of B-TDFP and U-TDFP with respect to distance by varying the path distance (1 to 300 miles) between s and d. Figure 6.7 shows the speed-up with respect to distance. We observe that the speed-up is significantly more especially for long distance queries. The reason is that for short distances the computational overhead incurred by B-TDFP is not worthwhile as U-TDFP visits less number of nodes anyway. Figure 6.7: Speed-up Ratio Analysis 86 Chapter 7 Conclusion and Future Work 7.1 Conclusion In this thesis, we studied the problem of time-dependent k nearest neighbor and fastest path queries in spatial networks where the weight of each edge is a function of time. We formulated a generalized type of k nearest neighbor query where we, unlike the exist- ing studies, assume the edge weights of the network are time varying rather than fixed. We introduced indexing schemes that partition the network space based on object loca- tions using time-independent network distance metrics between the objects. This facili- tates the localization of the search space, hence reduces the invocation of the expensive fastest-path computation between the query point and data objects in time-dependent spatial networks. Our proposed index structures are independent of density and dis- tribution of the data objects, and effectively handle the database updates where nodes, links, and data objects are added or removed. To process the fastest path queries in time-dependent road networks, we study a bidirectional A* search based on a novel heuristic function. Since the number of fastest paths between any pair of nodes in time-dependent road networks can be theoretically super-polynomial, it is infeasible to extend existing shortest path precomputation tech- niques proposed for static road networks to time-dependent road networks. With our approach, we partition the road network to non-overlapping partitions and precom- pute the intra (node-to-border) and inter (border-to-border) partition distance labels with respect to time-independent lower-bound graph. We use the combination of intra and 87 inter distance labels as a heuristic function to efficiently prune the search space during the online computation. 7.2 Discussion of Future Work We plan to extend our work in three main directions. First, although our proposed approach is efficient for predefined static data objects, it will not scale in case of mobile data objects due to frequent node reconstruction of index structures. Therefore, we will address this fundamental challenge in kNN queries with non-predefined data objects. Second, the correctness and the performance evaluation of our proposed algorithm and it’s extensions rely on the accurate and realistic modeling of time-dependent spatial networks. Thus, we will focus on developing a framework that generates realistic and well-defined data for the time-dependent spatial networks. Third, our time-dependent fastest path computation is based on a bidirectional A* search. However, we observe that response time of this approach may degrade with relatively long paths queries (e.g., Los Angeles to Seattle). To improve the performance of our proposed algorithm we plan to study hierarchical fastest path computation in time-dependent road networks. Below we elaborate each of the three future tasks in turn. Construction of Dynamic Index Structures for Time-dependent Road Networks We plan to extend our time-dependentkNN algorithm to support arbitrarily moving data and/or query objects. This extension will enable us to support continuous monitoring of kNN queries on time-dependent road networks. As an example of continuous kNN problem, consider that the queries correspond to pedestrians, and the data objects are transportation vehicles (such as taxis, buses, and trains). As transportation vehicles and pedestrians move, each pedestrian wishes to know his/herk closest transportation 88 vehicles in terms of time-dependent travel-time. This problem and its variations have been extensively studied in Euclidean space (e.g., [37, 38]) and there exists a few studies assuming static road networks [9, 40]. To the best of our knowledge, continuouskNN queries for moving objects and queries on time-dependent road networks is an open problem. The main challenge with continuouskNN monitoring on time-dependent road networks is that the server(s) need to continuously compute and update the results of each query in real-time (or close to real-time). One approach to address this challenge is to design dynamic index structures that are updatable with minimal incremental cost as the data objects move. The basic idea is to localize the update cost of the index structure to the local tight or loose sub-networks in order to avoid global reconstruction of the indexes with each update. In addition, with lazy update techniques we can control the accuracy of the index versus its efficiency. Accurate Modeling of Time-dependent Road Networks The accurate modeling of time-dependent road networks is critical for the following rea- sons. First, the design, development and correctness of the time-dependent road network query algorithms depend on the correctness of the model. Second, due to the availability of traffic data, we envision that many researchers in industry and academia will consider developing new time-dependent query algorithms. Clearly, the performance evaluation and comparison of these proposed algorithms under various conditions are critical to the success of this research area. Therefore, to enable systematic and comprehensi- ble evaluation and comparison of the proposed algorithms, there is a need for a model that produces realistic and well-defined test data for the time-dependent road networks. Last but not least, full-scale and realistic modeling of the time-varying edge weights on road networks will enable more accurate Intelligent Transportation Systems (ITS) and 89 Advanced Traveler Information Systems (ATIS) development, and facilitate transporta- tion systems analysis to improve mobility and throughput. In [13], we conducted some preliminary work in creating traffic flow models for the Los Angeles County freeways. However, through our experiments in [DKS09a], we observed that the freeway network model is not sufficient to evaluate time-dependent query algorithms. This is because 1) number of freeways in a city, even a megacity such as Los Angeles, is too small to evaluate the scalability of our algorithms, and 2) only a small number of traffic patterns exist on freeways. Moreover, in [13], we only used spatial characteristics of road seg- ments (e.g., if a segment is in residential or business area) to cluster and model their corresponding traffic data. Clearly, we need to bring temporal clustering to the mix, for example, cluster segments based on the time of the day (e.g., rush hour). Hence, we pro- pose to create a framework for realistic and accurate modeling of time-dependent edge weights for the entire road network of Los Angeles County considering both spatial and temporal characteristics of the data. Our initial study suggests that there are a finite number of profiles for a given city (e.g., about 59 for Los Angeles). The idea is whether we can parameterize these profiles for different cities; for example, include the start and end of the morning and afternoon rush-hours as parameters to a single weight-profile to generate traffic patterns for different cities. Hierarchical Time-dependent Fastest Path Computation We plan to develop a hierarchical path planning approach (H-TDSP) that exploits the road hierarchies (e.g., freeways, arterials, alleys) inherit in real-world road networks. This solution offers to make a trade-off between the computation time and the optimality of the shortest path. The intuition here is that in real-world, given that source and des- tination are sufficiently far away, humans tend to select paths hierarchically, i.e., drive on the nearest main street which connects to freeway, drive on the freeway to a location 90 close to the destination, exit freeway and reach the destination via a main street. The hierarchical path may not necessarily be optimal but largely considered as the best path due to the fact that higher level road segments offer uninterrupted (e.g., smaller number of turns, no traffic lights) and safer travel. Because of their simplicity and popularity in real-world traveling, hierarchical route planning approaches have been widely deployed by industry for static road networks. To the best of our knowledge, hierarchical route planning solution in time-dependent road networks has not yet been studied. We plan to develop an approach where we start the path search at the first level and then carry out exploring the hierarchy in the network in ascending order. Since the number of nodes at each level shrinks rapidly, the total number of explored nodes is considerably smaller than that of the plain shortest path algorithm. The main challenge however with hierar- chical path planning in time-dependent networks is to identify the transit node through which we start searching the adjacent higher levels. This is due to the fact that the travel-time from the search node to the potential transit nodes can change significantly throughout the day, and hence unlike their static counterpart solutions, it is not possi- ble to have predetermined transit nodes. We will address this challenge by extending our tight and loose cell methods that pre-compute two sub-networks around each tran- sit node. 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Semantic caching in location-dependent query processing. In SSTD, 2001. 97 Appendix A Indexing Network Voronoi Diagrams The latest developments in wireless technologies as well as the widespread use of GPS- enabled mobile devices have led to the recent prevalence of location-based services. An important class of location based queries consists of proximity queries such as k Nearest Neighbor(kNN) query [30, 34, 54, 62, 76] and its variations, e.g., Reverse k Nearest Neighbor (RkNN) [64, 73], k Aggregate Nearest Neighbor (kANN) [72]. The proximity queries in general search for data objects that minimize a distance-based function with reference to one or more query objects. With proximity queries, potentially the distance between the query point and every object in the database (e.g., all the points-of-interest) must be computed in order to find the closest (or thek closest) object(s) to the query point. Hence, the main research focus has been on indexing the objects to avoid the exhaustive search. Earlier studies assumed Euclidean distance as the distance function and hence indexed the objects in Euclidean space (e.g., [62, 65, 75, 76]) using R-tree like index structures. With the advent of online mapping systems such as Google Maps and Mapquest and the availability of accurate nation-wide road network data, the proximity queries have been extended from Euclidean space to the road network space as natural artifact. The challenge in process- ing proximity queries on road networks is that the computation of the distance function is complex and hence the indexing techniques incorporated some sort of pre-computation of distances (in network) into their structures. One such approach is based on using network V oronoi diagrams [44]. 98 A network V oronoi diagram is a specialization of a V oronoi diagram in which the locations of objects are restricted to the network edges and the distance between objects is defined as the length of the shortest network distance (e.g., shortest path or shortest time), instead of the Euclidean distance. Any network node located in a V oronoi cell has a shortest path to its corresponding V oronoi generator that is always shorter than that to any other V oronoi generator. A large number of studies adopted network V oronoi diagrams [44] to evaluate variety of proximity queries on road networks (e.g., [34, 43, 45, 55, 71]). For example, in [45] Okabe et al. introduced six different types of network V oronoi diagrams (each corresponds to very important real-world applications) whose generators are based on points, sets of points, lines and polygons, and whose distances are given by inward/outward distances, and additively/multiplicatively weighted shortest path distances. Given a query pointq and network V oronoi diagram (NVD), the first step in answer- ing any proximity query is to locate the network V oronoi cell NVC(p i ) that contains q (the generator p i of NVC(p i ) is the nearest neighbor of q). We refer to this opera- tion ascontain(q) in the rest of the paper. Considering the large size of the underlying space (e.g., a continental size road network) with numerous data objects as well as the online nature of the queries that requires fast response-time, an index structure is neces- sary to efficiently access the portion ofNVD associated withq. Although the existing approaches successfully used network V oronoi diagrams as a pre-computation approach for partitioning the network space, they overlooked the indexing techniques that enable efficient evaluation ofcontain(q). Currently, indexing network V oronoi diagram with R-tree (referred as V oronoi R-tree or VR-tree for short) is the only known method for locating the network V oronoi cell that contains a particular point or edge of the network. VR-tree is first proposed in [34] and later used in many other approaches based on NVD (e.g., [43, 55, 71]). 99 Figure A.1: Network V oronoi Diagram In this paper, we show that VR-tree has two main problems. First, VR-tree may yield inaccurate results due to the way the V oronoi cells are formed in network space, i.e., although aNVD is generated based on the network distance metric, its V oronoi cells are created and indexed as regular polygons in Euclidean space. This inconsistency may result in a network edge belonging to a cellNVC(p i ), to be classified as a member of the cellNVC(p j ) because due to the network topology, the edge falls inside the polygon ofNVC(p j ) even though its network distance is closer to the generator ofNVC(p i ). For example, Figure A.1 depicts the network V oronoi diagram of a hypothetical road network where each line style corresponds to network V oronoi cells of the generators p 1 , p 2 and p 3 . With VR-tree the network V oronoi cells are formed by connecting the border points (i.e., {b 1 ;b 2 ;:::;b 7 }) 1 and bounded by straight line segments (i.e., bold lines in the Figure). As shown, the edges marked by false-negative are included in the V oronoi cell ofp 1 NVC(p 1 ), however the network distance from any point on the false-negative edges top 3 is shorter than that top 1 . Second, VR-tree is inefficient because of the non-disjoint partitioning of the space. Specifically, VR-tree splits the network space with hierarchically nested and largely overlapping minimum bounding rectangles (MBR) created around network V oronoi 1 We discuss the network V oronoi diagram generation in Section A.2.1 100 cells. The overhead of executing contain(q) query is prohibitively high particularly in large networks with a dense (but perhaps large) set of data objects. This is because VR-tree has to redundantly visit the parent node(s) of the overlapping MBRs (aka, back- tracking problem) in the index structure. To address both of the aforementioned drawbacks, we propose a new index- ing approach for network V oronoi diagrams based on region Quad-tree [56], termed Voronoi-Quad-tree or VQ-tree for short. VQ-tree, unlike VR-tree that approximates network V oronoi cells using regular polygons in the Euclidean space, enables exact rep- resentation of the network V oronoi cells based on quad-tree blocks in the network space, and hence always yields correct results. VQ-tree does not suffer from the backtracking problem of VR-tree. This is because VQ-tree enables disjoint decomposition of the network space and encodes each of the quad-tree blocks to indicate the identity of the network V oronoi cell of which it is a member. Thus, once the quad-tree block containing q is located, VQ-tree immediately identifies the nearest V oronoi generator based on the encoded value of that block. Our experiments with real-world datasets show that the ratio of false-negative edges is %16 on average with respect to the total number of edges in the network and VQ-tree outperforms VR-tree with 12 times improved response time (see Section 6.3). The remainder of this section is organized as follows. In Section A.1, we overview Network V oronoi diagrams and it’s properties. In Section A.2, we establish the theo- retical foundation of the proposed solution for indexing Network V oronoi diagrams for efficient and accurate processing of proximity queries in spatial networks. In Section 6.3, we present the results of our experiments with a variety of spatial networks with large number of query and data objects. 101 A.1 Preliminaries In this section, we review the principles of Euclidean and Network V oronoi diagrams. We first introduce 2-dimensional Euclidean space V oronoi diagrams and describe the properties of V oronoi diagrams. We then explain the network V oronoi diagram. We refer readers to [44] for a comprehensive discussion of Euclidean and network V oronoi diagrams. Voronoi Diagrams LetP :fp 1 ;p 2 ;::;p n g be a set ofn distinct sites (i.e., generator points) distributed in the Euclidean space. These generator points can be considered any spatial type of objects (e.g., gas station, restaurant). We define the Voronoi diagram of P as the subdivision of the space inton cells, one for each site inP , with the property that a pointq lies in the cell corresponding to a sitep i if and only ifdistance(q;p i ) < distance(q;p j ) for eachp j 2P withj6=i. Figure A.2 shows the ordinary V oronoi diagram of eight points where the distance metric is Euclidean. Figure A.2: V oronoi diagram in Euclidean space We refer to the region containing the pointp i as its V oronoi cellVC(p i ) or V oronoi polygon (seeVC(p 4 ) in the Figure). In Euclidean space,VC(p i ) is a convex polygon. 102 Each edge of VC(p i ) is a segment of the perpendicular bisector of the line segment connecting p to another point of the set P . We call each of these edges a V oronoi edge. The V oronoi cells that have common edges are called adjacent cells and their generators are called adjacent generators. The V oronoi cells are collectively exhaustive and mutually exclusive except their boundaries (i.e., V oronoi edges). We define the V oronoi cell and V oronoi diagram as follows. Definition 17 ConsiderP :fp 1 ;p 2 ;::;p n g where 2 n andp i 6= p j fori6= j,i;j2 I n = 1;:::n. The region given by VC(p i ) = pjd(p;p i )(p;p j ) where d(p;p i ) is the minimum Euclidean distance betweenp andp i is called the Voronoi Cell (VC) associated withp i . Definition 18 The set of Voronoi cells given by VD(P) =fVC(p 1 );:::;VC(p n )g is called the Voronoi Diagram (VD) generated byP . Network Voronoi Diagrams With network V oronoi diagrams (NVD), the VD described above is generalized by replacing the Euclidean space with a spatial network (e.g., road network), hence the distance with the network distance (e.g., shortest-path) between the objects. Definition 19 A road network is represented as a directional weighted graphG(N;E), where N is a set of nodes representing intersections and terminal points, and E (E NN) is a set of edges representing the network edges each connecting two nodes. Each edge e is denoted as e(n i ;n j ) where n i and n j are starting and ending nodes, respectively. In this study, we consider planar graph where edges intersect only at their endpoints. We assume that V oronoi generators are located on the network segments as the graph 103 nodes. Each edge connecting nodesp i , p j stores the network distanced N (p i ;p j ). For nodes that are not directly connected,d N (p i ;p j ) is the length of the shortest path from p i top j . Given a weighted graph G(N;E) consisting of a set of nodes N = fp 1 ;:::p n ;p n+1 ;::p o g where the firstn nodes represent the V oronoi generators and a set of edgesE =fe 1 ;:::e k g that connects the nodes, we define the set dominance region and border points as follows, Definition 20 The dominance region ofp i overp j Dom(p i ;p j ) =fpjp2 k F o=1 e o ;d N (p;p i ) d N (p;p j )g represents all points in all edges in E that are closer (or equal distance) top i thanp j . Definition 21 The border points between p i and p j b(p i ;p j ) = fpjp 2 k F o=1 e o ;d N (p;p i )=d N (p;p j )g represent all points in all edges that are equally distanced fromp i andp j . Definition 22 Based on the above definitions, the Voronoi edge set V edge of p i as V edge (p i )= F j2Innfig Dom(p i ;p j ) represents all the points in all edges inE that are closer top i than any other generator point inN. Consequently, we define network Voronoi dia- gramNVD(P) w.r.t set of pointsP asNVD(P)=fV edge (p 1 );::::;V edge (p n )g. Similar toVD described in Section A.1, the elements ofNVD are mutually exclu- sive and collectively exhaustive. A.2 Indexing Network Voronoi Cells In this section, we will first explain how to construct a network V oronoi diagram in road networks and then discuss two different index structures, namely the V oronoi R-tree and 104 V oronoi Quad-tree that efficiently identifies the subdivision of the network space that contains a particular query point or network edge. A.2.1 Network Voronoi Diagram Construction The network V oronoi diagrams can be constructed using parallel Dijkstra algorithm [17] with the V oronoi generators as multiple sources. Specifically, one can expand shortest path trees from each V oronoi generator simultaneously and stop the expansions when the shortest path trees meet. (a) Road Network (b) Network V oronoi Diagram Figure A.3: A Road network and network V oronoi diagram Figure A.3 shows an example of road network and the corresponding network V oronoi diagram. Figure A.3a depicts the original weighted graphG(N;E) which con- sists ofN =fp 1 ;p 2 ;p 3 ;p 4 ;:::p 16 g nodes wherep 1 ;p 2 , andp 3 are the V oronoi generators (i.e., data objects such as restaurants, hotels) and p 4 to p 16 are the intersections on a road network that are interconnected by a set of edges. Figure A.3b shows the NVD of the road network where each line style corresponds to the shortest path tree based on the generators {p 1 ;p 2 ,p 3 }. Each shortest path tree composes a network V oronoi cell and some edges (e.g.,e(p 4 ;p 5 )) can be partially contained in different network V oronoi 105 cells. The border pointsb 1 tob 7 are the nodes where the shortest path trees meet as a result of the parallel Dijkstra algorithm. The border points between any two generator p i andp j are equally distanced fromp i andp j . Figure A.4 shows a real network V oronoi diagram with respect to 50 data objects in Los Angeles road network. Each network node marked with a different color corresponds to a network V oronoi cell. Figure A.4: Network V oronoi diagram with P = fp 1 ;:::;p 50 g in Los Angeles road network. A.2.2 Index Generation on Network Voronoi Diagram As we discussed, to answer any proximity query with respect to a query pointq, one first needs to find the V oronoi cell that containsq. There remains a basic question concerning how to efficiently access the portion of the NVD associated with a particular query point q. This can be achieved by utilizing a spatial index structure that is generated on V oronoi cells. Below, we discuss two types of spatial index structures that can be used to index NVCs, namely, the V oronoi R-tree(VR-tree) and V oronoi Quad-tree (VQ-tree). 106 The Voronoi R-tree (VR-tree) VR-tree is first introduced in [34] where NVD is used to evaluate kNN queries in road networks. VR-tree is based on the R-tree that splits the network space with hierarchi- cally nested Minimum Bound Rectangels (MBR) generated around network V oronoi cells. Given the location of a query point q, a contain(q) query invoked on VR-tree starts from the root node and iteratively checks the MBRs (of NVCs) with respect to aq to decide whether or not to further search the child nodes. (a) NVC in VR-tree (b) False-negative edges Figure A.5: Network V oronoi cell construction in VR-tree VR-tree has two main shortcomings. First, VR-tree may yield inaccurate results for a contain(q) query. This is because VR-tree makes the simplifying assumption that although the NVD is computed based on the network distance metric, its NVCs are treated as regular polygons (by connecting border points of NVCs) and indexed using R-tree that is designed for the Euclidean distance metric. However, such approach may cause misclassification of the network edges (i.e., false-negative edges) in the network V oronoi cells, and hence inaccurate results. Specifically, a network edge belonging to a network V oronoi cell of p i NVC(p i ) may be classified as a member of another network V oronoi cell NVC(p j ). For instance, continuing with our running example 107 in Figure A.3, Figure A.5(a) shows how adjacent border points are connected to each other: if two adjacent border points are between two similar generators (e.g.,b 5 andb 7 are between p 1 and p 3 ), they can be connected with an arbitrary line. Three or more adjacent border points (e.g., b 2 , b 3 and b 5 ) can be connected to each other through an arbitrary auxiliary point (e.g., v in the figure). As a result, similar to its Euclidean counterpart, the NVCs are represented with polygons in the network space. However, to illustrate why VR-tree may fail to yield correct results, consider Figure A.5(b) where we introduce two new edges (as an extension of p 12 ) to the road network. As shown, although the new edges (marked by false-negative edges in the Figure) are included inside the V oronoi cell ofp 1 , the network distance from any point on the false-negative edges top 3 is shorter than that top 1 . Thus, with VR-tree, whenq is located on false- negative edges, acontain(q) will return incorrect V oronoi generator as the NN. With our example we only show one particular case that can happen in real-world road networks. Arguably, it is possible to increase the number of such examples under different road network topologies. Figure A.6 depicts the NVC of a particular data object in Los Angeles road network where border nodes and false-negative edges are marked by light blue and red color, respectively. One naive solution to the inaccuracy problem of VR-tree is to perform an additional refinement step. Specifically, one can maintain false-negative edges (along with their corresponding V oronoi generators) in a separate index structure and, for eachcontain(q) query, checkq against this index structure. Ifq is located in any of the false-negative edges, the corresponding V oronoi generator is returned as the nearest neighbor. Other- wise, VR-tree continues the search based on MBRs of the V oronoi cells as explained above. 108 Figure A.6: False-negative edges of a NVC in Los Angeles road network Second, VR-tree is inefficient due to non-disjoint partitioning of the space. Specif- ically, with VR-tree the hierarchy of NVCs is enforced by minimum bounding rectan- gles created around network V oronoi cells. Depending on the different topologies of the road network and the distribution of the objects on the network segments, the overlap- ping areas of MBRs of network V oronoi cells may be quite large, and hence significant computation overhead in traversing R-tree forcontain(q) query. For example, Figure A.7 illustrates the MBRs of network V oronoi cells in Figure A.4. For the sake of clar- ity, we do not include the V oronoi cells in the picture. As shown, the MBRs around network V oronoi cells result in a non-disjoint decomposition of the underlying space which means that the location occupied by a V oronoi cell may be contained in sev- eral bounding boxes. This degrades the search performance in VR-tree because of the backtracking [25] problem, i.e., the parent node(s) of the overlapping MBRs have to be accessed repeatedly in order to search the child nodes that containq. Thus, with VR-tree the amount of work often depends on the overlapping areas of MBRs. We also imple- mented VR-tree with R+ tree [60] to reduce the impact of overlapping areas. However, 109 we observe that the performance of VR+ tree is still less as compared to VQ-tree (see Section A.3). Figure A.7: Minimum bounding rectangles on network V oronoi cells The Voronoi Quad-tree (VQ-tree) The alternative to VR-tree is to index network V oronoi cells using Quad-tree [19, 56], termed V oronoi Quad-tree (VQ-tree), that enables disjoint decomposition of the under- lying space. The main observation behind VQ-tree is that each color coded area in Fig- ure A.4 is a spatially contiguous region in the network space. The regions are mutually exclusive as they do not have any overlapping areas and collectively exhaustive as every location in the network space is associated with at least one generator. Therefore, an exact approximation of the network V oronoi diagram can be obtained by using a region quad-tree [56] where the leaf nodes of the quad-tree correspond to a region in a V oronoi cell in NVD. In particular, with VQ-tree the root node represents the rectangular region enclosing the entire span of the road network (and hence NVD) under consideration. We subdivide this rectangular region into four equal quadrants where each quadrant is one of the four child nodes of the root. Subsequently, we recursively subdivide the quad- rants until each quadrant contains only one network V oronoi cell information. That is, 110 Figure A.8: VQ-tree on Los Angeles road network for each quadrant, we search for two (or more) different color-coded nodes 2 . If we find such a quadrant (meaning that the quadrant includes more than one network V oronoi cell), we subdivide that quadrant into four subquadrants. This subdivision process con- tinues recursively until all nodes in a quadrant have the same color code. Figure A.8 illustrates the quad-blocks generated on the road network in Figure A.4. We note that the leaf nodes of VQ-tree does not store any information about the network nodes. As shown in Figure A.9, the leaf nodes only store the region information (i.e., coordinates) of the quad-blocks as well as a single value (e.g, a color code or a integer number) which indicates the identity of the network V oronoi cell of which the quad-tree block is a member. We note that a leaf node in the quad-tree corresponds to a particular subdivision of a network V oronoi cell. As shown in A.8, each network V oronoi cell NVC i consists of disjoint quad-tree blocks. The disjoint decomposition of the network V oronoi diagram with VQ-tree 2 During NVD construction parallel Dijkstra algorithm can encode each node with a V oronoi cell identifier, e.g., a color 111 Figure A.9: VQ-tree addresses the two drawbacks of VR-tree. Specifically, unlike VR-tree that roughly esti- mates the network V oronoi cells with polygons in the Euclidean space, VQ-tree enables the exact representation of the network V oronoi cells using quad-tree blocks and hence always yield correct results. VQ-tree does not suffer from the backtracking problem of VR-tree, and hence fast response time forcontain(q). This is due to non-overlapping partitioning of the network V oronoi cells: once the quad-tree block containing q is located in the leaf nodes, VQ-tree immediately identifies the nearest V oronoi genera- tor based on the value (e.g, a color code) of that block. Algorithm 6 presents the outline for VQ-tree. Given a set of N nodes with their color codes and bounding box[x 1 ;x 2 ]x[y 1 ;y 2 ] that contains N as an input, Algorithm 6 creates VQ-tree by recursively splitting the quadrants until all the nodes in a quadrant have the same color code. 112 Algorithm 6: VQ-Tree Algorithm VQuadTree(N;x 1 ;x 2 ;y 1 ;y 2 )f /* Scan distinct color codes in the region cellColor[](checkRegion(N;x 1 ;x 2 ;y 1 ;y 2 ); /* If there exist more than one color-code then split ifcellColor:length> 1 then /*Initialize intermediate node node(QuadTreeNode(); /*Set Quadrants node:SE(VQuadTree(N;x 1 ;(x 2 +x 1 )=2;y 1 ;(y 1 +y 2 )=2); node:SW(VQuadTree(N;(x 2 +x1)=2;x 2 ;y 1 ;(y 1 +y 2 )=2); node:NE(VQuadTree(N;x 1 ;(x 2 +x 1 )=2;(y 1 +y 2 )=2;y 2 ); node:NW(VQuadTree(N;(x 2 +x 1 )=2;x 2 ;(y 1 +y 2 )=2;y 2 ); else /*Create leaf node QuadTreeLeafNode(cellColor[0]); end if } A.3 Performance Evaluation We conducted experiments with different spatial networks and various parameters to evaluate the performance of VQ-tree and VR-tree. We measured the ratio of false- negative edges with varying object cardinality (i.e., number of V oronoi generators) and object distribution in the road network. In addition, we compared the precomputation, index rebuilding (for dynamic environments) and response time of VQ-tree and VR-tree with respect to different network sizes and object cardinality. As of our dataset, we used California (CA), Los Angeles (LA) and San Joaquin County (SJ) road network data (obtained from Navteq [42]) with approximately 1,965,300, 304,162 and 24,123 nodes, respectively. Since the experimental results with LA and SJ networks differ insignificantly, we only present the results from theCA andLA datasets. We conducted our experiments on a workstation with 2.7 GHz Pentium Core Duo processor and 12GB 113 RAM memory. For each set of experiments, we only vary one parameter and fix the remaining to the default values in Table 1. Table A.1: Experimental parameters Parameters Default Range Object Cardinality 100 10,50,100,500,1000 Road Network LA SJ, LA, CA Object Distribution Uniform Uniform, Gaussian Ratio of False-negative Edges First, we study the ratio of false-negative edges with respect to object cardinality (i.e., number of V oronoi generators) and object distribution. To identify false-negative edges, we compare the encoded values (i.e., color code) of each node based on VR-tree and VQ-tree. Specifically, we first encode each edge to its corresponding V oronoi generator by using VR-tree polygons and then compare the encoded values to that we obtained from VQ-tree. We repeat each experiment 100 times and report the average number of incorrectly encoded (i.e., false-negative) edges with respect to total number of edges in the network. Figure A.10(a) shows the ratio of false-negative edges of both networks where the object cardinality ranging from 10 to 1000. As illustrated, the ratio of incor- rectly identified edges is %16 on average in both networks. The maximum recorded false-negative edge ratio for LA and CA road networks is %24 and %29, respectively. Figure A.10(b) illustrates the ratio of false-negative edges with different object distri- bution for both CA and LA road networks. We observe that the number of false-negative edges is less in Gaussian distribution. This is because as objects are clustered in the spa- tial network with Gaussian distribution, the corresponding shortest path trees would be less disperse and hence spatially close border points. As mentioned, with VR-tree we encode the edges based on the Euclidean polygon generated by connecting the border 114 points. The more spatially close border points provides the more accurate presentation of the NBCs and hence less false-negative edges. (a) Impact of object cardinality (b) Impact of object distribution Figure A.10: Impact of object cardinality and distribution Precomputation Time With another set of experiments, we compare the precomputation (i.e., index construc- tion) time of VR-tree and VQ-tree with varying network sizes and number of objects. In order to evaluate the impact of network size, we conducted experiments with the sub-networks of CA dataset ranging from 50K to 250K segments. We set the the node size of VR-tree to 4K bytes in all cases. Figure A.11(a) shows the precomputation time of VQ-tree and VR-tree in CA road network with varying network size. The results indicate that the precomputation time increases with the network size in both methods where VQ-tree outperforms VR-tree with all numbers of edges. This is because as the network size increases the perimeters of the polygons (and hence the number of con- nected line segments that form a polygon) grow in VR-tree. Arguably, the overhead of generating MBRs (to be used in VR-tree) around the polygons composed of numerous connected line strings is time-consuming as the coordinates (that form the lines) needs to be scanned to find the ultimate corners of the MBR. On the other hand, VQ-tree is 115 constructed based the underlying space (rather than objects in VR-tree) by recursively dividing the road network to quad-blocks each corresponding to one NVC. Figure A.11(b) illustrates the impact of object cardinality over precomputation time in LA road network (the results are similar in CA network and hence not presented). We observe that as the number of objects in the road network increases, the preprocessing time for both approaches increases. As shown, the precomputation time for VQ-tree outperforms VR-tree. The reason is that the time for hierarchically clustering polygons in VR-tree for a large datasets is relatively expensive. We also observe that the depth of VQ-tree increases with the increasing number of data objects. This is because large number of data objects yields smaller VCs and hence more splits. (a) Impact of network size (b) Impact of object cardinality Figure A.11: Impact of network size Index Reconstruction Next, we compare the index reconstruction overhead of VR-tree and VQ-tree with respect to object updates. In this set of experiments, we update the location of the randomly selected data objects and measure the index reconstruction overhead in both VR-tree and VQ-tree. Figure A.12(a) shows the index reconstruction time of both index structures with varying object update ratio (i.e., the percentage of data objects whose locations changed). We observe that VQ-tree outperforms VR-tree with respect to index 116 reconstruction. This is because the insert operations in VR-tree are expensive. When new data objects are inserted into VR-tree, besides updating leaf nodes, it is likely that updates are also required to non-leaf nodes (i.e., more than one branch of the tree maybe expanded), which leads to a large overhead during insertion. On the other hand, with VQ-tree we observe that most of the index updates take place in the leaf nodes. (a) Index reconstruction (b) Impact of object cardinality Figure A.12: Response time vs object cardinality and Index reconstruction Response Time In this experiment, we compare the performance (i.e., the response time forcontaing(q) query) of VQ-tree and VR-tree with varying object cardinality. We determine the loca- tion of the query objectq uniformly at random and report average of 100 queries. As we mentioned the original VR-tree proposed in [34] may yield inaccurate results. In order to provide correct results with VR-tree, we modify VR-tree by adding an addi- tional index structure that maintains false-negative edges. Specifically, we construct a R-tree on the false-negative network edges along with their V oronoi generators. With eachcontain(q) query, we checkq against this index structure. If we locateq on any of the false-negative edges, the corresponding data object is returned as the first NN. Oth- erwise, VR-tree continues the search based on the polygons explained in A.2.2. Figure A.12(b) plots the average response time forcontaing(q) query. The results indicate that 117 VQ-tree outperforms VR-tree with all data objects and scales better with large number of data objects. The response time of VQ-tree is approximately 12 times better than that of VR-tree with more than 200 data objects. This is because of the overlapping MBRs of network V oronoi cells. With VR-tree the amount of work often depends on the size of the overlapping areas. In particular, the overlapping areas may belong to more than one NVC and hence during the search the parent node(s) of the overlapping MBRs have to be accessed repeatedly. Moreover, with eachcontaing(q) query VR-tree performs an additional step to check ifq is located on false-negative edges; it seems that the overhead of this operation is not negligible. 118

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Creator Demiryurek, Ugur (author)

Core Title Query processing in time-dependent spatial networks

School Viterbi School of Engineering

Degree Doctor of Philosophy

Degree Program Computer Science

Publication Date 11/27/2014

Defense Date 09/16/2012

Publisher University of Southern California (original), University of Southern California. Libraries (digital)

Tag k nearest neighbor search,OAI-PMH Harvest,road networks,shortest path,spatial networks,time-dependent road networks,time-dependent shortest path

Language English

Contributor Electronically uploaded by the author (provenance)

Advisor Shahabi, Cyrus (committee chair), Giuliano, Genevieve (committee member), Nakano, Aiichiro (committee member)

Creator Email demiryur@usc.edu,demiyurek@hotmail.com

Permanent Link (DOI) https://doi.org/10.25549/usctheses-c3-120837

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