Scalable Katz Ranking Computation in Large Static and Dynamic Graphs

LIPICS - Leibniz International Proceedings in Informatics, Aug 2018

Network analysis defines a number of centrality measures to identify the most central nodes in a network. Fast computation of those measures is a major challenge in algorithmic network analysis. Aside from closeness and betweenness, Katz centrality is one of the established centrality measures. In this paper, we consider the problem of computing rankings for Katz centrality. In particular, we propose upper and lower bounds on the Katz score of a given node. While previous approaches relied on numerical approximation or heuristics to compute Katz centrality rankings, we construct an algorithm that iteratively improves those upper and lower bounds until a correct Katz ranking is obtained. We extend our algorithm to dynamic graphs while maintaining its correctness guarantees. Experiments demonstrate that our static graph algorithm outperforms both numerical approaches and heuristics with speedups between 1.5 x and 3.5 x, depending on the desired quality guarantees. Our dynamic graph algorithm improves upon the static algorithm for update batches of less than 10000 edges. We provide efficient parallel CPU and GPU implementations of our algorithms that enable near real-time Katz centrality computation for graphs with hundreds of millions of nodes in fractions of seconds.

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Scalable Katz Ranking Computation in Large Static and Dynamic Graphs

E S A Scalable Katz Ranking Computation in Large Static and Dynamic Graphs Alexander van der Grinten 0 1 2 3 4 5 0 Oded Green School of Computational Science and Engineering, Georgia Institute of Technology , USA 1 Elisabetta Bergamini Karlsruhe Institute of Technology (KIT) , Germany 2 Department of Computer Science, Humboldt-Universität zu Berlin , Germany 3 Funding Most of the work was done while AvdG was affiliated with University of Cologne, Germany, and HM with Karlsruhe Institute of Technology and University of Cologne. Additionally, EB, AvdG and HM were partially supported by grant ME 3619/3-2 within German Research 4 Henning Meyerhenke Department of Computer Science, Humboldt-Universität zu Berlin , Germany 5 David A. Bader School of Computational Science and Engineering, Georgia Institute of Technology , USA Network analysis defines a number of centrality measures to identify the most central nodes in a network. Fast computation of those measures is a major challenge in algorithmic network analysis. Aside from closeness and betweenness, Katz centrality is one of the established centrality measures. In this paper, we consider the problem of computing rankings for Katz centrality. In particular, we propose upper and lower bounds on the Katz score of a given node. While previous approaches relied on numerical approximation or heuristics to compute Katz centrality rankings, we construct an algorithm that iteratively improves those upper and lower bounds until a correct Katz ranking is obtained. We extend our algorithm to dynamic graphs while maintaining its correctness guarantees. Experiments demonstrate that our static graph algorithm outperforms both numerical approaches and heuristics with speedups between 1.5× and 3.5×, depending on the desired quality guarantees. Our dynamic graph algorithm improves upon the static algorithm for update batches of less than 10000 edges. We provide efficient parallel CPU and GPU implementations of our algorithms that enable near real-time Katz centrality computation for graphs with hundreds of millions of nodes in fractions of seconds. 2012 ACM Subject Classification Theory of computation → Dynamic graph algorithms, Theory of computation → Parallel algorithms Related Version A full version of the paper is available at [20], 03847. and phrases network analysis; Katz centrality; top-k ranking; dynamic graphs; parallel algorithms - Foundation (DFG) Priority Programme 1736. Funding was also provided by Karlsruhe House of Young Scientists via the International Collaboration Package. Funding for OG and DB was provided in part by the Defense Advanced Research Projects Agency (DARPA) under Contract Number FA8750-17-C-0086. The content of the information in this document does not necessarily reflect the position or the policy of the Government, and no official endorsement should be inferred. The U.S. Government is authorized to reproduce and distribute reprints for Government purposes notwithstanding any copyright notation here on. 1 Introduction Finding the most important nodes of a network is a major task in network analysis. To this end, numerous centrality measures have been introduced in the literature. Examples of well-known measures are betweenness (which ranks nodes according to their participation in the shortest paths of the network) and closeness (which indicates the average shortest-path distance to other nodes). A major limitation of both measures is that they are based on the assumption that information flows through the networks following shortest paths only. However, this is often not the case in practice; think, for example, of traffic on street networks: it is easy to imagine reasons why drivers might prefer to take slightly longer paths. On the other hand, it is also quite unlikely that much longer paths will be taken. Katz centrality [9] accounts for this by summing all walks starting from a node, but weighting them based on their length. More precisely, the weight of a walk of length i is αi, where α is some attenuation factor smaller than 1. Thus, naming ωi(v) the number of walks of length i starting from node v, the Katz centrality of v is defined as (1) (2) c(v) := ∞ X ωi(v) αi i=1 (I − αA) z = I~ , or equivalently: c = Pi∞=1 Ai αi I~, where A is the adjacency matrix of the graph and I~ is the vector consisting only of 1s. This can be restated as a Neumann series, resulting in the closed-form expression c = αA(I − αA)−1I~, where I is the identity matrix. Thus, Katz centrality can be computed exactly by solving the linear system followed by evaluating c = αA z. We call this approach the linear algebra formulation. In practice, the solution to Eq. (2) is numerically approximated using iterative solvers for linear systems. While these solvers yield solutions of good quality, they can take hundreds of iterations to converge [17]. Thus, in terms of running time, those algorithms can be impractical for today’s large networks, which often have millions of nodes and billions of edges. Instead, Foster et al.’s [7] algorithm estimates Katz centrality iteratively by computing partial sums of the series from Eq. (1) until a stopping criterion is reached. Although very efficient in practice, this method has no guarantee on the correctness of the ranking it finds, not even for the top nodes. Thus, the approach is ineffective for applications where only a subset of the most central nodes is needed or when accuracy is needed. As this is indeed the case in many applications, several top-k centrality algorithms have been proposed recently for closeness [2] and betweenness [13]. Recently, a top-k algorithm for Katz centrality [17] was suggested. That algorithm still relies on solving Eq. (2); however, it reduces the numerical accuracy that is required to obtain a top-k rating. Similarly, Zhan et al. [21] propose a heuristic method to exclude certain nodes from top-k rankings but do not present algorithmic improvements on the actual Katz computation. Dynamic graphs. Furthermore, many of today’s real-world networks, such as social networks and web graphs, are dynamic in nature and some of them evolve over time at a very quick pace. For such networks, it is often impractical to recompute centralities from scratch after each graph modification. Thus, several dynamic graph algorithms that efficiently update centrality have been introduced for closeness [3] and betweenness [12]. Such algorithms usually work well in practice, because they reduce the computation to the part of the graph that has actually been affected. This offers potentially large speedups compared to recomputation. For Katz centrality, dynamic algorithms have recently been proposed by Nathan et al. [15, 16]. However, those algorithms rely on heuristics and are unable to reproduce the exact Katz ranking after dynamic updates. Our contribution. We construct a vertex-centric algorithm that computes Katz centrality by iteratively improving upper and lower bounds on the centrality scores (see Section 3 for the construction of this algorithm). While the computed centralities are approximate, our algorithm guarantees the correct ranking. We extend (in Section 4) this algorithm to dynamic graphs while preserving the guarantees of the static algorithm. An extensive experimental evaluation (see Section 5) shows that (i) our new algorithm outperforms Katz algorithms that rely on numerical approximation with speedups between 1.5× and 3.5×, depending on the desired correctness guarantees, (ii) our algorithm has a speedup in the same order of magnitude over the widely-used heuristic of Foster et al. [7] while improving accuracy, (iii) our dynamic graph algorithm improves upon static recomputation of Katz rankings for batch sizes of less than 10000 edges and (iv) efficient parallel CPU and GPU implementations of our algorithm allow near real-time computation of Katz centrality in fractions of seconds even for very large graphs. In particular, our GPU implementation achieves speedups of more than 10× compared to a 20-core CPU implementation. 2 2.1 Preliminaries Notation Graphs. In the following sections, we assume that G = (V, E) is the input graph to our algorithm. Unless stated otherwise, we assume that G is directed. For the purposes of Katz centrality, undirected graphs can be modeled by replacing each undirected edge with two directed edges in reverse directions. For a node x ∈ V , we denote the out-degree of x by deg(x). The maximum out-degree of any node in G is denoted by degmax. Katz centrality. The Katz centrality of the nodes of G is given by Eq. (1). With ci(v) we denote the i-th partial sum of Eq. (1). Katz centrality is not defined for arbitrary values 1 , where σmax is the largest singular value of α. In general, Eq. (1) converges for α < σmax of the adjacency matrix A (see [9]). Katz centrality can also be defined by counting inbound walks in G [9, 18]. For this definition, ωi(x) is replaced by the number of walks of length i that end in x ∈ V . Indeed, for applications like web graphs, nodes that are the target of many links intuitively should be considered more central than nodes that only have many links themselves1. However, as inbound Katz centrality coincides with the outbound Katz centrality of the reverse graph, we will not specifically consider it in this paper. 1 This is a central idea behind the PageRank [5] metric. 2.2 Related work Most algorithms that are able to compute Katz scores with approximation guarantees are based on the linear algebra formulation and compute a numerical solution to Eq. (2). Several approximation algorithms have been developed in order to decrease the practical running times of this formulation (e.g. based on low-rank approximation [1]). Nathan et al. [17] prove a relationship between the numerical approximation quality of Eq. (2) and the resulting Katz ranking quality. While this allows computation of top-k rankings with reduced numerical approximation quality, no significant speedups can be expected if full Katz rankings are desired. Foster et al. [7] present a vertex-centric heuristic for Katz centrality: They propose to determine Katz centrality by computing the recurrence ci+1 = αA ci + I~. The computation is iterated until either a fixed point2 or a predefined number of iterations is reached. This algorithm performs well in practice; however, due to the heuristic nature of the stopping condition, the algorithm does not give any correctness guarantees. Another paper from Nathan et al. [16] discusses an algorithm for a “personalized” variant of Katz centrality. Our algorithm uses a similar iteration scheme but differs in multiple key properties of the algorithm: Instead of considering personalized Katz centrality, our algorithm computes the usual, “global” Katz centrality. While Nathan et al. give a global bound on the quality of their solution, we are able to compute per-node bounds that can guarantee the correctness of our ranking. Finally, Nathan et al.’s dynamic update procedure is a heuristic algorithm without correctness guarantee, although its ranking quality is good in practice. In contrast to that, our dynamic algorithm reproduces exactly the results of the static algorithm. 3 3.1 Iterative improvement of Katz bounds Per-node bounds for Katz centrality The idea behind our algorithm is to compute upper and lower bounds on the centrality of each node. Those bounds are iteratively improved. We stop the iteration once an application-specific stopping criterion is reached. When that happens, we say that the algorithm converges. Per-node upper and lower bounds allow us to rank nodes against each other: Let `r(x) and ur(x) denote respectively lower and upper bounds on the Katz score of node x after iteration r. An explicit construction of those bounds will be given later in this section; for now, assume that such bounds exist. Furthermore, let w and v be two nodes; without loss of generality, we assume that w and v are chosen such that `r(w) ≥ `r(v). If `r(w) > ur(v), then w appears in the Katz centrality ranking before v and we say that w and v are separated by the bounds `r and ur. In this context, it should be noted that per-node bounds do not allow us to prove that the Katz scores of two nodes are equal3. However, as the algorithm still needs to be able to rank nodes x that share the same `r(x) and ur(x) values, we need a more relaxed concept of separation. Therefore: I Definition 1. In the same setting as before, let > 0. We say that w and v are -separated, if and only if `r(w) > ur(v) − . (3) 2 Note that a true fixed point will not be reached using this method unless the graph is a DAG. 3 In theory, the linear algebra formulation is able to prove that the score of two nodes is indeed equal. However, in practice, limited floating point precision limits the usefulness of this property. Intuitively, the introduction of makes the -condition easier to fulfill than the separation condition: Indeed, separated pairs of nodes are also -separated for every > 0. In particular, -separation allows us to construct Katz rankings even in the presence of nodes that have the same Katz score: Those nodes are never separated, but they will eventually be -separated for every > 0. In order to actually construct rankings, it is sufficient to notice that once all pairs of nodes are -separated, sorting the nodes by their lower bounds `r yields a correct Katz ranking, except for pairs of nodes with a difference in Katz score of less than . Thus, using this definition, we can discuss possible stopping criteria for the algorithm: Ranking criterion. Stop once all nodes are -separated from each other. This guarantees that the ranking is correct, except for nodes with scores that are very close to each other. Top-k criterion. Stop once the top-k nodes are -separated from each other and from all other nodes. For k = n this criterion reduces to the ranking criterion. Score criterion. Stop once the difference between the upper and lower bound of each node becomes less than . This guarantees that the Katz centrality of each node is correct up to an additive constant of . Pair criterion. Stop once two given nodes u and v are -separated. First, we notice that a simple lower bound on the Katz centrality of a node v can be obtained by truncating the series in Eq. (1) after r iterations, hence, `r(v) := Pr i=1 ωi(v)αi is a lower bound on c(v). For undirected graphs, this lower bound can be improved to Pir=1 ωi(v)αi + ωr(v)αr+1, as any walk of length r can be extended to a walk of length r + 1 with the same starting point by repeating its last edge with reversed direction. I Theorem 2. Let γ = 1−dαegdmegamxax . For any r ≥ 1, v ∈ V and α < deg1max , the value r ur(v) := X αiωi(v) + αr+1ωr(v)γ i=1 is an upper bound on c(v). Proof. First, let Si(v) be the set of nodes x for which there exists a walk of length i starting in v and ending in x. Each walk of length i + 1 is the concatenation of a walk of length i ending in x ∈ Si(v) and an edge (x, y), where y is some neighbor of x. Let ωi(v, x) denote the number of walks of length i that start in v and end in x. Thus, we can write ωi+1(v) = X deg(x) ωi(v, x) ≤ x∈Si(v) X x∈Si(v) ωi+j(v) ≤ (degmax)jωi(v) . By applying induction to the previous inequality, it is easy to see that, for any j > 1, Discarding the first r terms of the sum in Eq. (1) then yields ∞ X αiωi(v) ≤ i=r+1 X∞ αr+j(degmax)jωr(v) = αrωr(v) X∞(α degmax)j j=1 j=1 = αrωr(v) 1 1 − α degmax − 1 = αr+1ωr(v)γ . For the second to last equality, we rewrite the infinite series as a geometric sum. J degmax ωi(v, x) = degmax ωi(v) . (4) The following lemma (proof in the full version of this paper [20]) shows that we can indeed iteratively improve the upper and lower bounds for each node x ∈ V : I Lemma 3. For each x ∈ V , `i(x) is non-decreasing in i and ui(x) is non-increasing in i. Theorem 2 requires us to choose α < deg1max , which is a restriction compared to the more 1 . For our experiments in the later sections of this paper, general requirement of α < σmax we set α = 1+de1gmax in order to satisfy this condition. Aside from enabling us to apply the theorem, this choice of α has some additional advantages: First, because Theorem 2 gives an upper bound on Eq. (1), Katz centrality is guaranteed to converge for this value of the α parameter4. degmax is also much easier to compute than σmax, an operation that is comparable in complexity to computing the Katz centrality itself5. Finally, α = 1+de1gmax is widely-used in existing literature [4, 7], with Foster et al. calling it the “generally-accepted default attenuation factor” [7]. It is worth remarking (proof in the full version of this paper [20]) that graphs exist for which the bound from Theorem 2 is sharp: I Lemma 4. If G is a complete graph, ui(x) = c(x) for all x ∈ V and i ∈ N. 3.2 Efficient rankings using per-node bounds In the following, we state the description of our Katz algorithm for static graphs. As hinted earlier, the algorithm estimates Katz centrality by computing ur(v) and `r(v). These upper and lower bounds are iteratively improved by incrementing r until the algorithm converges. To actually compute cr(v), we use the well-known fact that the number of walks of length i starting in node v is equal to the sum of the number of walks of length i − 1 starting in the neighbors of v, in other words: Thus, if we initialize ω1(v) to deg(v) for all v ∈ V , we can then repeatedly loop over the edges of G and compute tighter and tighter lower bounds. We focus here on the top-k convergence criterion. It is not hard to see how our techniques can be adopted to the other stopping criteria mentioned at the start of the previous subsection. To be able to efficiently detect convergence, the algorithm maintains a set of active nodes. These are the nodes for which the lower and upper bounds have not yet converged. Initially, all nodes are active. Each node is deactivated once it is -separated from the k nodes with highest lower bounds `r. It should be noted that, because of Lemma 3, deactivated nodes will stay deactivated in all future iterations. Thus, for the top-k criterion, it is sufficient to check whether (i) only k nodes remain active and (ii) the remaining active nodes are -separated from each other. This means that each iteration will require less work than its previous iteration. Algorithm 1 depicts the pseudocode of the algorithm. Computation of ωr(v) is done by evaluating the recurrence from Eq. (5). After the algorithm terminates, the -separation property guarantees that the k nodes with highest `r(v) form a top-k Katz centrality ranking (although `r(v) does not necessarily equal the true Katz score). 4 This was already noticed by Katz [9] and can alternatively be proven through linear algebra. 5 Indeed, the popular power iteration method to compute σmax for real, symmetric, positive-definite matrices has a complexity of Ω(r |E|), where r denotes a number of iterations. ωi(v) = X v→x∈E ωi−1(x) . (5) Algorithm 1 Katz centrality bound computation for static graphs. γ ← degmax /(1 − α degmax) Initialize c0(x) ← 0 ∀x ∈ V Initialize r ← 0 and ω0(x) ← 1 ∀x ∈ V Initialize set of active nodes: M ← V while not converged() do Set r ← r + 1 and ωr(x) ← 0 ∀x ∈ V for all v ∈ V do for all v → u ∈ E do ωr(v) ← ωr(v) + ωr−1(u) cr(v) ← cr−1(v) + αrωr(v) if G undirected then `r(v) ← cr(v) + αr+1ωr(v) else `r(v) ← cr(v) ur(v) ← cr(v) + αr+1ωr(v)γ function converged() partialSort(M , k, `r, decreasing) for all i ∈ {k + 1, . . . , |V |} do if ur(M [i]) − < `r(M [k]) then M ← M \ {v} if |M | > k then return false for all i ∈ {2, . . . , min(|M |, k)} do if ur(M [i]) − ≥ `r(M [i − 1]) then return false return true The converged procedure in Algorithm 1 checks whether the top-k convergence criterion is satisfied. In this procedure, M denotes the set of active nodes. The procedure first partially sorts the elements of M by decreasing lower bound `r. After that is done, the first k elements of M correspond to the top-k elements in the current ranking (which might not be correct yet). Note that it is not necessary to construct the entire ranking here; sorting just the top-k nodes is sufficient. The procedure tries to deactivate nodes that cannot be in the top-k and afterwards checks if the remaining top-k nodes are correctly ordered. These checks are performed by testing if the -separation condition from Eq. (3) is true. Complexity analysis. The sequential worst-case time complexity of Algorithm 1 is O(r |E| + r C), where r is the number of iterations and C is the complexity of the convergence checking procedure. It is easy to see that the loop over V can be parallelized, yielding a complexity of O(r |Vp | degmax +r C) on a parallel machine with p processors. The complexity of converged, the top-k ranking convergence criterion, is dominated by the O(|V | + k log k) complexity of partial sorting. Both the score and the pair criteria can be implemented in O(1). It should be noted that – for the same solution quality – our algorithm converges at least as fast as the heuristic of Foster et al. that computes a Katz ranking without correctness guarantee. Indeed, the values of cr yield exactly the values that are computed by the heuristic. However, Foster et al.’s heuristic is unable to accurately assess the quality of its current solution and might thus perform too many or too few iterations. 4 Updating Katz centrality in dynamic graphs In this section, we discuss how our Katz centrality algorithm can be extended to compute Katz centrality rankings for dynamically changing graphs. We model those graphs as an initial graph that is modified by a sequence of edge insertions and edge deletions. We do not explicitly handle node insertions and deletions as those can easily be supported by adding enough isolated nodes to the initial graph. Before processing any edge updates, we assume that our algorithm from Section 3 was first executed on the initial graph to initialize the values ωi(x) for all x ∈ V . The dynamic graph algorithm needs to recompute ωi(x) for i ∈ {1, . . . , r}, where r is the number of iterations that was reached by the static Katz algorithm on the initial graph. The main observation here is that if an edge u → v is inserted into (or deleted from) the initial graph, ωi(x) only changes for nodes x in the vicinity of u. More precisely, ωi(x) can only change if u is reachable from x in at most i − 1 steps. Algorithm 2 Dynamic Katz update procedure. E ← E \ D S ← ∅, T ← ∅ for all w → v ∈ I ∪ D do S ← S ∪ {w} T ← T ∪ {v} for all i ∈ {1, . . . , r} do updateLevel(i) for all w ∈ S do Recompute `r(w) and ur(w) from cr(w) for all w ∈ V do if ur(w) ≥ minx∈M `r(x) − then M ← M ∪ {w} . Reactivation E ← E ∪ I while not converged() do Run more iterations of static algorithm procedure updateLevel(i) for all v ∈ S ∪ T do ωi0(v) ← ωi(v) for all v ∈ S do for all w → v ∈ E do S ← S ∪ {w} ωi0(w) ← ωi0(w) − ωi−1(v) + ωi0−1(v) for all w → v ∈ I do ωi0(w) ← ωi0(w) + ωi0−1(v) for all w → v ∈ D do ωi0(w) ← ωi0(w) − ωi−1(v) for all w ∈ S do ci(w) ← ci(w) − αiωi(w) + αiωi0(w) Algorithm 2 depicts the pseudocode of our dynamic Katz algorithm. I denotes the set of edges to be inserted, while D denotes the set of edges to be deleted. We assume that I ∩E = ∅ and D ⊆ E before the algorithm. Effectively, the algorithm performs a breadth-first search (BFS) through the reverse graph of G and updates ωi for all nodes nodes that were reached in steps 1 to i. After the update procedure terminates, the new upper and lower bounds can be computed from cr as in the static algorithm. We note that ωi0(x) matches exactly the value of ωi(x) that the static Katz algorithm would compute for the modified graph. Hence, the dynamic algorithm reproduces the correct values of cr(x) and also of `r(x) and ur(x) for all x ∈ V . In case of the top-k convergence criterion, some nodes might need to be reactivated afterwards: Remember that the top-k criterion maintains a set M of active nodes. After edge updates are processed, it can happen that there are nodes x that are not -separated from all nodes in M anymore. Such nodes x need to be added to M in order to obtain a correct ranking. The ranking itself can then be updated by sorting M according to decreasing `r. It should be noted that there is another related corner case: Depending on the convergence criterion, it can happen that the algorithm is not converged anymore even after nodes have been reactivated. For example, for the top-k criterion, this is the case if the nodes in M are not -separated from each other anymore. Thus, after the dynamic update we have to perform a convergence check and potentially run additional iterations of the static algorithm until it converges again. Assuming that no further iterations of the static algorithms are necessary, the complexity of the update procedure is O(r |E| + C), where C is the complexity of convergence checking (see Section 3). In reality, however, the procedure can be expected to perform much better: Especially for the first few iterations, we expect the set S of vertices visited by the BFS to be much smaller than |V |. However, this implies that effective parallelization of the dynamic graph algorithm is more challenging than the static counterpart. We mitigate this problem, by aborting the BFS if |S| becomes large and just update the ωi scores unconditionally for all nodes. Finally, it is easy to see that the algorithm can be modified to update ω in-place instead of constructing a new ω0 matrix. For this optimization, the algorithm needs to save the value of ωi for all nodes of S before overwriting it, as this value is required for iteration i + 1. For readability, we omit this modification in the pseudocode. Implementation details. The new algorithm in this paper is hardware independent and as such we can implement it on different types of hardware with the right type of software support. Specifically, our dynamic Katz centrality requires a dynamic graph data structure. On the CPU we use NetworKit [19]; on the GPU we use Hornet6. The Hornet data structure is architecture independent, though at time of writing only a GPU implementation exists. NetworKit consists of an optimized C++ network analysis library and bindings to access this library from Python. NetworKit contains parallel shared-memory implementations of many popular graph algorithms and can handle networks with billions of edges. The Hornet [6], an efficient extension to the cuSTINGER [8] data structure, is a dynamic graph and matrix data structure designed for large scale networks and to support graphs with trillions of vertices. In contrast to cuSTINGER, Hornet better utilizes memory, supports memory reclamation, and can be updated almost ten times faster. In our experiments, we compare our new algorithm to Foster et al.’s heuristic and a conjugate gradient (CG) algorithm (without preconditioning) that solves Eq. (2). The performance of CG could be possibly improved by employing a suitable preconditioner; however, we do not expect this to change our results qualitatively. Both of these algorithms were implemented in NetworKit and share the graph data structure with our new Katz implementation. We remark that for the static case, both CG and our Katz algorithm could be implemented on top of a CSR matrix data structure to improve the data locality and speed up the implementation. Experimental setup. We evaluate our algorithms on a set of complex networks. The networks originate from diverse real-world applications and were taken from SNAP [14] and KONECT [11]. Details about the exact instances that we used can be found in the full version of this paper [20]. In order to be able to compare our algorithm to the CG algorithm, we turn the directed graphs in this test set into undirected graphs by ignoring edge directions. This ensures that the adjacency matrix is symmetric and CG is applicable. Our new algorithm itself would be able to handle directed graphs just fine. 6 Hornet can be found at, while NetworKit is available from https: // Both projects are open source, including the implementations of our new algorithm. urntime ict-Paens osc-pke. idmacs10-uk. with 10 cores per socket7 and 128 GiB RAM. Our GPU experiments are conducted on an NVIDIA P100 GPU which has 56 Streaming Multiprocessors (SMs) and 64 Streaming Processors (SPs) per SM (for a total of 3584 SPs) and has 16GB of HBM2 memory. To effectively use the GPU, the number of active threads need to be roughly 8 times larger than the number of SPs. The Hornet framework has an API that enables such parallelization (with load balancing) such that the user only needs to write a few lines of code. 5.1 Evaluation of the static Katz algorithm In a first experiment, we evaluate the running time of our static Katz algorithm. In particular, we compare it to the running time of the linear algebra formulation (i.e. the CG algorithm) and Foster et al.’s heuristic. We run CG until the residual is less than 10−15 to obtain a nearly exact Katz ranking (i.e. up to machine precision; later in this section, we compare to CG runs with larger error tolerances). For Foster’s heuristic, we use an error tolerance of 10−9, which also yields an almost exact ranking. For our own algorithm, we use the ranking convergence criterion (see Section 3) and report running times and the quality of our correctness guarantees for different values of . All algorithms in this experiment ran in single-threaded mode. fraction of separated pairs of nodes. This value represents a lower bound on the correctness of ranking. Note that pairs of nodes that have the same Katz score will never be separated. Indeed, this seems to be the case for about 0.001% of all pairs of nodes (as they are never separated, not even if is very low). Taking this into account, we can see that our algorithm already computes the correct ranking for 99% of all pairs of nodes at our algorithm outperforms the other Katz algorithms considerably. = 10−3. At this , Furthermore, Table 1 shows that the average running time of our algorithm is smaller than the running time of the Foster et al. and CG algorithms. However, the graphs in our instance set vastly differ in size and originate from different applications; thus, the average running time alone does not give good indication for performance on individual graphs. In = 10−1 is taken as baseline and the running times of all other algorithms are reported relative to this baseline. In the ≤ 10−3 setups, our Katz algorithm outperforms the CG and 7 Hyperthreading was disabled for the experiments. aNthn Top-10 1 0 0 1 0 1 1 0 2 1 0 3 1 0 4 1 0 5 aBtch size psedu eMan arnkig ufl Foster et al. algorithms on all instances. Foster et al.’s algorithm is faster than our algorithm for = 10−5 on three out of ten instances. On the depicted instances, CG is never faster than our algorithm, although it can outperform our algorithm on some small instances and for very low . Finally, in Figure 2, we present results of our Katz algorithm while using the top-k convergence criterion. We report (geometric) mean speedups relative to the full ranking criterion. The figure also includes the approach of Nathan et al. [17]. Nathan et al. conducted experiments on real-world graphs and concluded that solving Eq. (2) with an error tolerance of 10−4 in practice almost always results in the correct top-100 ranking. Thus, we run CG with that error tolerance. However, it turns out that this approach is barely faster than our full ranking algorithm. In contrast to that, our top-k algorithm yields decent speedups for k ≤ 1000. 5.2 Evaluation of the dynamic Katz algorithm In our next experiment, we evaluate the performance of our dynamic Katz algorithm to compute top-1000 rankings using = 10−4. We select b random edges from the graph, delete them in a single batch and run our dynamic update algorithm on the resulting graph. We vary the batch size b from 100 to 105 and report the running times of the dynamic graph algorithm relative to recomputation. Similar to the previous experiment, we run the algorithms in single-threaded mode. Note that while we only show results for edge deletion, edge insertion is completely symmetric in Algorithm 2. Figure 3 summarizes the results of the experiment. For batch sizes b ≤ 1000, our dynamic algorithm offers a considerable speedup over recomputation of Katz centralities. As many of the graphs in our set of instances have a small diameter, for larger batch sizes (b > 10000), almost all of the vertices of the graph need to be visited during the dynamic update procedure. Hence, the dynamic update algorithm is slower than recomputation in these cases. 5.3 Real-time Katz computation using parallel CPU and GPU implementations Our last experiment concerns the practical running time and scalability of efficient parallel CPU and GPU implementations of our algorithm. For this, we compare the running times of our shared-memory CPU implementation with different numbers of cores. Furthermore, we report results of our GPU implementation. Because of GPU memory constraints, we could PCU urntime PCU x1 oradNet-CA PCU x2 not process all of the graphs on the GPU. Hence, we provide the results of this experiment only for a subset of graphs that do fit into the memory of our GPU. The graphs in this subset have between 1.5 million and 120 million edges. We use the top-10000 convergence criterion with CPU implementation as a baseline. We report the relative running times of the 2, 4, 8 and 16 core CPU configurations, as well as the GPU configuration, to this baseline. While the parallel CPU configurations yield moderate speedups over the sequential implementation, the GPU gives a significant speedup over the 16 core CPU configuration8. Even compared to a 20 core CPU configuration (not depicted in the plots; see the full version of this paper [20]), the GPU achieves a (geometric) mean speedup of 10×. The CPU implementation achieves running times in the range of seconds; however, our GPU implementation reduces this running time to a fraction of a second. In particular, the GPU running time varies between 20 ms (for roadNet-PA) and 213 ms (for com-orkut), enabling near real-time computation of Katz centrality even for graphs with hundreds of millions of edges. 6 Conclusion In this paper, we have presented an algorithm for Katz centrality that computes upper and lower bounds on the Katz score of individual nodes. Experiments demonstrated that our algorithm outperforms both linear algebra formulations and approximation algorithms, with speedups between 150% and 350% depending on desired correctness guarantees. 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Alexander van der Grinten, Elisabetta Bergamini, Oded Green, David A. Bader, Henning Meyerhenke. Scalable Katz Ranking Computation in Large Static and Dynamic Graphs, LIPICS - Leibniz International Proceedings in Informatics, 2018, 42:1-42:14, DOI: 10.4230/LIPIcs.ESA.2018.42