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Approximability of Combinatorial Optimization Problems with Submodular Cost Functions Pushkar Tripathi Georgia Institute of Technology Based on joint work with Gagan Goel, Chinmay Karande, and Wang Lei Motivation Network Design Problem f g h Objective: Find minimum spanning tree that can be built collaboratively by these agents Functions which capture Additive Cost Function economies of scale cost(a) = 1 cost(a) = 1 cost(b) = 1 cost(b) = 1 cost(a,b) = 2 cost(a,b) = 1.5 How to mathematically model these functions? - We use Submodular Functions as a starting point. Can one design efficient approximation algorithms under Submodular Cost Functions? Assumptions over cost functions ØNormalized: ØMonotone: ØDecreasing Marginal: Submodularity + ≥ + Submodular Functions General Framework Ground set X and collection C µ 2X C: set of all tours, set of all spanning trees k agents, each specifies fi: 2X → R+ fi is submodular and monotone S Find S1, …, Sk such that: f(S) ORACLE [ Si 2 C i fi(Si) is minimized Our Results Single Agent Multiple Agents Problem Upper Bound Lower Bound Upper Bound Lower Bound Vertex Cover 2 2-ϵ 2. log n (log n) Shortest Path O(n2/3 ) (n2/3 ) O(n2/3 ) (n2/3 ) Spanning Tree n (n) n (n) Perfect Matching n (n) n (n) Lower Bounds : Information theoretic Upper Bounds : Rounding of configurational LPs, Approximating sumdodular functions and Greedy Selected Related Work Ø [Grötschel, Lovász, Schrijver 81] Minimizing non-monotone submodular function is poly-time Ø [Feige, Mirrokni, Vondrak 07] Maximizing non-monotone function is hard. 2/5-Approximation Algorithm. Ø [Calinescu, Chekuri, Pal, Vondrak 08] Maximizing monotone function subject to Matroid constraint: 1-1/e Approximation. Ø [Svitkina, Fleischer 09] Upper and lower bounds for Submodular load balancing, Sparsest Cut, Balanced Cut Ø[Iwata, Nagano 09] Bounds for Submodular Vertex Cover, Set Cover Ø[Chekuri, Ene 10] Bounds for Submodular Multiway Partition In this talk Submodular Shortest Path with single agent ü O(n 2/3) approximation algorithm ü Matching hardness of approximation In this talk Submodular Shortest Path with single agent ü O(n 2/3) approximation algorithm ü Matching hardness of approximation Submodular Shortest Path t s G=(V,E) |V| =n , |E| =m Given: Graph G, Two nodes s and t f : 2E → R+ Submodular, Monotone Goal: Find path P s.t. f(P) is minimized Attempt 1: Approximate by Additive function Let we = f({e}) Idea : e 2 OPT · OPT · 2we we e OPT t 1. Guess e* = argmax{we| e 2OPT } s 2. Pruning: Remove edges costlier than e* 3. Search: Find the shortest length s-t path in the residual graph ALG · diameter(G’).we* · diameter(G’).OPT Attempt 2: Ellipsoid Approximation John’s theorem : For every polytope P, there exists an ellipsoid contained in it that can be scaled by a factor of O(√n) to contain P P [GHIM 09]: If the convex body is a polymatroid , then there is a poly-time algorithm to compute the ellipse. Attempt 2: Ellipsoid Approximation P [GHIM 09]: If the convex body is a polymatroid , then there is a poly-time algorithm to compute the ellipse. ∀S: ∑e 2 S x(e) ≤ f(S) ∀e: x(e) ≥ 0 Polymatroid f: Submodular, monotone Approximating Submodular Functions Polynomial d4 time d5 |X| = n d1 d6 f : Monotone submodular function d3 d2 g(S) = √ e S de 2 X g(S) · f(S) · √ n g(S) Attempt 2: Ellipsoid Approximation STEP 1: [GHIM ‘09] f: 2E → R+ {de} Submodular, Monotone g(S): = √ de STEP 2: Min g(S) s.t. S 2 PATH(s,t) * Minimizing over g(S) is equivalent to minimizing just the additive part Analysis: f(P) ≤ g(P) √E P: Optimum path under g ≤ √E g(O) ≤ √E f(O) O: Optimum path under f Recap. ü Approximating by linear functions : Works for graphs with small diameter ü Approximating by ellipsoid functions : Works for sparse graphs n/2 n/2 Dense Graph with large diameter Algorithm for Shortest Path STEP 1: Pruning - Guess edge e* = argmax {we | e ϵ OPT path} - Remove edges costlier than we* Algorithm for Shortest Path STEP 1: Pruning - Guess edge e* = argmax {we | e ϵ OPT path} - Remove edges costlier than we* STEP 2 : Contraction - if ∃ v , s.t. degree(v) > n1/3, contract neighborhood of v - repeat s s t t Dense connected component Algorithm for Shortest Path STEP 1: Pruning - Let we = f({e}) - Guess edge e* = argmax {we | e ϵ OPT path} - Remove edges costlier than we* STEP 2 : Contraction - if ∃ v , s.t. degree(v) < n1/3, contract neighborhood of v - repeat STEP 3 : Ellipsoid Approximation - Calculate ellipsoidal approximation (d,g) for the residual graph Algorithm for Shortest Path STEP 1: Pruning - Let we = f({e}) - Guess edge e* = argmax {we | e ϵ OPT path} - Remove edges costlier than we* STEP 2 : Contraction - if ∃ v , s.t. degree(v) < n1/3, contract neighborhood of v - repeat STEP 3 : Ellipsoid Approximation - Calculate ellipsoidal approximation (d,g) for the residual graph STEP 4 : Search - Find shortest s-t path according to g. s t Algorithm for Shortest Path STEP 1: Pruning - Let we = f({e}) - Guess edge e* = argmax {we | e ϵ OPT path} - Remove edges costlier than we* STEP 2 : Contraction - if ∃ v , s.t. degree(v) < n1/3, contract neighborhood of v - repeat STEP 3 : Ellipsoid Approximation - Calculate ellipsoidal approximation (d,g) for the residual graph STEP 4 : Search - Find shortest s-t path according to g. STEP 5 : Reconstruction - Replace the path through each contracted vertex with one having the fewest edges. Path having fewest edges s t Analysis s R P1 P2 t Bounding the cost of P1 P1 s P2 R f(P1) ≤ √ E(R) .g(P1) ≤ √ E(R).g(OPT) Has at most ≤ √ E(R) .f(OPT) n4/3 edges ≤ n2/3 f(OPT) t Bounding the cost of P2 G1 Diam(Gi) · |Gi|/n1/3 s G2 f(P2) ≤ (dia(G1) +.. +dia(Gk) ) we* ≤ (|G1| / n1/3 + …. ) we* G3 ≤ (n / n1/3) we* t ≤ n2/3 f(OPT) In this talk Submodular Shortest Path with single agent ü O(n 2/3) approximation algorithm ü Matching hardness of approximation Information Theoretic Lower Bound S1 f(S1) S2 f f(S2) S3 f(S2) Polynomial number of queries to the oracle Algorithm is allowed unbounded amount of time to process the results of the queries Not contingent on P vs NP General Technique Cost functions f , g satisfying OPT( f ) >> OPT( g ) f (S) = g(S) for ‘most’ sets S A – any randomized algorithm f(Q ) = g( Q ) with high probability for every query Q made by A. Probability over random bits in A. Yao’s Lemma f(Q) = g(Q) with high probability for every query Q made by randomized algorithm A. f and a distribution D from which we choose g, such that for an arbitrary query Q , f(Q) = g(Q) with high probability Non-combinatorial Setting X : Ground set f(S) = min{ |S|, ® } D : R µ X, |R| = ® gR(S) = min{| S Å Rc| + min( S Å R, ¯ ) } Optimal Query Claim : Optimal query has size ® Case 1 : |Q| < ® Probability can only increase if we increase |Q| Case 2 : |Q| > ® Probability can only increase if we decrease |Q| Optimal query size to distinguish f and gR is ® Distinguishing f and gR Chernoff Bounds f and g are hard to distinguish ¯ = (1+ ±) E[|Q Å R|] Hardness of learning submodular functions Set ® = n1/2log n Optimal query size = ® = n1/2log n |R| = ® = n1/2log n E[ Q Å R] = log2n Super logarithmic ¯ = (1+±) E[ Q Å R] = (1+±)log2n f and g are indistinguishable f(R) = min{ |R|, ® } = |R| = ® = n1/2 log n gR(R) = min{| R Å Rc| + min( R Å R, ¯ ) } = ¯ = log2 n Corollary : Hard to learn a submodular function to a factor better than n1/2/log n in polynomial value queries. Difficulty in Combinatorial Setting Randomly chosen set may not be a feasible solution in the combinatorial setting. Eg. Randomly chosen set of edges rarely yield a s-t path. Solution : 1. Do not choose R randomly from the entire domain X. 2. Use a subset of R as a proxy for the solution. Base Graph G s …... t n1/3 vertices n2/3 levels Functions f and g Y s … … t …. …. B f(S) = f( S Å B ) & g(S) = g( S Å B ) Functions f and g Y s … … t …. …. B f(S) = min( |S Å B|, α) Functions f and g Solution : 1. Do not choose R randomly from the entire domain X. 2. Use a subset of R as a proxy for the solution. Y s … … t … … . . Uniform random B subset of B of size ® gR(S) = min{| S Å R Å B| + min( S Å R Å B, ¯ )} Functions f and g Solution : 1. Do not choose R randomly from the entire domain X. 2. Use a subset of R as a proxy for the solution. Y R = n2/3 log2 n s … … t …. …. B gR(S) = min{| S Å R Å B| + min( S Å R Å B, ¯ ) Setting the constants Set ® = n2/3 log2 n Optimal Query size = ® = n2/3 log2 n ¯ = log2 n f and g are indistinguishable f(OPT) = min{ |R|, ® } = |R| = ® = O(n2/3 log2 n) gR(OPT) = min{| R Å Rc| + min( R Å R, ¯ ) } = ¯ = log2 n Theorem : Submodular Shortest Path problem is hard to approximate to a factor better than O(n2/3) Summary Single Agent Multiple Agents Problem Upper Bound Lower Bound Upper Bound Lower Bound Vertex Cover 2 2-ϵ 2. log n (log n) Shortest Path O(n2/3 ) (n2/3 ) O(n2/3 ) (n2/3 ) Spanning Tree n (n) n (n) Perfect Matching n (n) n (n) n: # of vertices in graph G What’s the right model to study economies of scale? Newer Models Discount Models f E R g E R E hR Payment Sub modular functions Cost f(a) + f(b) + f(c) …. Task: Minimize sum of payments Approximability under Discounted Costs[GTW 09] Problem Lower Bound Upper bound Edge Cover O(log n) O(log n) Spanning Tree O(log n) O(log n) Shortest Path O(poly log n) n Minimum Perfect Matching O(poly log n) n Shortest Path : O(logc n) hardness S U s t Agents - Cost of every edge is 1 1 1 Set Cover Instance Claim : Set cover of size |S| ↔ Shortest path of length |S| Hardness Gap Amplification s s t Original Instance • Replace each edge by a copy of the original graph. • Edges of the same color get the same copy. • Edges of different colors gets t copies with new colors(agents) Harder Instance Claim : The new instance has a solution of cost α2 iff the original instance has a solution of cost α. For any fixed constant c iterate this construction c times to further amplify the lower bound to O(logcn). Q.E. Why is it so hard to distinguish f and g ? Observation: fR(S ) is at most g(S ) for any set S. Case 1: ‘Small’ size queries - |Q | ≤ n This probability can only increase if we increase |Q | Case 2: ‘Large’ size queries - |Q | ≥ n This probability can only increase if we decrease |Q | Combinatorial Optimization ØC - Ground set Øf - Valuation function over subsets of C ØX - Collection of some subsets C having a special property ØTask - Find the set in X that has minimum cost under a given valuation function. General Technique cont. S1 f(S1) = g(S1) S2 f f(S2) = g(S2) S3 f(S3) = g(S3) A cannot distinguish Output is at least between f and g OPT( g ) α ≥ OPT( g ) OPT( f ) Plan Fix a cost function f Fix a distribution D of functions such that for every g in D OPT(f ) >> OPT (g) For an arbitrary query Q , f(Q) = g(Q) with high probability Optimal size queries Queries of size n- |Q | = n

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