Complexity Theory of PolynomialTime Problems


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1 Complexity Theory of PolynomialTime Problems Lecture 1: Introduction, Easy Examples Karl Bringmann and Sebastian Krinninger
2 Audience no formal requirements, but: NPhardness, satisfiability problem, how to multiply two matrices, dynamic programming, allpairsshortestpath problem, Dijkstra s algorithm if you (vaguely) remember at least half of these things, then you should be able to follow the course
3 NPhardness suppose we want to solve some problem fastest algorithm that we can come up with: O(2 $ ) should we search further for an efficient algorithm? prove NPhardness! an efficient algorithm would show P=NP we may assume that no efficient algorithm exists relax the problem: approximation algorithms, fixed parameter tractability, average case, heuristics
4 What about polynomial time? suppose we want to solve another problem we come up with an O(n ' ) algorithm polynomial time = efficient big data: input is too large for O(n ' ) algorithm should we search for faster algorithms? P vs NP is too coarse we need finegrained complexity to study limits of big data
5 Conditional Lower Bounds which barriers prevent subquadratic algorithms? consider a wellstudied problem P: bestknown running time O(n * ) conjecture that it has no O(n *+, ) algorithm for any ε > 0 relate another problem Q to P via a reduction conditional lower bound
6 Hard problems SAT: OV: APSP: 3SUM: given a formula in conj. normal form on n variables is it satisfiable? conjecture: no O(2 :+, $ ) algorithm (SETH) given n vectors in {0,1} 8 (for small d) are any two orthogonal? conjecture: no O(n '+, ) algorithm given a weighted graph with n vertices compute the distance between any pair of vertices conjecture: no O(n 3+, ) algorithm given n integers do any three sum to 0? conjecture: no O(n '+, ) algorithm
7 Relations = Reductions transfer hardness of one problem to another one by reductions problem P instance I size n reduction r(n) time problem Q instance J size s(n) I is a yes instance () J is a yes instance t(n) algorithm for Q implies a r n + t(s n ) algorithm for P if P has no r n + t(s n ) algorithm then Q has no t(n) algorithm
8 Relations = Reductions transfer hardness of one problem to another one by reductions problem P instance I size n reduction total time r(n) problem Q instance I 1 size s : (n) instance I k size s E (n) t(n) algorithm for Q implies a r n + E DF: t(s D n ) algorithm for P
9 Showcase Results longest common subseq. edit distance, longest palindromic subsequence, Fréchet distance... O(n ' ) SETHhard n '+, [B.,Künnemann 15, Abboud,Backurs,VWilliams 15] given two strings x, y of length n, compute the longest string z that is a subsequence of both x and y a b b c a d a c d a a b d
10 Showcase Results longest common subseq. edit distance, longest palindromic subsequence, Fréchet distance... O(n ' ) SETHhard n '+, [B.,Künnemann 15, Abboud,Backurs,VWilliams 15] we can stop searching for faster algorithms! in this sense, conditional lower bounds replace NPhardness O(n '+, ) algorithms are unlikely to exist improvements are at least as hard as a breakthrough for SAT
11 Showcase Results longest common subseq. edit distance, longest palindromic subsequence, Fréchet distance... O(n ' ) SETHhard n '+, [B.,Künnemann 15, Abboud,Backurs,VWilliams 15] bitonic TSP O(n ' ) O n log N n longest increasing subsequence, matrix chain multiplication... [de Berg,Buchin,Jansen,Woeginger 16] given n points in the plane, compute a minimum tour connecting all points among all tours consisting of two xmonotone parts
12 Showcase Results longest common subseq. edit distance, longest palindromic subsequence, Fréchet distance... O(n ' ) SETHhard n '+, [B.,Künnemann 15, Abboud,Backurs,VWilliams 15] bitonic TSP O(n ' ) O n log N n longest increasing subsequence, matrix chain multiplication... [de Berg,Buchin,Jansen,Woeginger 16] maximum submatrix O(n 3 ) APSP hard n 3+, minimum weight triangle, graph centrality measures... [Backurs,Dikkala,Tzamos 16] given matrix A over Z, choose a submatrix (consisting of consecutive rows and columns of A) maximizing the sum of all entries
13 Showcase Results longest common subseq. edit distance, longest palindromic subsequence, Fréchet distance... O(n ' ) SETHhard n '+, [B.,Künnemann 15, Abboud,Backurs,VWilliams 15] bitonic TSP O(n ' ) O n log N n longest increasing subsequence, matrix chain multiplication... [de Berg,Buchin,Jansen,Woeginger 16] maximum submatrix O(n 3 ) APSP hard n 3+, minimum weight triangle, graph centrality measures... [Backurs,Dikkala,Tzamos 16] colinearity motion planning, polygon containment... O(n ' ) 3SUM hard n '+, [Gajentaan,Overmars 95] given n points in the plane, are any three of them on a line?
14 Showcase Results longest common subseq. edit distance, longest palindromic subsequence, Fréchet distance... O(n ' ) SETHhard n '+, [B.,Künnemann 15, Abboud,Backurs,VWilliams 15] bitonic TSP O(n ' ) O n log N n longest increasing subsequence, matrix chain multiplication... [de Berg,Buchin,Jansen,Woeginger 16] maximum submatrix O(n 3 ) APSP hard n 3+, minimum weight triangle, graph centrality measures... [Backurs,Dikkala,Tzamos 16] colinearity motion planning, polygon containment... O(n ' ) 3SUM hard n '+, [Gajentaan,Overmars 95] Open: optimal binary search tree O(n ' ) knapsack O(nW) many more...
15 Complexity Inside P SAT 2 n 3SUM n 2 OV n 2 Colinearity n 2 3SUMhard APSP n 3 APSP equivalent Radius n 3 LCS n 2 SETHhard Negative Triangle n 3 EDIT n 2 Fréchet n 2 classification of polynomial time problems diameter n 2 problemcentric view on complexity theory
16 Machine Model complexity theory is (to some extent) independent of the machine model but only up to polynomial factors we have to fix a machine model! Turing Machine: any (singletape) Turing machine takes time Ω(n ' ) for recognizing palindromes abbcaacbba this does not apply to real computers!
17 Machine Model Random Access Machine (RAM): CPU memory of O(1) cells can perform all reasonable operations on two cells in O(1) time (arithmetic + logical operations ) can read/write any cell in O(1) time storage: cell / word consisting of Θ(log n) bits can recognize palindromes in time O(n) the details do not matter!
18 More Discussion What about unconditional lower bounds? no tools available beyond Ω(n log n) What if the hypotheses are wrong? NPhardness was in the same situation 40 years ago relations between problems will stay suggests ways to attack a problem + which problems to attack
19 Conditional Lower Bounds allow to classify polynomial time problems are an analogue of NPhardness yield good reasons to stop searching for faster algorithms should belong to the basic toolbox of theoretical computer scientists allow to search for new algorithms with better focus improve SAT before longest common subsequence nonmatching lower bounds suggest better algorithms motivate new algorithms relax the problem and study approximation algorithms, parameterized running time,
20 Content of the Course we study core problems SAT, OV, APSP, 3SUM, and others conditional lower bounds: from each of these hypotheses algorithms: learn fastest known algorithms for core problems finegrained complexity is a young field of research we will see many open problems & possibilities for BA/MAtheses
21 II. An Example for OVhardness
22 Orthogonal Vectors Hypothesis Input: Task: Sets A, B 0,1 8 of size n Decide whether there are a A, b B such that a b 8 [ a D \ b D = 0 DF: for all 1 i d: a D = 0 or b D = 0 A = { 1,1,1, 1,1,0, 1,0,1, 0,0,1 } B = { 0,1,0, 0,1,1, 1,0,1, 1,1,1 } trivial O(n ' d) algorithm best known algorithm: O(n '+:/e(fgh *) ) where d = c log n [Lecture 03] OVHypothesis: no O(n '+, poly(d)) algorithm for any ε > 0 OV has no O(n '+, ) algorithm, even if d = polylog n
23 Graph Diameter Problem Input: Task: An unweighted graph G = (V, E) Compute the largest distance between any pair of vertices = max s,t u d v(u, v) diameter 2 Easy algorithm: Singlesourceshortestpaths: Dijkstra s algorithm: O m + n log n Allpairsshortestpaths: Dijkstra from every node: O n m + n log n O(n m log n) from this information we can compute the diameter in time O(n ' )
24 OVHardness Result OV sets A, B 0,1 8 of size n reduction time O(dn) diameter graph G O(n) vertices O(dn) edges O(n '+, poly(d)) algorithm O((nm) :+, ) algorithm Thm: Diameter has no O((nm) :+, ) algorithm unless the OVHypothesis fails. [Roditty,VWilliams 13]
25 Proof OV sets A, B 0,1 8 of size n reduction time O(dn) diameter graph G O(n) vertices O(dn) edges Proof: can assume: every vector has at least one 1 put edge if a D = d i put edge if b D = 1 d a, b = 2 a, b not orthogonal A a b B diameter = 3 there exists an orthogonal pair
26 Proof OV sets A, B 0,1 8 of size n reduction time O(dn) diameter graph G O(n) vertices O(dn) edges Remark: Even deciding whether the diameter is 2 or 3 has no O((nm) :+, ) algorithm unless OVH fails. There is no ( 3 ' ε)approximation for the diameter in time O((nm) :+, ) unless OVH fails.
27 III. Another Example for OVhardness
28 NFA Acceptance Problem nondeterministic finite automaton G accepts input string s if there is a walk in G from starting state to some accepting state, labelled with s 0,1 0,1 start 1 1 string: dynamic programming algorithm in time O( s G ): T i set of states reachable via walks labelled with s[1.. i] T 0 {starting state} T i v u T i 1 and transition u v labelled s i }
29 OVHardness Result OV sets A, B 0,1 8 of size n reduction time O(dn) NFA acceptance grammar G, string s G = O(dn) s = O(dn) O(n '+, poly(d)) algorithm O(( s G ) :+, ) algorithm Thm: NFA acceptance has no O(( s G ) :+, ) algorithm unless OVH fails. [Impagliazzo]
30 Proof OV sets A, B 0,1 8 of size n reduction time O(dn) NFA acceptance grammar G, string s G = O(dn) s = O(dn) Proof: fix some a A: in string s: 0011 fix some b B: in NFA G: = a : a ' a 8 if b D = 1 if b D = 0
31 Proof OV sets A, B 0,1 8 of size n reduction time O(dn) NFA acceptance grammar G, string s G = O(dn) s = O(dn) Proof: fix some a A: in string s: 0011 fix some b B: in NFA G: = a : a ' a 8 if b D = 1 if b D = 0
32 Proof fix some a A: in string s: 0011 fix some b B: in NFA G: string s = $1100$0110$ $0011$ (for all a A) equivalent to OV instance size s = G = O(dn) NFA G: start 0,1, $ $ (for all b B) 0 1 $ 0,1, $ $ $ 1 1
33 VI. Four Russians
34 Method of Four Russians Arlazarov, Dinic, Kronrod, and Faradzev 1970 Algorithm for Boolean matrix multiplication Not all of them were Russian Better name: Four Soviets??
35 Boolean Matrix Multiplication Input: Boolean (0/1) matrices A and B Output: A B where + is OR and is AND A B C n = n
36 Naïve Algorithm A i, 1 B 1,j A i, k B k, j B Running time: time O n per inner product #inner products: n ' O n 3 total time A C C i, j = A i, k B k, j : E $ Is there a k such that A i, k = 1 and B k, j = 1?
37 Main Idea Divide A into blocks of size t t We use t = 0.5 log n t Number of blocks: n ' = n' t log ' n t 1. Preprocess blocks and construct lookup table 2. Speed up naïve algorithm using lookup table
38 Preprocessing a Block X 1. For every tdimensional vector v: precompute product X v v # inner products: t X O t ' time 2. Store results in lookup table: retrieve X v in time O(t) # tdimensional vectors: 2 = 2. fgh$ = n. Total time: $ ' 2 t ' = n ' 2 = n '. = O $ fgh $ #blocks #vectors inner product
39 Multiplying with Lookup Table For each of output: 1 O(t) $ times 2 O(t) $ times O n n t n t t = O n3 t = O n3 log n
40 Summary Key property: small number of possible cell entries Main idea: speedup from lookup tables after preprocessing Discussion: In preprocessing: need to prepare for all tdimensional input vectors Counterintuitive at first Only some of the tdimensional vectors might really appear in B # different tdimensional vectors in general: 2 = n. # tdimensional vectors per block in matrix multiplication: n Reusability outweighs preprocessing cost Four Russians in some sense a charging trick: Charge running time to the tdim. vectors instead of the columns
41 Beyond Four Russians Very general technique: Matrix problems, dynamic programming problems (e.g. LCS) Shaving off logarithmic factors popular for certain problems Race to fastest algorithm Example: Allpairs shortest paths (APSP) FloydWarshall: O(n 3 ) State of the art: O $ ' ( ª«) The Polynomial Method 2 fgh$ grows faster than log * n for any constant c Ideas from circuit complexity We will teach it in May
42 Transitive Closure Problem Directed graph G, n nodes TC matrix TC i, j = 1 iff i can reach j Thm: BMM in time O(T n ) TC in time O(T n ) Naïve Algorithm: O n 3 by breadthfirst search from each node
43 Reduction: From BMM (T (n)) to TC (T(n)) A B n 2layered graph: edge from i to j iff A i, j = 1 i j TC = A B n T ± n = T 3n
44 Reduction: From TC (T (n)) to BMM (T(n)) A: I: adjacency matrix of G A i, j = 1 if and only if G has edge (i, j) identity matrix Fact: (A I) E i, j = 1 if and only if path from i to j of length at most k Fact: TC = (A I) $ Repeated squaring: M ' = M M M N = M ' M ' M ³ = M N M N Lem: TC can be computed using log n Boolean matrix multiplications
45 Drawback of Reduction Lem: TC can be computed with log n Boolean matrix multiplications T ± n = T n log n = n3 log n log n = n3 Logfactor improvement of Four Russians is gone! Goal: Better reduction with T ± n = O(T n )
46 Reducing Problem to DAG I 1 III IV II 4 1. Compute strongly connected components (SCCs) i and j in same component iff i can reach j and j can reach i Suffices to solve problem on graph of SCCs Graph of SCCs is directed acyclic graph (DAG) 2. Compute topological order on DAG: edge i, j in DAG i j O(n ' )
47 Recursive Transitive Closure Input: DAG G with n nodes in topological order X: First n 2 edges in topological order Y: Last n RECURSE 2 edges in topological order M: Adjacency matrix of edges between X and Y X ± X Y Y ± TC G = TC X M TC(Y) Master Theorem* T ± n = 2T ± $ ' + O T n = O(T n ) * 2T ± ct (n) for some 0 < c < 1
48 Summary BMM and TC have same asymptotic time complexity Same status: Allpairs shortest paths (APSP) and Minplus matrix multiplication
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