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Handbook of Knowledge Representation Chapter 2: Satisfiability Solvers Junya Yamaguchi May 2, 2018 Tokyo Institute of Technology, Inoue Lab [email protected]

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SAT Problem Example • An assignment is a function ϕ : V → {0, 1} F = ( x literal ∨ y ∨ z ) ∧ (¬x ∨ ¬y) clause ∧ (¬y ∨ ¬z) ∧ (¬z ∨ ¬x) CNF formula assignment ϕ(x, y, z) = (0, 0, 1) derives F = 1. • A formula is satisfiable (SAT) iff there exists an assignment that evaluates it TRUE, which is called model. • Otherwise, we say a formula is unsatisfiable (UNSAT). 1

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SAT ιϧόͷ׆༂ • ࠷ѱܭࢉྔ͕ࢦ਺ؔ਺Ͱ͋Δ͜ͱͰ༗໊͕ͩɺSAT ιϧό͸ଟ͘ͷྖҬͰޭ ੷Λ࢒͍ͯ͠Δ • ιϑτ΢ΣΞ/ϋʔυ΢ΣΞݕূ • ࣗಈςετύλʔϯੜ੒ • ϓϥϯχϯά • εέδϡʔϦϯά • ೥ʹҰճ SAT ͷίϯϖ͕͋Δ (SAT Competition) • ͦ͜Ͱ਺ଟ͘ͷૉ੖Β͍͠ SAT ιϧό͕஀ੜͨ͠ • Ϟμϯͳ SAT ιϧόʹͳΔͱɺ਺ສม਺ɾ਺ઍ੍໿͔Β੒Δඇৗʹ೉͍͠໰ ୊Ͱ΋ղ͚Δ 2

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SAT ͱ஌ֶࣝशͷؔ܎ • ࠓͰ͸ͨ͘͞ΜͷԠ༻͕͋Δ͕ɺͦͷىݯ͸ ஌ֶࣝश • ஌ࣝͷදݱྗͱܭࢉྔͷτϨʔυΦϑʹؔ͢Δݚڀ (ୈ 3 ষͰѻ͏) ͕ϝΠϯ • Ծఆ: ࠷ѱܭࢉྔ͕ଟ߲ࣜ࣌ؒͰ͋Γͳ͕Β΋ɺΊͪΌͪ͘ΌΤϨΨϯτͰදݱ ྗͷߴ͍දݱݴޠΛΈ͚͍ͭͨ • 90 ೥୅ͷ͸͡Ίʹɺ2 ͭͷ࿦จ͕͜ͷԾఆʹ௅Μͩ • ͋Δಛघͳ໰୊Λআ͚͹ɺ΄ͱΜͲͷϥϯμϜ SAT ໰୊͸ͱͯ΋؆୯ʹͱ͚Δ ͜ͱΛࣔͨ͠࿦จ • ͦͷಛघͳ೉͍͠໰୊Ͱ͑͞΋ɺϩʔΧϧαʔνͷςΫχοΫΛ࢖͑͹؆୯ʹղ ͘͜ͱ͕Ͱ͖Δ͜ͱΛࣔͨ͠࿦จ 3

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ࢲͨͪ͸࠷ѱܭࢉྔʹͱΒΘΕͯͳ͍͔ʁ • ਺ඦສ΋ͷม਺Λ΋ͭݱ࣮ͷ SAT ໰୊Ͱͷ੒ޭ • యܕతͳ SAT ໰୊΍ݱ࣮తͳ NP ׬શͳ໰୊ (ͷܭࢉྔ) ʹରͯ͠͸ɺޮ཰Α͘ ղ͘͜ͱ͕Ͱ͖ΔҰൠղ๏͕͋Δ • ࠷ѱܭࢉྔʹϏϏΓ͗ͯ͢͸͍͚ͳ͍ • ͜ͷ·· SAT ιϧό͕੒௕͢Ε͹ɺ΋ͬͱෳࡶͳ஌ࣝදݱͷݴޠΛѻ͑ΔΑ ͏ʹͳΔ͸ͣ • NO: ࠷ѱ࣌ܭࢉྔ͕ଟ߲ࣜͳΞϧΰϦζϜͰͷදݱ • YES: SAT ιϧόͰղ͚Δൣғͷදݱ 4

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ষͷߏ੒ 1. ࣍ͷ 2 छྨͷιϧόʹ͓͍ͯ࢖ΘΕ͍ͯΔओͳςΫχοΫ • ׬શ (complete) SAT ιϧό • ෆ׬શ (incomplete) SAT ιϧό 2. ࣮ફతͳ SAT ූ߸Խʹର͢Δ্هςΫχοΫͷ༗ޮੑ 3. SAT ιϧόͷকདྷ΁ͷల๬ 5

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ఆٛ • ໋୊࿦ཧࣜ (propositional or Boolean formula) ͸ɺม਺ͷू߹্ʹఆٛ ͞ΕΔ࿦ཧࣜ • ֤ม਺͸ {FALSE, TRUE} ͷͲͪΒ͔ͷ஋ΛऔΔ • ศ্ٓ {0, 1} Ͱද͢ࣄ͕ଟ͍ • ม਺ͷू߹ V ʹର͢Δ ׬શׂ౰ͯ (truth assignment)1͸ɺࣸ૾ σ : V → {0, 1} ͷ͜ͱɻ • ಛʹɺ໋୊࿦ཧࣜΛ 1 ʹධՁ͢ΔΑ͏ͳ׬શׂ౰ͯΛ satisfying assignment ΍ Ϟσϧ (model) ͱݺͿɻ • SAT ιϧόͰѻ͏ SAT ໰୊͸ CNF (conjunctive normal form) ͱݺ͹Ε Δಛผͳܗࣜͷ໋୊࿦ཧ͚ࣜͩʹ੍ݶ 1 ׬શׂ౰ͯ͸୯ʹʮׂ౰ͯʯͱݺ͹ΕΔ͜ͱ΋͋Γ·͕͢ɺຊεϥΠυͰ͸෦෼ׂ౰ͯͷ͜ͱΛׂ౰ͯͱݺͼɺ׬શׂ౰ͯͱ۠ผ͠·͢ɻ 6

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ॆ଍Մೳੑ൑ఆ໰୊ Boolean satisfiability testing (SAT) ໰୊ ೖྗ CNF ܗࣜͷ໋୊࿦ཧࣜ (CNF ࣜ) F ࣭໰ F ʹ͸Ϟσϧ͕͋Δ͔ʁ CNF ࣜ અͷબݴ F = C1 ∧ C2 · · · ∧ Cm ·ͨ͸ F = {Ci }m i=1 અ Ϧςϥϧͷ࿈ݴ C = l1 ∨ l2 · · · ∨ ln ·ͨ͸ C = {li }n i=1 Ϧςϥϧ ม਺ x ͔ͦͷ൱ఆ ¬x. ม਺ x ∈ {0, 1} 7

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ఆٛ • અʹؚ·ΕΔϦςϥϧ਺ΛɺઅͷαΠζͱݺͿɻ • e.g. (x ∨ ¬y ∨ z) ͳΒαΠζ͸ 3 • αΠζ͕ 0 ͷઅΛۭઅ (empty clause)ɺαΠζ͕ 1 ͷઅΛ୯Ґઅ (unit clause)ɺͦͯ͠αΠζ͕ 2 ͷઅΛόΠφϦઅ (binary clause) ͱͦΕͧΕ ݺͿɻ • ͢΂ͯͷઅͷαΠζ͕ͪΐ͏Ͳ k Ͱ͋Δ CNF ࣜΛ k-SAT ͱݺͿɻ • 2-SAT ͸ଟ߲ࣜ࣌ؒͰٻղՄೳ • 3-SAT Ҏ্ʹͳΔͱ NP ׬શ (x1 ∨ x2 ∨ x3 ) ∧ (x4 ∨ x5 ∨ x6 ) ∧ (x7 ∨ x8 ∨ x9 ) 8

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ఆٛ • ෦෼ׂ౰ͯ (partial assignment) ͸ɺม਺ू߹ͷ෦෼ू߹ʹର͢Δ׬શׂ ౰ͯͷ͜ͱɻ • CNF ࣜ F ΁ͷ෦෼ׂ౰ͯ ρ ʹରͯ͠ɺρ Λ୅ೖͯ͠ಘΒΕͨࣜͷ͜ͱΛ simplified ͞ΕͨࣜͱݺͼɺF|ρ Ͱද͢ɻ • 1 ͭҎ্ͷϦςϥϧ͕ 1 ʹධՁ͞Εͨ͢΂ͯͷઅΛ࡟আ • 0 ʹධՁ͞Εͨ͢΂ͯͷϦςϥϧΛ࡟আ ࠓޙɺೖྗͱͯ͠༩͑Δ໋୊࿦ཧࣜʹ͸ CNF ࣜΛ҉໧ͷ͏ͪʹԾఆ͢Δ͕ɺ ଟ͘ͷ৔߹͜ͷԾఆ͸໰୊ʹͳΒͳ͍ɻ 9

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2.2 SAT Solver TechnologyŠ Complete Methods

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Complete Methods • A complete SAT solver, given the input formula F, either produces a satisfying assignment for F or proves that F is unsatisfiable. • Recent complete methods remain variants of a process introduced several decades ago, DPLL • DPLL procedure: • Was introduced in the early 1960`s. • Can prune of the search space based on falsified clauses. • Performs a backtrack search in the space of partial truth assignments. • Main improvements to DPLL: • Smart branch selection heuristics • Extensions like clause learning and randomized restarts • Well-crafted data structures such as lazy implementations and watched literals 10

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2.2.1 The DPLL Procedure

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DPLL Procedure 11

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DPLL Procedure • Repeatedly select an unassigned literal l. • The step to choose l is called branching step. • Setting l to TRUE or FALSE is called a decision • decision level is used to reffer the recursion depth at that stage • Recursively search for a satisfying assignment for F|l and F|¬l . 12

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DPLL Procedure • Repeatedly select an unassigned literal l. • The step to choose l is called branching step. • Setting l to TRUE or FALSE is called a decision • decision level is used to reffer the recursion depth at that stage • Recursively search for a satisfying assignment for F|l and F|¬l . 12

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DPLL Procedure • Repeatedly select an unassigned literal l. • The step to choose l is called branching step. • Setting l to TRUE or FALSE is called a decision • decision level is used to reffer the recursion depth at that stage • Recursively search for a satisfying assignment for F|l and F|¬l . 12

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DPLL Procedure • Repeatedly select an unassigned literal l. • The step to choose l is called branching step. • Setting l to TRUE or FALSE is called a decision • decision level is used to reffer the recursion depth at that stage • Recursively search for a satisfying assignment for F|l and F|¬l . 12

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DPLL Procedure 13

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DPLL Procedure – Unit Propagation • Initial state: F = (x) ∧ (¬x ∨ y) ∧ (¬x ∨ ¬y ∨ z) ρ = ∅ • There exists an unit clause (x). Execute unit propagation. F = (x)∧(¬x∨y) ∧ (¬x∨¬y ∨ z) ρ = {x} • Now, another unit clause (y) found, let’s more simplify the formula. F = (x) ∧ (¬x ∨ y)∧(¬x ∨ ¬y∨z) ρ = {x, y} • Execute unit propagation in the same way. F = (x) ∧ (¬x ∨ y) ∧ (¬x ∨ ¬y ∨ z) ρ = {x, y, z} • UnitPropagate() ends since there are no unit clauses. 14

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • a.k.a. decision strategy • MOMS, BOHM maximize a moderately complex function of the dcurrent var. and cls. state. • DLIS selects and fix the literal occuring most frequently in the yet unsatisfied clauses • VSIDS chooses a literal based on its weight which preiodically decays but is boosted if a clause in which it appears is used in deriving a conflict. 15

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • Critical role in the success of modern complete SAT solvers. • The idea here is to • cache lcauses of conflictz as learned clauses. • utilize this information to prune the search in a different part of the search space encountered later. 16

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • is a implementaion technique to accelerate unit propagation, introduced in zChaff. • which key idea is to maintain and lwatchztwo special literals for each not yet satisfied clause. • {U, U} : Not yet statisfied • {0, U} : Unit clause • {0, 0} : Empty clause • has high compatibility with clause learning. 17

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • Backtrack just returns to previous branching point. • Backjump repeats backtrack safely as long as possible. 18

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • It lets a solver jump directly to a lower decision level d where; • even one branch leads to a conflict involving variables at levels d or lower only. • for completeness, the level d is not marked as unsatisfiable. • While conflict-directed backjumping is always beneficial, fast backjumping may not be so. 19

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • It’s a technique to learn smaller and more pertinent learnt clauses. • When a conflict occurs because of a clause C′, and the size of learnt clause C exceeds a certain threshold length; • the solver backtracks to almost the highest decision level of the literals in C, • it then starts assigning to FALSE the unassigned literals of C′ until a new conflict is encountered. 20

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • The idea is to try to reduce the size of a learned conflict clause C by repeatedly identifying and removing any literals of C that are implied to be FALSE when the rest of the literals in C are set to FALSE. a ∨ b, ¬a ∨ c b ∨ c 21

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Key Features of Modern DPLL-Based SAT Solvers • Variable selection heuristic • Clause learning • The watched literals scheme • Conflict-directed Backjump • Fast backjump • Assignment stack shrinking • Cnflict clause minimization • Randomized restrts • It allows a SAT solver to arbitrarily stop the search and restart their branching process from decision level zero. • Most of the current SAT solvers, employ aggressive restart strategies, sometimes restarting after as few as 20 to 50 backtracks. 22

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Clause Learning and Iterative DPLL 23

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Clause Learning and Iterative DPLL 23

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Clause Learning and Iterative DPLL 23

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Clause Learning and Iterative DPLL 23

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Clause Learning and Iterative DPLL 24

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Clause Learning and Iterative DPLL 24

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Clause Learning and Iterative DPLL 24

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Clause Learning and Iterative DPLL 24

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Clause Learning and Iterative DPLL 24

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Cleause Learning • The idea of clause learning came from Artificial Intelligence; • seeking to improve backtrack search algorithms by generating explanations for failure points, • For general constraint satisfaction problems the explanations are called lconflictszor lno-goodsz. • In early works, clause learning could only obtain unusable clauses, but various studies have led to success. 25

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Implication graph Implication graph The implication graph G at a given stage of DPLL is a directed acyclic graph with edges labeled with sets of clauses. 26

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How to build implication graph 1. Create a node for each decision literal, labeled with that literal. 2. If there is an unit clause C when regarding nodes as the partial assingment; C = k ∨ i=1 li ∨ l, where {li } are assigned to FALSE • Add a node l (if not in G) • Add edges (li , l) for all i = 1, . . . , k (if not in G) • These edges are labeld with C to specify the cause of implication. 3. Repeat step 2 while no such clause found. 4. Add a special ’conflict’ node ¯ Λ. For any variable x s.t. x, ¬x are both in G, add edges (x, ¯ Λ) and (¬x, ¯ Λ). 27

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Conflict on Implication Graph • An implication graph may not contain a conflict at all, or it may contain many conflict variables and several ways of deriving any single literal. 28

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Conflict Graph A conflict graph H = (VH , EH ) is a subgraph of the implication graph with the following propeties: • H contains ¯ Λ and exactly one conflict variable. • ∀l ∈ VH have a path to ¯ Λ • ∀l ∈ VH \ { ¯ Λ } satisfies either of the following: 1. l is a decision literal 2. ∃¬l1 , . . . , ¬lk ∈ VH s.t. ∀¬li holds (¬li , l) ∈ EH and there exists a known clause (l1 ∨ · · · ∨ lk ∨ l) 29

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Conflict Graph A conflict graph H = (VH , EH ) is a subgraph of the implication graph with the following propeties: • H contains ¯ Λ and exactly one conflict variable. • ∀l ∈ VH have a path to ¯ Λ • ∀l ∈ VH \ { ¯ Λ } satisfies either of the following: 1. l is a decision literal 2. ∃¬l1 , . . . , ¬lk ∈ VH s.t. ∀¬li holds (¬li , l) ∈ EH and there exists a known clause (l1 ∨ · · · ∨ lk ∨ l) 29

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Reason/Conflict Side 30

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Learning Scheme - Relsat • Relsat uses the cut whose conflict side consists of all implied variables at the current decision level. • The learnt clause (¬a ∨ ¬b) has exactly one variable from the current decision level (i.e. b). • After learning it and backtraking until unassigning b (to decision level 2), filpping the value b to ¬b. 31

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Learning Scheme - Relsat • Relsat uses the cut whose conflict side consists of all implied variables at the current decision level. • The learnt clause (¬a ∨ ¬b) has exactly one variable from the current decision level (i.e. b). • After learning it and backtraking until unassigning b (to decision level 2), filpping the value b to ¬b. 32

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Unique Implication Points (UIPs) • An UIP (unique implication point) is a node at the latest decision level s.t. • every path from the lastest decision variable to conflict literal must go through the node. • Intuitively, an UIP is single reason at the current decisioin level. • There can be more than one UIP , but not 0. • The decision variabe is the obvious UIP. 33

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Application of UIPs • Learning scheme with UIP: • relsat uses the decision variable as UIP • Grasp and zChaff use FirstUIP, which is the UIP closest to the conflict variable. • Grasp also uses all UIPs to learn multiple clauses. • kUIP – Generalized UIP to k previous decision levels. • However, the study showed 1UIP is quite robust and outperforms all other schemes they consider on most of the benchmarks. 34

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A Proof Complexity Perspective

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Propositional proof complexity • Propositional proof complexity is the study of the proofs of validity of mathematical statements expressed in a Boolean form. • A propositional proof system is an algorithm A Input A propositional statement S and a purported proof π Returns Rejection / Acceptance • The crucial property of A is that: • For all invalid S, A rejects the pair (S, π) for all π • For all valid S, A accepts the pair (S, π) for some π 35

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Resolution • Resolution is a very simple proof system with only one rule. a ∨ b, ¬a ∨ c b ∨ c • Repeated application of this rule can derive an empty disjunction iff the initial formula is unsatisfiable; • Such a derivation serves as a proof of unsatisfiability of the formula. • There are several resolution types due to the implementaion restriction. 36

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General Resolution 1 Paul Beame University of Washington Satisfiability and Unsatisfiability: Proof Complexity and Algorithms. http://slideplayer.com/slide/4978670/ 37

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Tree-like Resolution 1 Paul Beame University of Washington Satisfiability and Unsatisfiability: Proof Complexity and Algorithms. http://slideplayer.com/slide/4978670/ 38

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Variation of Resolution • Tree-like resolution uses non-empty derived clauses exactly once in the proof and is equivalent to an optimal DPLL procedure. • Regular resolution allows any variable to be resolved upon at most once along any “path” in the proof from an initial clause to Λ, allowing (restricted) reuse of derived clauses. • While these and other refinements are sound and complete as proof systems, they differ vastly in efficiency. 39

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SAT Solver as Proof System • Most of today`s complete SAT solvers implement a subset of the resolution proof system. • Till recently, it was not clear: • Where do they fit in the proof system hierarchy? • How do we compare them to refinements of resolution, e.g. regular resolution? • Clause learning and random restarts was considered to be important. • Despite overwhelming empirical evidence, for many years not much was known of the ultimate strengths and weaknesses of the two. 40

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FIrstNewCut • Beame, Kautz, and Sabharwal answered several of these questions in a formal proof complexity framework. • They characterized clause learning as a proof system called CL, and proposed a new learning scheme called FirstNewCut • FirstNewCut can provide exponentially shorter proofs than refinement of general resolution satisfying a natural self-reduction property. • These include regular and ordered resolution. • These are already known to be much stronger than the ordinary DPLL procedure. • They also showed that a slight variant of clause learning with unlimited restarts is as powerful as general resolution itself. 41

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Ralation Between SAT/UNSAT Proof Complexity • INTEREST is a familiy of unsatisfiable formulas, BECAUSE only proofs of unsatisfiability can be large. • The minimum proofs of satisfiability are O(n). • However, in practice, many formulas are satisfiable. • Achlioptas, Beame, and Molloy have shown • Negative proof complexity results for unsatisfiable formulas can be used; • for the run-time lower bounds of specific inference algorithms, especially DPLL, running on satisfiable formulas as well. • The key observation in their work is that before hitting a satisfying assignment, an algorithm is very likely to explore a large unsatisfiable part of the search space due to uninformative initial assignments. 42

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Branching Heuristic for SAT/UNSAT proofs. • Proof complexity does not capture everything we intuitively mean by the power of a reasoning system because it says nothing about how difficult it is to find shortest proofs. • However, it is a good notion with which to begin our analysis because the size of proofs provides a lower bound on the running time of any implementation of the system. Negative Positive Proofs Size Must be large Small proofs exist Perfect Branching Hopeless Possible 43

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DPLL and Tree-like Resolution For a CNF formula F, the size of the smallest DPLL refutation of F is equal to the size of the smallest tree-like resolution refutation of F. • The best performances are the same. • Considering a propositional proof system CL is exponentially stronger than tree-like resolution, clause learning is very effective. 44

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Natural Refinement / Proper • CS (F) : the length of a shortest refutation of a formula F under a proof system S • tree-like, regular, linear, positive, negative, semantic, and ordered resolution are natural refinements. • tree-like, reglar, and ordered resolution are exponentially-proper. 45

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The Power of CL • (pros) There are some formulas that CL is more effective than some exponentially-propered resolutions. • (cons) CL may not be able to simulate all regular resolution proofs! 46

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CL– with non-redundant vs. RES • CL– – : CL allowed to branch on a literal whose value is already assigned. • Such a branch can: • lead to an immediate conflict; • allow one to learn a key conflict clause that would otherwise have not been learned. • A clause learning scheme will be called non-redundant if on a conflict, it always learns a clause not already known. • Most of the practical clause learning schemes are non-redundant. 47

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Symmetry Breaking

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Symmetry • Symmetry in real-world problems: • in FPGA routing, all available wires or channels used for connecting two switch boxes are equivalent; • in circuit modeling, all gates of the same ltypezare interchangeable; • in planning, all identical boxes that need to be moved from city A to city B are equivalent; • in multi-processor scheduling, all available processors are equivalent; • in cache coherency protocols in distributed computing, all available identical caches are equivalent. • A key property of such objects is that when selecting k of them, we can choose any k, without loss of generality . • We would like to treat without-loss-of-generality reasoning in an automatic fashion. 48

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Semantic Meaning • A CNF formula consists of constraints over different kinds of variables that typically represent tuples of these high level objects. • During the problem modeling phase, we could have a Boolean variable zw,c that is TRUE iff the first end of wire w is attached to connector c. • When the formula is converted into DIMACS format, the semantic meaning of the variables is discarded. 49

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Symmetry Breaking Predicates (SBPs) • Symmetry Breaking Predicates (SBPs) is a method to deal with this problem: • Symmetries can be broken by adding one lexicographic ordering constraint, called llex-constraintsz. • This constraint removes all redundant solutions but the lexically-first solutions. • The idea is to identify the set of permutations of variables {πi }, where πi keeps the formula unchanged. • Let σ be a lexically-first solution, πi (σ) is also a satisfying assignment. 50

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Symmetry in SAT – Graph Isomorphism • Shatter by Aloul et al. improves the number of lex-constraints from O(n2) to O(n), where n is the number of variables. • They use techniques of graph isomorphism to generate SBPs. • Since there is no polynomial-time algorithms to praph isomorphism, the number os SBPs to break all symetries can be exponential. • Shatter handled this by discarding “large” symmetries. 51

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Symmetry in SAT – Non-CNF Formulation • Utilizing non-CNF formulations , known as pseudo-Boolean inequalities, is another approach for symmetry problem. • PBS by Aloul et al. • pbChaff by Dixon et al. • Galena by Chai and Kuehlmann • Instead of resolution system, they use Cutting Planes proof system. • It is difficult to implement, so pseudo-Boolean solvers often implement only a subset of it. 52

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2.3 SAT Solver Technology - Incomplete Methods

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Incomplete SAT Solver • What about an incomplete method? • It can provide the model of satisfiable formula, but it is not always possible to solve; • It can’t say anything about unsatisfiability. • Incomplete methods significantly outperform DPLL-base methods for some problems. • They are generally based on stochastic local search, while typical complete methods are based on exhaustive branching and backtracking search. 53

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Local Search on Satisfiability • Original impetus of local search on satisfiability problems was the successful application for finding solutions to N-queens problems. • Researchers regarded N-queens was an easy problem. • They felt that it would fail in practice for SAT. 54

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GSAT • GSAT is based on randomized local search technique. • Initialize truth assignment randomlly. • Then, greedily flip the assignment of the variable that leads to the greatest decrease in the total number of UNSAT clauses. • Finish the search if a model is found, otherwise repeat fliping. 55

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GSAT • GSAT outperformed DPLL-based solvers on various SAT problem. • From randomly generated problems to SAT encodings of graph coloring problems. • Most time is spent searching plateaus • Plateaus are possible sideways movements that don’t increase/ decrease the # of UNSAT clauses too much. 56

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Walksat 57

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Walksat • Walksat is a refinement of GSAT, employing random walk moves of a standard Metropolis search. • For the purpose to increase randomness. • Always select the variable to flip from unsatisfied clauses. • If there is a variable that can be flipped without affecting satisfied clauses, pick it up. • Otherwise, perform a p-greedy strategy. • This idea is inpired by the O(n2) randomized algorithm for 2-SAT. • It guarantees that one can reach a model by selecting only from variables on unsatified clauses. 58

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Walksat vs. GSAT http://ai.cs.unibas.ch/_files/teaching/fs16/ai/slides/ai32.pdf 59

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Phase Transition Phenomenon

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Clause-to-Variable Ratio • Random k-SAT with n variables and m clauses. • For each clause, select k unique variables uniformaly. • For each variable, negate it with probablity 0.5 • clause-to-variable ratio α = m n • The difficulity of random k-SAT problems is characterized by α. • For random 3-SAT, the run-time of problems with α > 4.26 tends to be much larger than the opposite one. 60

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Phase Transition • Left: The run-time • Right: The ratio of unsatisfiable problem • As n grows, the phenomenon becomes shaper and shaper. 61

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2.3.2 Survey Propagation

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Survey Propagation • Survey Propagation (SP) is a new algorithm for solving hard combinational problems. • It showed good results on random 3-SAT instances with one million variables. • SP has a connection to Blief Propagation (BP): • BF is a method to calculate marginal probability on graphical models (factor-graph). • Both using an iterative process of local “message” update. • Unfortunately, the success of SP is limited to random SAT instances. 62

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Survey Propagation http://slideplayer.com/slide/4559400/ 63

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2.4 Runtime Variance and Problem Structure

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Exceptionally Hard Problem • The performance of backtrack-style modhods can vary: • The way to select the next variable to branch on. • In what order the possible values are assigned to the variable. • There are exceptionally hard problems. • They ignore the “easy-hard-easy” phenomenon. • They occur in the under-constrained area. • However, subsequent research showed that such instances are not inherently difficult. • Simply renaming the variables or by considering a different search heuristic such can easily solved such problems. • The “hardness” is in the combination of the instance with the search method, rather than in the instances per se2. • This is why the median is used to measure the “hardness” of problems. 2Adverb: by or in itself or themselves; intrinsically. 64

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Fat and Heavy Tailed Behavior • Runtime distributions of search methos provides a better characterization. • Quite often complete backtrack search methods exhibit fat and heavy-tailed behavior. • Looks like lognormal, Wibull dist. • Such distributions can be observed: • when running a deterministic backtracking procedure on a distribution of random instances; • by repeated runs of a randomized backtracking procedure on a single instance. 65

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Runtime of CSP Summarizing CSP Hardness with Continuous Probability Distributions. Proceedings of the National Conference on Artificial Intelligence. 327-333. 66

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Heavy Tail • fat-tailedness – • kurtosis = µ4 µ2 2 , where µ2 , µ4 are 2nd and 4th moment resp. • A distribution with igh central peak and long tails has generally large kurtosis. • A heavy-taild distribution is “heavier” than fat-tailed one. • It have some infinite moments, e.g. mean, variance, etc. • DPLL style complete backtrack search methods have been shown to exhibit heavy-tailed behavior. • They can solve some problems so quickly, meanwhile do some problems super slowly. 67

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2.4.2 Backdoors

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Backdoors • The idea of heavy-tailed behavior comes from backdoor variables . • Variables which gives us a polynomial subproblem when set. • This is for explaining how a backtrack search method can get “lucky” on certain runs. • When backdoor variables are identified early on in the search and set the right way, we say it’s lucky. • The definition of a backdoor depends on a particular algorithm, reffered to as sub-solver. • It solves a tractable sub-case of the general constraint satisfaction problem. 68

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Sub-Solver and Backdoor • Intuitively, the backdoor are variables such that when set correctly, the sub-solver can solve the remaining problem. • A stronger notion of backdoors, strong backdoor, considers both satisfiable and unsatisfiable problems. 69

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Backdoor Detection Complexity • Nishimura et al. provided that detecting backdoor sets where the sub-solver solves Horn or 2-CNF formulas, both of which are linear time problems. • They proved that the detection of such a strong backdoor set is fixed-parameter tractable; • while the detection of a weak backdoor set is not. • Dilkina et al. proved that the complexity of backdoor detection jumps from NP to NP-hard/coNP-hard when adding certain obvious inconsistency checks to the underlying class. 70

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Backdoor Size • The backdoor size of random formulas is large. • For random 3-SAT near the phase transition, about 30% of variables are backdoors. • That’s why DPLL-based search algorithm can’t solve these problems • Structured (real-world) problems have very few backdoor variables. 71

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2.4.3 Restarts

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Restart Strategy • Another way to exploit heavy-tailed behavior is to add restarts to backtracking. • Gomes et al. proposed randomized rapid restarts (RRR) to take advantage of heavy-tailed behavior and boost the efficiency of complete backtrack search. • They also showed that restart strategy with a fixed cutoff eliminates heavy-tail behavior and has finite moments. 72

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Luby’s Restarts • Luby et al. showed that: • When the distriution is known, the optimal restart policy is the fixed cutoff. • When there is no a priori to the distribution, a universal strategy minimizes the expected cost. 1, 1, 2, 1, 1, 2, 4, 1, 1, 2, 4, 8, . . . • Theoretically the universal strategy is within a log factor of the optimal one, but the convergence is too slowly in practice. • State-of-the-art SAT solver’s restart: • Default cutoff value, which is increased linearly 73

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Bayesian Framework • Horvitz et al. introduced a Bayesian framework for learning predictive models of randomized backtrack solvers based on “knowledgeless” situation. • Kautz et al. extended it, enabling restart policies to use information of real-time observations about a solver. • They also showed the optimal policy for dynamic restarts. • It can be implemented using dynamic programming technique. 74