DOCSLIB.ORG

Lecture 19,20 (Nov 15&17, 2011): Hardness of Approximation, PCP

Total Page：16

File Type：pdf, Size：1020Kb

Download full-text PDF Read full-text
Lecture 19,20 (Nov 15&17, 2011): Hardness of Approximation, PCP ,20 CMPUT 675: Approximation Algorithms Fall 2011 Lecture 19,20 (Nov 15&17, 2011): Hardness of Approximation, PCP theorem Lecturer: Mohammad R. Salavatipour Scribe: based on older notes 19.1 Hardness of Approximation So far we have been mostly talking about designing approximation algorithms and proving upper bounds. From no until the end of the course we will be talking about proving lower bounds (i.e. hardness of approximation). We are familiar with the theory of NP-completeness. When we prove that a problem is NP-hard it implies that, assuming P=NP there is no polynomail time algorithm that solves the problem (exactly). For example, for SAT, deciding between Yes/No is hard (again assuming P=NP). We would like to show that even deciding between those instances that are (almost) satisfiable and those that are far from being satisfiable is also hard. In other words, create a gap between Yes instances and No instances. These kinds of gaps imply hardness of approximation for optimization version of NP-hard problems. In fact the PCP theorem (that we will see next lecture) is equivalent to the statement that MAX-SAT is NP- hard to approximate within a factor of (1 + ǫ), for some fixed ǫ> 0. Most of hardness of approximation results rely on PCP theorem. For proving a hardness, for example for vertex cover, PCP shows that the existence of following reduction: Given a formula ϕ for SAT, we build a graph G(V, E) in polytime such that: 2 • if ϕ is a yes-instance, then G has a vertex cover of size ≤ 3 |V |; 2 • if ϕ is a no-instance, then every vertex cover of G has a size > α 3 |V | for some fixed α> 1. Corollary 1 The vertex cover problem can not be approximated with a factor of α unless P = NP. In this reduction we have created a gap of size α between yes/no instances. Figure 19.1: g is a gap-introducing reduction 19-1 19-2 Lecture 19,20: Hardness of Approximation, PCP theorem Definition 1 Suppose that L is an NP-complete problem and π is a minimization problem. Suppose that g is a function computable in polytime that maps Yes-instances of L into a set S1 of instances of π and No-instances of L into a set S2 of instances of π. Assume that there is a polytime computable function h such that: • for every Yes-instance x of L: OPT (g(x)) ≤ h(g(x)); • for every No-instance x of L: OPT (g(x)) > αh(g(x)). Then g is called a gap-introducing reduction from L to π and α is the size of the gap. The chromatic number of a graph G, denoted by χ(G), is the smallest integer c such that G has a c-coloring. It is known that deciding whether a graph is 3-colorable or not is NP-complete. Using this we easily get the following hardness of approximation for the optimization problem of computing the chromatic number of a given graph G. 4 Theorem 1 There is no better than 3 -approximation algorithm for computing the chromatic number, unless P = NP . Proof. Create a gap-introducing reduction from 3-colorability to the problem of computing chromatic numbers, as shown in the figure below. 4 Figure 19.2: No better than 3 -approximation algorithm for computing the chromatic number, unless P = NP . The reduction we gave for the (non-metric) traveling salesman problem in Lecture 5 was a gap-introducing reduction from Hamilton cycle. Once we have a gap-introducing reduction (e.g. from SAT) to π1, then we can prove a hardness for another problem π2 by giving a gap-preserving reduction from π1 to π2. Definition 2 A gap-preserving reduction g (g is polytime computable) from π1 to π2 has four parameters f1,α,f2,β. Given instance x of π1, g(x) is an instance of π2 such that • if OPT (x) ≤ f1(x) −→ OPT (g(x)) ≤ f2(g(x)); • if OPT (x) > αf1(x) −→ OPT (g(x)) > βf2(g(x)). Lecture 19,20: Hardness of Approximation, PCP theorem 19-3 19.1.1 MAX-SNP Problems and PCP Theorem Many problems, such as Bin Packing and Knapsack, have PTAS’s. A major open question was: does MAX-3SAT have a PTAS? A significant result on this line was the paper by Papadimitriou and M. Yannakakis (’92) where they defined the class MAX-SNP. All problems in MAX-SNP have constant approximation algorithms. They also defined the notion of completeness for this class and showed if a MAX-SNP-complete problem has a PTAS then every MAX-SNP problem has a PTAS. Several well-known problems including MAX-3SAT and 7 TSP are MAX-SNP-complete. We have seen a 8 -approximation for MAX-3SAT. It turns out that even if there 7 are not exactly 3 literals per clause this problem has a 8 -approximation algorithm. How about lower bounds? There is a trivial gap-introducing reduction form 3-SAT to MAX-3SAT which shows that formula ϕ is a Yes- m−1 instance iff more than m fraction of ϕ can be satisfied. Suppose that for each L ∈ NP there is a polytime computable g from L to instances of MAX-3SAT, such that • for Yes-instance y ∈ L, all 3-clauses in g(y) can be satisfied; • for No-instance y ∈ L, at most 99% of the 3-clauses of g(y) can be satisfied. There is a standard proof system for L ∈ NP , that is, a deterministic verifier that takes as input an instance y of L together with a proof π, reads π and accepts iff π is a valid proof (solution). We give an alternative proof system for NP using the gap-introducing reduction g assumed above. Define a proof that y ∈ L to be a truth assignment satisfying g(y) where g is the gap-instance reduction from L to MAX-3SAT. We define a randomized verifier V for L. V runs in polynomial time and • takes y and the “new” proof π; • accept iff π satisfies a 3-clause selected uniformly and randomly from g(y). If y ∈ L then there is a proof (truth assignment) such that all the 3-clauses in g(y) can be satisfied. For that proof Verifier V (y, π) accepts with probability 1. If y∈ / L then every truth assignment satisfies no greater of 99% of clauses in g(y). So verifier V (y, π) will reject 1 with probability not less than 100 . Note that under the assumption we made, we can achieve a constant separation between the two cases above by reading only 3 bits of the proof (regardless of the length of the proof). We can increase the probabability of success (i.e. decrease the probability of error) by simulating V a constant (but large) number of times. For 1 example, by repeating 100 times, the probability of success will be at least 2 in the second case and we are still checking a constant (300) bits of the proof. 19.2 PCP Theorem To be consistent with the PCP definition later on, we give a slightly modified definition for class NP which is clearly equivalent to the original one: Definition 3 A language L is ∈ NP if and only if there is a deterministic polynomial time verifier (i.e. algorithm) V that takes an input x and a proof y with |y| = |x|c for a constant c > 0 and it satisfies the following: 19-4 Lecture 19,20: Hardness of Approximation, PCP theorem • Completeness: if x ∈ L ⇒ ∃ y such that V (x, y)=1. • Soundness: if x∈ / L ⇒ ∀y, V (x, y)=0. Now we define probabilistic verifiers that are restricted to look at (query about) only a few bits of the alleged proof y and make their decision based on those few bits instead of reading the whole proof. Definition 4 A (r(n),b(n))-restricted verifier is a randomized verifier that uses at most r(n) random bits. It runs in probabilistic polynomial time and reads/queries at most b(n) bits of the proof. Finally we provide the main definition of a language ∈ P CP . This definition is for P CP 1 (r(n),b(n)). 1, 2 Definition 5 A language L is ∈ P CP (r(n),b(n)) if and only if there is a (r(n),b(n))-restricted verifier V such that given an input x with length |x| = n and a proof π, it satisfies the following: • Completeness: if x ∈ L ⇒ ∃ a proof π such that P r[V (x, π)=1]=1. 1 • Soundness: if x∈ / L ⇒ ∀π, P r[V (x, π) = 1] ≤ 2 . 1 The probabilities in completeness and soundness given in definition above are 1 and 2 , respectively. A more general definition of P CP languages is by allowing these probabilities be some other constant values: Definition 6 For 0 ≤ s<c ≤ 1, a language L is ∈ P CPc,s(r(n),b(n)) if and only if there is a (r(n),b(n))- restricted verifier V such that given an input x with length |x| = n and a proof π, it satisfies the following: • Completeness: if x ∈ L ⇒ ∃ a proof π such that P r[V (x, π) = 1] ≥ C. • Soundness: if x∈ / L ⇒ ∀π, P r[V (x, π) = 1] ≤ S. In the more general definition of P CP language, we need to have the following restrictions on parameters: • c and s are constants and 0 ≤ S < C ≤ 1.
Load more

Recommended publications
  • Ab Pcpchap.Pdf
    i Computational Complexity: A Modern Approach Sanjeev Arora and Boaz Barak Princeton University http://www.cs.princeton.edu/theory/complexity/ [email protected] Not to be reproduced or distributed without the authors’ permission ii Chapter 11 PCP Theorem and Hardness of Approximation: An introduction “...most problem reductions do not create or preserve such gaps...To create such a gap in the generic reduction (cf. Cook)...also seems doubtful. The in- tuitive reason is that computation is an inherently unstable, non-robust math- ematical object, in the the sense that it can be turned from non-accepting to accepting by changes that would be insignificant in any reasonable metric.” Papadimitriou and Yannakakis [PY88] This chapter describes the PCP Theorem, a surprising discovery of complexity theory, with many implications to algorithm design. Since the discovery of NP-completeness in 1972 researchers had mulled over the issue of whether we can efficiently compute approxi- mate solutions to NP-hard optimization problems. They failed to design such approxima- tion algorithms for most problems (see Section 1.1 for an introduction to approximation algorithms). They then tried to show that computing approximate solutions is also hard, but apart from a few isolated successes this effort also stalled. Researchers slowly began to realize that the Cook-Levin-Karp style reductions do not suffice to prove any limits on ap- proximation algorithms (see the above quote from an influential Papadimitriou-Yannakakis paper that appeared a few years before the discoveries described in this chapter). The PCP theorem, discovered in 1992, gave a new definition of NP and provided a new starting point for reductions.
    [Show full text]
  • Pointer Quantum Pcps and Multi-Prover Games
    Pointer Quantum PCPs and Multi-Prover Games Alex B. Grilo∗1, Iordanis Kerenidis†1,2, and Attila Pereszlényi‡1 1IRIF, CNRS, Université Paris Diderot, Paris, France 2Centre for Quantum Technologies, National University of Singapore, Singapore 1st March, 2016 Abstract The quantum PCP (QPCP) conjecture states that all problems in QMA, the quantum analogue of NP, admit quantum verifiers that only act on a constant number of qubits of a polynomial size quantum proof and have a constant gap between completeness and soundness. Despite an impressive body of work trying to prove or disprove the quantum PCP conjecture, it still remains widely open. The above-mentioned proof verification statement has also been shown equivalent to the QMA-completeness of the Local Hamiltonian problem with constant relative gap. Nevertheless, unlike in the classical case, no equivalent formulation in the language of multi-prover games is known. In this work, we propose a new type of quantum proof systems, the Pointer QPCP, where a verifier first accesses a classical proof that he can use as a pointer to which qubits from the quantum part of the proof to access. We define the Pointer QPCP conjecture, that states that all problems in QMA admit quantum verifiers that first access a logarithmic number of bits from the classical part of a polynomial size proof, then act on a constant number of qubits from the quantum part of the proof, and have a constant gap between completeness and soundness. We define a new QMA-complete problem, the Set Local Hamiltonian problem, and a new restricted class of quantum multi-prover games, called CRESP games.
    [Show full text]
  • Probabilistic Proof Systems: a Primer
    Probabilistic Proof Systems: A Primer Oded Goldreich Department of Computer Science and Applied Mathematics Weizmann Institute of Science, Rehovot, Israel. June 30, 2008 Contents Preface 1 Conventions and Organization 3 1 Interactive Proof Systems 4 1.1 Motivation and Perspective ::::::::::::::::::::::: 4 1.1.1 A static object versus an interactive process :::::::::: 5 1.1.2 Prover and Veri¯er :::::::::::::::::::::::: 6 1.1.3 Completeness and Soundness :::::::::::::::::: 6 1.2 De¯nition ::::::::::::::::::::::::::::::::: 7 1.3 The Power of Interactive Proofs ::::::::::::::::::::: 9 1.3.1 A simple example :::::::::::::::::::::::: 9 1.3.2 The full power of interactive proofs ::::::::::::::: 11 1.4 Variants and ¯ner structure: an overview ::::::::::::::: 16 1.4.1 Arthur-Merlin games a.k.a public-coin proof systems ::::: 16 1.4.2 Interactive proof systems with two-sided error ::::::::: 16 1.4.3 A hierarchy of interactive proof systems :::::::::::: 17 1.4.4 Something completely di®erent ::::::::::::::::: 18 1.5 On computationally bounded provers: an overview :::::::::: 18 1.5.1 How powerful should the prover be? :::::::::::::: 19 1.5.2 Computational Soundness :::::::::::::::::::: 20 2 Zero-Knowledge Proof Systems 22 2.1 De¯nitional Issues :::::::::::::::::::::::::::: 23 2.1.1 A wider perspective: the simulation paradigm ::::::::: 23 2.1.2 The basic de¯nitions ::::::::::::::::::::::: 24 2.2 The Power of Zero-Knowledge :::::::::::::::::::::: 26 2.2.1 A simple example :::::::::::::::::::::::: 26 2.2.2 The full power of zero-knowledge proofs ::::::::::::
    [Show full text]
  • Lecture 19: Interactive Proofs and the PCP Theorem
    Lecture 19: Interactive Proofs and the PCP Theorem Valentine Kabanets November 29, 2016 1 Interactive Proofs In this model, we have an all-powerful Prover (with unlimited computational prover) and a polytime Verifier. Given a common input x, the prover and verifier exchange some messages, and at the end of this interaction, Verifier accepts or rejects. Ideally, we want the verifier accept an input in the language, and reject the input not in the language. For example, if the language is SAT, then Prover can convince Verifier of satisfiability of the given formula by simply presenting a satisfying assignment. On the other hand, if the formula is unsatisfiable, then any \proof" of Prover will be rejected by Verifier. This is nothing but the reformulation of the definition of NP. We can consider more complex languages. For instance, UNSAT is the set of unsatisfiable formulas. Can we design a protocol so that if the given formula is unsatisfiable, Prover can convince Verifier to accept; otherwise, Verifier will reject? We suspect that no such protocol is possible if Verifier is restricted to be a deterministic polytime algorithm. However, if we allow Verifier to be a randomized polytime algorithm, then indeed such a protocol can be designed. In fact, we'll show that every language in PSPACE has an interactive protocol. But first, we consider a simpler problem #SAT of counting the number of satisfying assignments to a given formula. 1.1 Interactive Proof Protocol for #SAT The problem is #SAT: Given a cnf-formula φ(x1; : : : ; xn), compute the number of satisfying as- signments of φ.
    [Show full text]
  • QMA with Subset State Witnesses
    QMA with subset state witnesses Alex B. Grilo1, Iordanis Kerenidis1,2, and Jamie Sikora2 1 LIAFA, CNRS, Universite´ Paris Diderot, Paris, France 2Centre for Quantum Technologies, National University of Singapore, Singapore One of the notions at the heart of classical complexity theory is the class NP and the fact that deciding whether a boolean formula is satisfiable or not is NP-complete [6, 13]. The importance of NP-completeness became apparent through the plethora of combinatorial problems that can be cast as constraint satisfaction problems and shown to be NP-complete. Moreover, the famous PCP theorem [3, 4] provided a new, surprising description of the class NP: any language in NP can be verified efficiently by accessing probabilistically a constant number of bits of a polynomial- size witness. This opened the way to showing that in many cases, approximating the solution of NP-hard problems remains as hard as solving them exactly. An equivalent definition of the PCP theorem states that it remains NP-hard to decide whether an instance of a constraint satisfaction problem is satisfiable or any assignement violates at least a constant fraction of the constraints. Not surprisingly, the quantum analog of the class NP, defined by Kitaev [11] and called QMA, has been the subject of extensive study in the last decade. Many important properties of this class are known, including a strong amplification property and an upper bound of PP [14], as well as a number of complete problems related to the ground state energy of different types of Hamiltonians [11, 10, 15, 7, 8].
    [Show full text]
  • Notes for Lecture 1 1 a Brief Historical Overview
    U.C. Berkeley Handout N1 CS294: PCP and Hardness of Approximation January 18, 2006 Professor Luca Trevisan Scribe: Luca Trevisan Notes for Lecture 1 1 A Brief Historical Overview This lecture is a “historical” overview of the field of hardness of approximation. Hundreds of interesting and important combinatorial optimization problems are NP-hard, and so it is unlikely that any of them can be solved by a wost-case efficient exact algorithm. Short of proving P = NP, when one deals with an NP-hard problem one must accept a relaxation of one or more of the requirements of having a an optimal algorithm that runs in polynomial time on all inputs. Possible ways of “coping” with NP-hardness are to design algorithms that 1. Run in “gently exponential” time, such as, say, O(1.1n). Such algorithms can be reasonably efficient on small instances; 2. Run in polynomial time on “typical” instances. Such algorithms are extremely useful in practice. A satisfactory theoretical treatment of such algorithms has been, so far, elusive, because of the difficulty of giving a satisfactory formal treatment of the notion of “typical” instances. (Most average-case analysis of algorithms is done on distributions of inputs that are qualitatively different from the distributions arising in practice.) 3. Run in polynomial time on all inputs but return sub-optimal solutions of cost “close” to the optimum. In this course we focus on approximation algorithms, that are algorithms of the third kind with a provably good worst-case ratio between the value of the solution found by the algorithm and the true optimum.
    [Show full text]
  • Revisiting Alphabet Reduction in Dinur's
    Revisiting Alphabet Reduction in Dinur’s PCP Venkatesan Guruswami Computer Science Department, Carnegie Mellon University, Pittsburgh, PA, USA [email protected] Jakub Opršal Computer Science Department, Durham University, UK [email protected] Sai Sandeep Computer Science Department, Carnegie Mellon University, Pittsburgh, PA, USA [email protected] Abstract Dinur’s celebrated proof of the PCP theorem alternates two main steps in several iterations: gap amplification to increase the soundness gap by a large constant factor (at the expense of much larger alphabet size), and a composition step that brings back the alphabet size to an absolute constant (at the expense of a fixed constant factor loss in the soundness gap). We note that the gap amplification can produce a Label Cover CSP. This allows us to reduce the alphabet size via a direct long-code based reduction from Label Cover to a Boolean CSP. Our composition step thus bypasses the concept of Assignment Testers from Dinur’s proof, and we believe it is more intuitive – it is just a gadget reduction. The analysis also uses only elementary facts (Parseval’s identity) about Fourier Transforms over the hypercube. 2012 ACM Subject Classification Theory of computation → Problems, reductions and completeness; Theory of computation → Approximation algorithms analysis Keywords and phrases PCP theorem, CSP, discrete Fourier analysis, label cover, long code Digital Object Identifier 10.4230/LIPIcs.APPROX/RANDOM.2020.34 Category APPROX Funding Venkatesan Guruswami: Supported in part by NSF grants CCF-1526092, CCF-1814603, and CCF-1908125. Jakub Opršal: Received funding from the UK EPSRC grant EP/R034516/1.
    [Show full text]
  • Lecture 9: Dinur's Proof of the PCP Theorem 9.1 Hardness Of
    CS369E: Expanders May 31, 2005 Lecture 9: Dinur’s Proof of the PCP Theorem Lecturer: Prahladh Harsha Scribe: Krishnaram Kenthapadi In this lecture, we will describe a recent and remarkable proof of the PCP theorem, due to Irit Dinur [Din]. This proof is a beautiful application of expanders for “soundness amplification” of randomized algorithms without using too many random bits. Before describing the proof, we will first look at the PCP theorem, by relating it with the theory of NP-completeness. 9.1 Hardness of Optimization Problems The theory of NP-completeness, as developed by Cook, Levin, and Karp, states that any language, L in NP is reducible to the Boolean satisfiability problem, 3SAT. By this, we mean that for every instance, x of the language L, we can obtain a satisfiability instance, φ such that x ∈ L if and only if φ is satisfiable. Thus, 3SAT is at least as hard as any other problem in NP. Karp further showed that 3SAT can be reduced to other problems such as CLIQUE and 3-COLORABILITY and hence that these problems are at least as hard as any problem in NP. In other words, solving these problems optimally is as hard as solving any other problem in NP optimally. However the question of the hardness of approximation was left open. For instance, can the following be true – finding a satisfying assignment for 3SAT is NP-hard, however it is easy to find an assignment that satisfies 99% of the clauses. Questions such as Other examples: can we approximate the clique size in a graph? Or, can we obtain a 3-coloring that satisfies 99% of the edge constraints? In other words, is the approximation version of some of NP-hard problems easier than the optimization versions.
    [Show full text]
  • PCP and Hardness of Approximation
    Chapter 19 PCP and Hardness of Approximation “...most problem reductions do not create or preserve such gaps...To create such a gap in the generic reduction (cf. Cook)...also seems doubtful. The intuitive reason is that compu- tation is an inherently unstable, non-robust mathematical object, in the the sense that it can be turned from non-accepting to ac- cepting by changes that would be insignificant in any reasonable metric.” Papadimitriou and Yannakakis, 1991 [?] The PCP Theorem provides an interesting new characterization for NP, as the set of languages that have a “locally testable” membership proof. It is reminiscent of —and was motivated by— results such as IP =PSPACE. Its essence is the following: Suppose somebody wants to convince you that a Boolean formula is satisfiable. He could present the usual certificate, namely, a satisfying as- signment, which you could then check by substituting back into the formula. However, doing this requires reading the entire certificate. The PCP The- orem shows an interesting alternative: this person can easily rewrite his certificate so you can verify it by probabilistically selecting a constant num- ber of locations—as low as 3 bits— to examine in it. Furthermore, this probabilistic verification has the following properties: (1) A correct certifi- cate will never fail to convince you (that is, no choice of your random coins will make you reject it) and (2) If the formula is unsatisfiable, then you are guaranteed to reject every claimed certificate with high probability. Web draft 2006-09-28 18:10 405 Complexity Theory: A Modern Approach. c 2006 Sanjeev Arora and Boaz Barak.
    [Show full text]
  • Lecture 25: Sketch of the PCP Theorem 1 Overview 2
    A Theorist's Toolkit (CMU 18-859T, Fall 2013) Lecture 25: Sketch of the PCP Theorem December 4, 2013 Lecturer: Ryan O'Donnell Scribe: Anonymous 1 Overview The complexity class NP is the set of all languages L such that if x 2 L, then there exists a polynomial length certificate that x 2 L which can be verified in polynomial time by a deterministic Turing machine (and if x2 = L, no such certificate exists). Suppose we relaxed the requirement that the verifier must be deterministic, and instead the verifier could use randomness to \approximately" check the validity of such a certificate. The notion of a probabilistically checkable proof (PCP) arises from this idea, where it can be shown that for any L 2 NP , a polynomial length certificate that x 2 L can be verified or disproven, with probability of failure say at most 1=2, by only looking at a constant number (independent of L) of bits in the certificate. More details are given in Section 2, where the existence of such a verifier is implied by the so-called PCP theorem; this theorem was first proven by [2, 3], but here we will cover a proof given by Dinur that is simpler [4]. Some of the more interesting applications of the PCP theorem involve showing hardness of approximability, e.g. it is NP-Hard to approximate Max-3SAT within a 7=8 factor, but we do not discuss such applications here. 2 PCP We give a statement of the PCP Theorem; for a more complete description, please see [4, 1].
    [Show full text]
  • The PCP Theorem by Gap Amplification
    The PCP Theorem by Gap Amplification Irit Dinur∗ February 13, 2007 Abstract The PCP theorem [3, 2] says that every language in NP has a witness format that can be checked probabilistically by reading only a constant number of bits from the proof. The celebrated equiva- lence of this theorem and inapproximability of certain optimization problems, due to [12], has placed the PCP theorem at the heart of the area of inapproximability. In this work we present a new proof of the PCP theorem that draws on this equivalence. We give a combinatorial proof for the NP-hardness of approximating a certain constraint satisfaction problem, which can then be reinterpreted to yield the PCP theorem. Our approach is to consider the unsat value of a constraint system, which is the smallest fraction of unsatisfied constraints, ranging over all possible assignments for the underlying variables. We de- scribe a new combinatorial amplification transformation that doubles the unsat-value of a constraint- system, with only a linear blowup in the size of the system. The amplification step causes an increase in alphabet-size that is corrected by a (standard) PCP composition step. Iterative application of these two steps yields a proof for the PCP theorem. The amplification lemma relies on a new notion of “graph powering” that can be applied to sys- tems of binary constraints. This powering amplifies the unsat-value of a constraint system provided that the underlying graph structure is an expander. We also extend our amplification lemma towards construction of assignment testers (alterna- tively, PCPs of Proximity) which are slightly stronger objects than PCPs.
    [Show full text]
  • Approximate Low-Weight Check Codes and Circuit Lower Bounds for Noisy Ground States
    Approximate low-weight check codes and circuit lower bounds for noisy ground states Chinmay Nirkhe, Umesh Vazirani, and Henry Yuen Electrical Engineering and Computer Sciences University of California, Berkeley fnirkhe, vazirani, [email protected] Abstract The No Low-Energy Trivial States (NLTS) conjecture of Freedman and Hastings (Quantum Information and Computation 2014), which asserts the existence of local Hamiltonians whose low-energy states cannot be generated by constant-depth quantum circuits, identifies a funda- mental obstacle to resolving the quantum PCP conjecture. Progress towards the NLTS conjec- ture was made by Eldar and Harrow (Foundations of Computer Science 2017), who proved a closely related theorem called No Low-Error Trivial States (NLETS). In this paper, we give a much simpler proof of the NLETS theorem and use the same technique to establish super- polynomial circuit size lower bounds for noisy ground states of local Hamiltonians (assuming QCMA 6= QMA), resolving an open question of Eldar and Harrow. We discuss the new light our results cast on the relationship between NLTS and NLETS. Finally, our techniques imply the existence of approximate quantum low-weight check (qLWC) codes with linear rate, linear distance, and constant weight checks. These codes are similar to quantum LDPC codes except (1) each particle may participate in a large number of checks, and (2) errors only need to be corrected up to fidelity 1 − 1/poly(n). This stands in contrast to the best-knownp stabilizer LDPC codes due to Freedman, Meyer, and Luo which achieve a distance of O( n log1/4 n). The principal technique used in our results is to leverage the Feynman-Kitaev clock con- struction to approximately embed a subspace of states defined by a circuit as the ground space of a local Hamiltonian.
    [Show full text]