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An Integrality-Gap Lower Bound of 1.2325 for Minimum Multiway Cut

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Abstract

We prove that the integrality gap of the Călinescu-Karloff-Rabani relaxation for Minimum Multiway Cut is at least 1.23251.2325 when the number of terminals is large. By a result of Manokaran, Naor, Raghavendra, and Schwartz (2008), it is therefore NP-hard, under the Unique Games Conjecture, to approximate Minimum Multiway Cut within a factor of 1.23251.2325. This improves the previously best known lower bound of 1.200161.20016 due to Bérczi, Chandrasekaran, Király, and Madan (2020).

Introduction

Multiway Cut is the problem of separating kk terminals in a weighted graph by removing edges of minimum total weight. We are given an undirected graph G=(V,E)G=(V,E) with nonnegative edge weights wuvw_{uv} and a set of kk terminals t1,…,tk∈Vt_1,\ldots,t_k\in V. A kk-way cut is a labeling C:V→[k]C:V\to[k] with C(ti)=iC(t_i)=i for i∈[k]i\in[k], and its cost is the weight of the edges whose endpoints receive different labels,

cost⁡(C)=∑uv∈Ewuv 1[C(u)≠C(v)].(1)\operatorname{cost}(C)=\sum_{uv\in E}w_{uv}\,1[C(u)\ne C(v)]. \tag*{(1)}

We write OPT for the minimum cost of a kk-way cut.

When k=2k=2, Multiway Cut is the Minimum ss–tt Cut problem, which is solvable in polynomial time. However, Dahlhaus, Johnson, Papadimitriou, Seymour, and Yannakakis [12, 13] showed that the problem is APX-hard for every fixed k≥3k\ge3. They also gave a simple combinatorial (2−2/k)(2-2/k)-approximation algorithm for the problem. A major advance was made by Călinescu, Karloff, and Rabani [10], who introduced a geometric approach to Multiway Cut based on linear programming. Their relaxation, now commonly known as the CKR relaxation, is defined as follows:

minimize∑uv∈Ewuvd(xu,xv)subject toxu∈Δkfor every u∈V,xti=eifor every i∈[k].(2)\begin{aligned} \text{minimize}\quad&\sum_{uv\in E}w_{uv}d(x^u,x^v)\\ \text{subject to}\quad&x^u\in\Delta_k &&\text{for every }u\in V,\\ &x^{t_i}=e_i &&\text{for every }i\in[k]. \tag*{(2)} \end{aligned}

Here Δk={x∈R≥0k:x1+⋯+xk=1}\Delta_k=\{x\in\mathbb{R}_{\ge0}^k:x_1+\cdots+x_k=1\} is the simplex, eie_i is the iith vertex, and we use the ℓ1\ell_1 metric scaled by 1/21/2, i.e., d(x,y)=12∥x−y∥1d(x,y)=\frac{1}{2}\lVert x-y\rVert_1. Every kk-way cut CC gives a feasible solution xu=eC(u)x^u=e_{C(u)} of the same cost, so the optimal value LP of (2) is at most OPT. Călinescu, Karloff, and Rabani rounded their relaxation to obtain a (3/2−1/k)(3/2 - 1/k)-approximation, and every approximation algorithm since rounds the same relaxation with a different rounding procedure [15, 6, 18, 4]. In particular, Karger, Klein, Stein, Thorup, and Young [15] settled the case of k=3k = 3 with approximation ratio 12/1112/11, and 1.34381.3438 for general kk. Buchbinder, Naor, and Schwartz [6] introduced the method of exponential clocks and designed a rounding scheme with ratio 1.32391.3239, and Sharma and Vondrák [18] obtain a 1.29651.2965-approximation with a mixture of rounding schemes. Buchbinder, Schwartz, and Weizman [8] gave a simpler rounding scheme nearly matching the approximation guarantee of [18]. The current best approximation ratio for large kk is 1.27871.2787 due to Brakensiek, Huang, Potechin, and Zwick [4].

Recall that the integrality gap of an instance is the ratio OPT/LP. Let GkG_k denote the supremum of the integrality gaps over all instances with kk terminals. Manokaran, Naor, Raghavendra, and Schwartz [17] showed that, assuming the Unique Games Conjecture, the best approximation ratio achievable in polynomial time for Minimum Multiway Cut with kk terminals equals GkG_k.

Compared to algorithmic results, lower bounds on the integrality gap are much less developed. The gap is known exactly only for k=3k = 3, where it is 12/1112/11 [15, 11]. For general kk, the known lower bounds analyze a low-dimensional skeleton of the high-dimensional simplex Δk\Delta_k. The dd-dimensional skeleton is the union of the faces of dimension at most dd, that is, the points of Δk\Delta_k with at most d+1d + 1 positive coordinates. Freund and Karloff [14] analyzed the one-dimensional skeleton and gave an asymptotic lower bound of 8/78/7. Angelidakis, Makarychev, and Manurangsi [2] introduced the technique of non-opposite cuts, which makes it possible to analyze a low-dimensional skeleton of Δk\Delta_k without considering the entire simplex, and improved the bound to 1.21.2 on the two-dimensional skeleton. The current best lower bound of 1.200161.20016 is due to Bérczi, Chandrasekaran, Király, and Madan [3].

Our contribution. We obtain the following asymptotic lower bound on the integrality gap for large kk.

Theorem 1.1. Let GkG_k denote the supremum of the integrality gap of the CKR relaxation (2) over all instances with kk terminals. Then, we have3

lim⁡k→∞Gk≥159389999561650025129321858575204877>1.2325.\lim_{k \to\infty} G_k \ge\frac{159389999561650025}{129321858575204877} > 1.2325.

Equivalently, for every η>0\eta> 0 and all sufficiently large kk, there is a finite kk-terminal instance whose integrality gap is at least the constant above minus η\eta.

This result combined with the hardness result of [17], immediately yields the following corollary.

Corollary 1.2. Assuming the Unique Games Conjecture, for every ϵ>0\epsilon> 0 and all sufficiently large kk, it is NP-hard to approximate Multiway Cut with kk terminals within a factor 1.2325−ϵ1.2325 - \epsilon.

The gap in Theorem 1.1 is in part obtained with a computer-assisted search via linear programming. However, we can get the slightly weaker gap of 1.223761.22376 without resorting to numerical optimization tools (Theorem 3.5). This is still better than the previous best lower bound of 1.200161.20016 [3].

Our approach. A rounding scheme for (2) is a random kk-way cut of the simplex itself, and by [15] the gap GkG_k is exactly the best approximation factor of such a cut, so a lower bound on the gap is a statement about every random kk-way cut of Δk\Delta_k. To remove the dependence on kk, we follow Angelidakis, Makarychev, and Manurangsi [2] and restrict the cut to a face with a fixed number qq of vertices. Up to an error that vanishes as kk grows, this turns it into a non-opposite cut of Δq\Delta_q. Such a cut has one extra label ∗\ast, which stands for all terminals outside the face, and every point receives a label from its support or ∗\ast. A lower bound on the approximation factor of every random non-opposite cut of Δq\Delta_q is then a lower bound on lim⁡k→∞Gk\lim_{k\to\infty} G_k.

On the faces of Δq\Delta_q we build a Multiway Cut instance together with label flows, one per label, fixed before the cut. A cut selects at every point the flow of the label it assigns there, and the flows satisfy a conservation and a capacity condition, as in max-flow min-cut. Conservation makes the flow of a label carry the weight of that label’s terminal edges, so a cut that keeps a terminal edge must let that flow into the region of its label, across a boundary where the label changes. Capacity makes the edges of the instance crossing that boundary weigh at least the difference of the two flows there.

Hence, a random non-opposite cut with finite approximation factor costs, in expectation, nearly as much as the cut that labels every point by ∗\ast, and the ratio of this cost to the fractional cost of the instance is a lower bound on its approximation factor (Theorem 3.4). We prove this for every qq, so each explicit bound in Theorems 1.1 and 3.5 is a single flow instance, on Δ11\Delta_{11} and on Δ10\Delta_{10}, respectively.

Roadmap. We introduce the notion of label flows and give an overview of our approach in § 3. In § 4 we exhibit a flow instance for q=3q=3. § 5 proves the lower bound formally. § 6 and § 7 give the two flow instances on Δ10\Delta_{10} and Δ11\Delta_{11}, respectively.

Preliminaries on Non-Opposite Cuts

In this section we give an overview of the non-opposite cuts of [2], and explain why we can restrict our attention to them.

A solution of (2) places each vertex uu at a point xu∈Δkx^u\in\Delta_k, and a random kk-way cut of the simplex, that is, a random labeling C:Δk→[k]C:\Delta_k\to[k] with C(ei)=iC(e_i)=i for all i∈[k]i\in[k], rounds it by labeling uu with C(xu)C(x^u). Such a cut has approximation factor α\alpha if

P[C(x)≠C(y)]≤αd(x,y)for all x,y∈Δk.(3)\mathbb{P}[C(x)\ne C(y)]\le\alpha d(x,y)\quad\text{for all }x,y\in\Delta_k. \tag*{(3)}

Then every edge uvuv is cut with probability at most αd(xu,xv)\alpha d(x^u,x^v), so the rounding is an α\alpha-approximation and Gk≤αG_k\le\alpha. Karger, Klein, Stein, Thorup, and Young [15] showed that some random kk-way cut of Δk\Delta_k has approximation factor GkG_k.

To remove the dependence on kk, Angelidakis, Makarychev, and Manurangsi [2 Lemma 13 of full version] restrict the cut to the face spanned by qq terminals chosen at random, which we identify with Δq\Delta_q, and replace the labels of all other terminals by a single label ∗\ast. We do not treat ∗\ast as another terminal, so a point labeled ∗\ast has all of its terminal edges cut. A point xx of the face rarely receives the label of a chosen terminal outside supp⁡(x)={i:xi>0}\operatorname{supp}(x)=\{i:x_i>0\}, so replacing such labels by ∗\ast changes the approximation factor by an error that vanishes as k→∞k\to\infty. The result is a cut of the following kind.

Definition 2.1 (Non-opposite cuts). A non-opposite cut of Δq\Delta_q is a labeling C:Δq→[q]∪{∗}C:\Delta_q\to[q]\cup\{\ast\} with C(ei)=iC(e_i)=i for every i∈[q]i\in[q] and C(x)∈supp⁡(x)∪{∗}C(x)\in\operatorname{supp}(x)\cup\{\ast\} for every x∈Δqx\in\Delta_q. A random non-opposite cut is a probability distribution over non-opposite cuts.

For example, the side x1=0x_1=0 of Δ3\Delta_3 allows the labels 22, 33, and ∗\ast but forbids 11, so no point of the side opposite e1e_1 receives its label. The approximation factor of a random non-opposite cut is defined by (3) with Δq\Delta_q in place of Δk\Delta_k. In the limit k→∞k \to\infty we obtain the following lemma, which we prove in Section A.

Lemma 2.2 (Non-opposite reduction). Fix q≥2q \ge2 and β>0\beta> 0. If every random non-opposite cut of Δq\Delta_q with approximation factor α\alpha satisfies α≥β\alpha\ge\beta, then lim⁡k→∞Gk≥β\lim_{k\to\infty} G_k \ge\beta, and for every η>0\eta> 0 and all sufficiently large kk there is a finite kk-terminal instance whose integrality gap is at least β−η\beta- \eta.

Technical Overview of the Approach

By Lemma 2.2, it suffices to fix qq and show that every random non-opposite cut of Δq\Delta_q has approximation factor at least β\beta. The argument below works for every qq, and Theorem 1.1 is the case q=11q = 11.

Let CC be a random non-opposite cut of Δq\Delta_q with approximation factor α\alpha, and consider a Multiway Cut instance whose vertices are points of Δq\Delta_q and whose terminals are the vertices e1,…,eqe_1,\ldots,e_q. Since CC severs an edge of length dd with probability at most αd\alpha d, the expected cost of CC on the instance is at most αL\alpha L, where LL is the cost of the fractional solution that places every vertex at its own point. Hence, if every such CC has expected cost at least MM on the instance, then α≥M/L\alpha\ge M/L. Our task is therefore to construct an instance together with a proof that every random non-opposite cut has expected cost at least MM on it, with M/LM/L as large as possible. We obtain such proofs by designing flows, one for each label, that are fixed before the cut, and we call an instance together with these flows a flow instance.

A flow instance determines three quantities: (1) LL is the fractional cost above, (2) BB is the cost of the cut that assigns the label ∗\ast to every point, and (3) RR is an allowance, a slack in the flow conditions. Our main technical result, Theorem 3.4, states that every random non-opposite cut of Δq\Delta_q with approximation factor α\alpha satisfies B−R≤E[cost⁡(C)]≤αLB-R \le\mathbb{E}[\operatorname{cost}(C)] \le\alpha L, and therefore α≥(B−R)/L\alpha\ge(B-R)/L. So, up to the allowance, no non-opposite cut is cheaper in expectation than labeling every point by ∗\ast. We first define flow instances on a finite graph (§ 3.1), where the argument cannot give a bound above one, because it also applies to the fractional solution. We then move to the faces of the simplex (§ 3.2), where it can. Finally, we state Theorem 3.4 and reduce it to three lemmas (§ 3.3), and describe the two flow instances behind Theorems 1.1 and 3.5 (§ 3.4).

Flow Instances on a Graph

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