Theoretical Computer Science

Motivation

Clique-width is a graph parameter introduced by Courcelle and Olariu (Discrete Appl. Math. 101 (2000)) that measures how far a graph is from being built by a few labelled operations. Every problem expressible in monadic second-order logic with quantification over vertices and vertex sets (MSO$_1$) can be solved in linear time on graphs given together with a decomposition of bounded clique-width (Courcelle, Makowsky and Rotics, Theory Comput. Syst. 33 (2000)). Bounded clique-width is more general than bounded tree-width: complete graphs have unbounded tree-width but clique-width $2$.

For fixed $k$ there was, before this paper, no polynomial-time algorithm that either decides that a graph has clique-width at least $k+1$ or outputs a decomposition of clique-width bounded by a function of $k$; the best known algorithm, by Johansson (2001), gave width $2k\log n$. Oum and Seymour (J. Combin. Theory Ser. B 96 (2006)) closed this gap with approximation $2^{3k+2}-1$, through rank-width and a factor-3 approximation for the branch-width of symmetric submodular functions.

Timeline:

• 1991: Robertson and Seymour introduce branch-width of graphs and hypergraphs (J. Combin. Theory Ser. B 52).
• 2000: Courcelle and Olariu define clique-width; Courcelle, Makowsky and Rotics solve MSO$_1$ problems on graphs given with a $k$-expression.
• 2001: Johansson gives a $2k\log n$ approximation.
• 2006: Oum and Seymour define rank-width, prove $\mathrm{rwd}(G) \le \mathrm{cwd}(G) \le 2^{\mathrm{rwd}(G)+1}-1$, and give an $O(n^9 \log n)$ algorithm that outputs a $(2^{3k+2}-1)$-expression or certifies clique-width above $k$.

Setting

All graphs are finite and simple. For a finite set $V$, a function $f : 2^V \to \mathbb{Z}$ is submodular if $f(X)+f(Y) \ge f(X\cap Y)+f(X\cup Y)$ and symmetric if $f(X) = f(V\setminus X)$.

A branch-decomposition of $f$ is a pair $(T, L)$ where $T$ is a tree with at least two vertices and all degrees at most $3$, and $L$ is a bijection from $V$ onto the leaves of $T$. Removing an edge $e$ of $T$ splits the leaves in two; the width of $e$ is $f$ of the set of elements of $V$ on one side. The width of $(T, L)$ is the largest edge width, and the branch-width $\mathrm{bw}(f)$ is the least width of a branch-decomposition, with $\mathrm{bw}(f) = f(\emptyset)$ when $|V| \le 1$.

A set $W \subseteq V$ is well-linked with respect to $f$ if for every partition $(X, Y)$ of $W$ and every $Z$ with $X \subseteq Z \subseteq V\setminus Y$, $f(Z) \ge \min(|X|, |Y|)$.

Let $A(G)$ be the adjacency matrix of $G$ over $\mathrm{GF}(2)$. For disjoint $X, Y \subseteq V(G)$, $\mathrm{cutrk}^*_G(X, Y)$ is the rank of the submatrix of $A(G)$ with rows $X$ and columns $Y$, and the cut-rank function is $\mathrm{cutrk}_G(X) = \mathrm{cutrk}^*_G(X, V(G)\setminus X)$. The rank-width $\mathrm{rwd}(G)$ is $\mathrm{bw}(\mathrm{cutrk}_G)$.

A $k$-expression is a term built from constants $\cdot_i$ (a vertex with label $i \in \{1,\dots,k\}$), the operators $\eta_{i,j}$ ($i \ne j$; add all edges between labels $i$ and $j$), $\rho_{i\to j}$ (relabel $i$ into $j$) and disjoint union $\oplus$. Its value is the labelled graph it produces; $G$ has clique-width $\mathrm{cwd}(G) \le k$ if some $k$-expression has value isomorphic to $G$.

An interpolation of $f$ is a function $f^*$ on disjoint pairs $(X, Y)$ that agrees with $f$ on $(X, V\setminus X)$, is monotone, submodular in the sense $f^*(A,B)+f^*(C,D) \ge f^*(A\cap C, B\cup D) + f^*(A\cup C, B\cap D)$, and has $f^*(\emptyset,\emptyset)=f(\emptyset)$.

Formalization targets

Goal: Theorem 1.1, certificate form

For a graph $G$ with at least one vertex and an integer $k \ge 1$:

$$\exists\, W,\ |W| = 3k+1,\ W \text{ well-linked for } \mathrm{cutrk}_G \;\Longrightarrow\; \mathrm{cwd}(G) \ge k+1,$$ $$\nexists\, W,\ |W| = 3k+1,\ W \text{ well-linked for } \mathrm{cutrk}_G \;\Longrightarrow\; \mathrm{cwd}(G) \le 2^{3k+2}-1.$$

The same explicit condition decides which side of the approximation holds; this is what the paper's algorithm certifies.

Milestones

1. Proposition 4.1: properties of an interpolation, including that $X \mapsto f^*(X, B) - f(\emptyset)$ is a matroid rank function on $V\setminus B$ when $f(\{v\}) - f(\emptyset) \le 1$.
2. Proposition 4.2: $f_{\min}(X,Y) = \min_{X\subseteq Z\subseteq V\setminus Y} f(Z)$ is an interpolation.
3. Theorem 5.1: a well-linked set of size $k$ forces $\mathrm{bw}(f) \ge k/3$ (for $k \ne 1$).
4. Theorem 5.2: no well-linked set of size $k$ implies $\mathrm{bw}(f) \le k$, when $f(\{v\}) \le 1$.
5. Proposition 6.1: $\mathrm{rk}\,M[X_1,Y_1] + \mathrm{rk}\,M[X_2,Y_2] \ge \mathrm{rk}\,M[X_1\cup X_2, Y_1\cap Y_2] + \mathrm{rk}\,M[X_1\cap X_2, Y_1\cup Y_2]$.
6. Corollary 6.2: submodularity of $\mathrm{cutrk}^*_G$ and $\mathrm{cutrk}_G$.
7. Section 6 claim: $\mathrm{cutrk}_G$ is symmetric submodular and $\mathrm{cutrk}^*_G$ interpolates it.
8. Proposition 6.3: $\mathrm{rwd}(G) \le \mathrm{cwd}(G) \le 2^{\mathrm{rwd}(G)+1}-1$.

Significance

The dichotomy turns clique-width, for which no exact polynomial algorithm is known even for fixed $k$, into a parameter that can be approximated with an explicit witness in each direction. Downstream, every algorithm for graphs of bounded clique-width that needs a $k$-expression as input becomes applicable to graphs given without one, at the cost of an exponential blow-up of the width.

The result is proved in the literature; this mission formalizes it. To our knowledge none of the objects involved — branch-width of set functions, rank-width, cut-rank, $k$-expressions, clique-width — has been formalized in Mathlib, and the submodularity of submatrix rank (Proposition 6.1) is absent from Mathlib's Matrix.rank API. The formal development would give reusable definitions of branch-decompositions of arbitrary integer set functions, of cut-rank, and of clique-width, and a machine-checked link between the combinatorial and the linear-algebraic width parameters.

Difficulty

The upper bound in Theorem 5.2 is the core. The natural approach, growing a branch-decomposition one leaf split at a time while keeping the width at most $k$, gets stuck at a leaf carrying a set $B$ with $f(B) = k$: a split of $B$ into two parts of $f$-value below $k$ has to be found, and it must be found from the failure of well-linkedness of a set that is not obviously related to $B$. The paper's device is the interpolation $f^*$, which attaches a matroid to $B$ whose base has exactly $f(B)$ elements. Formalizing this requires handling partial branch-decompositions, their extensions, and a maximality argument over trees, none of which exists in Mathlib.

Proposition 6.3's upper bound is a second, independent difficulty: a rank-decomposition must be converted into a $k$-expression by an induction over a rooted binary tree, with a relabelling argument bounding the number of labels by the number of distinct nonzero rows of a $\mathrm{GF}(2)$ matrix of rank $k$. Its lower bound needs the tree structure of a $k$-expression to be read as a branch-decomposition.

Formalization scope

The ground set is a Fintype V with DecidableEq V; subsets are Finset V; set functions are Finset V → ℤ, as in the paper. A branch-decomposition is a tree T : SimpleGraph (Fin n) with $n \ge 2$, all neighbour sets of size at most $3$, and an injective map $L$ from $V$ onto the vertices of degree $1$; the side of an edge $uw$ is found by reachability from $u$ after deleting $uw$. Branch-width, rank-width and clique-width are never computed as minima: "$\mathrm{bw}(f) \le k$" is the predicate "$|V| \le 1$ and $f(\emptyset) \le k$, or a branch-decomposition of width at most $k$ exists", lower bounds say that every branch-decomposition has a wide edge, and "$\mathrm{cwd}(G) \le k$" is "$G$ has a $k$-expression". Labels $\{1,\dots,k\}$ are Fin k. The value of a $k$-expression has as vertex type the occurrences of constants (a nested sum type), and $\eta_{i,j}$ requires $i \ne j$. Cut-rank uses Matrix.rank over ZMod 2 of submatrices of SimpleGraph.adjMatrix. An interpolation is a function on all pairs of subsets whose axioms are imposed on disjoint pairs only.

Running time is not formalized. The paper's Theorem 1.1 asserts an $O(n^9\log n)$ algorithm; there is no cost model on the page, and the goal states the certificate the algorithm returns instead. Without the running time, "$\mathrm{cwd}(G) \ge k+1$ or $\mathrm{cwd}(G) \le 2^{3k+2}-1$" holds for every graph, so that reading is ruled out as a formalization of the goal; so are well-linkedness with respect to anything other than $\mathrm{cutrk}_G$, widths defined by an unguarded infimum (which is $0$ on an empty family), $k$-expressions whose value is not the graph up to isomorphism or whose $\eta$ may join equal labels, and Theorem 5.1 stated for $k = 1$.

Correction of Theorem 5.1. As printed, Theorem 5.1 fails for $k = 1$: a singleton is always well-linked, but the edgeless graph on two vertices has cut-rank identically $0$ and branch-width $0 < 1/3$. The milestone carries the hypothesis $k \ne 1$; the goal uses the theorem only at size $3k+1 \ge 4$.

The graph with no vertex is excluded from the goal and from the upper bound of Proposition 6.3, since it has no $k$-expression for any $k$. Contributions welcome: proofs of the milestones, lemmas on branch-decompositions (suppressing degree-2 vertices, extending partial decompositions), and submatrix-rank submodularity, which is reusable beyond this mission.

Selected references

• S. Oum and P. Seymour, Approximating clique-width and branch-width, J. Combin. Theory Ser. B 96 (2006) 514–528. https://doi.org/10.1016/j.jctb.2005.10.006
• B. Courcelle and S. Olariu, Upper bounds to the clique width of graphs, Discrete Appl. Math. 101 (2000) 77–114. https://doi.org/10.1016/S0166-218X(99)00184-5
• B. Courcelle, J. A. Makowsky and U. Rotics, Linear time solvable optimization problems on graphs of bounded clique-width, Theory Comput. Syst. 33 (2000) 125–150. https://doi.org/10.1007/s002249910009
• N. Robertson and P. D. Seymour, Graph minors. X. Obstructions to tree-decomposition, J. Combin. Theory Ser. B 52 (1991) 153–190. https://doi.org/10.1016/0095-8956(91)90061-N

Motivation

David S. Johnson's 1974 paper Approximation Algorithms for Combinatorial Problems (J. Comput. System Sci. 9 (1974) 256–278) was one of the first systematic worst-case analyses of polynomial-time heuristics for NP-complete optimization problems. Its Section 5 proves the harmonic bound $\sum_{j=1}^k 1/j$ for the greedy algorithm on minimum-cardinality set cover, a result that still underlies the standard $\ln n$ approximation guarantee.

Section 6, the subject of this mission, asks what happens when the cost of a cover is its total size rather than its number of sets. This problem, SET COVERING II (EC), is the optimization version of the EXACT COVER recognition problem of Karp's list (Karp 1972): a family has a disjoint subcover exactly when the optimum equals the number of covered points. Johnson shows that the change of measure breaks the cardinality greedy but that a greedy rule based on an overlap ratio recovers essentially the same guarantee. The same accounting (paying for each newly covered point) later became the standard analysis of greedy weighted set cover (Chvátal 1979).

Setting

An input is a finite family $F = \{S_1, \dots, S_p\}$ of finite sets. Its covered set is $T = \bigcup_{S \in F} S$. A subcover is a subfamily $F' \subseteq F$ with $\bigcup_{S\in F'} S = T$, and its measure is

$$m_{EC}(F') = \sum_{S \in F'} |S|.$$

The optimum $F^*$ is the least measure of a subcover; since every subcover has measure at least $|T|$, an optimal subcover is one with the least possible overlapping. The subproblem EC$(k)$ admits only families in which every set has at most $k$ points.

Algorithm C2 keeps a subfamily SUB (initially empty), the unused sets LEFT (initially $F$) and the uncovered points UNCOV (initially $T$). While UNCOV is nonempty it chooses $S' \in$ LEFT minimizing

$$\mathrm{Ratio}(S) = \frac{|S - \mathrm{UNCOV}|}{|S \cap \mathrm{UNCOV}|},$$

the number of already-covered points of $S$ per newly covered point, and moves $S'$ from LEFT to SUB, removing its points from UNCOV. When several sets tie, any of them may be chosen; a subcover is choosable by C2 if some sequence of admissible choices returns it.

The overlap of a chosen set is $|S' - \mathrm{UNCOV}|$ at the moment it is chosen, and the cumulative overlap $\mathrm{OV}(F_1)$ of a run returning $F_1$ is the sum of these overlaps.

Formalization targets

Goal: Theorem 6 (p. 271)

For all $k \ge 1$ and $n > 0$,

$$R[C2, EC(k)](n) \le 1 + \ln(k) \le \sum_{j=1}^k \frac1j + \frac12,$$

and for all sufficiently large $n$, $R[C2, EC(k)](n) \ge \sum_{j=1}^k (1/j)$. In the size-free form used here, for every $k \ge 1$:

1. every subcover $M$ choosable by C2 on an input of EC$(k)$ satisfies $m_{EC}(M) \le (1 + \ln k)\,F^*$;
2. $1 + \ln k \le \sum_{j=1}^k 1/j + 1/2$;
3. some input of EC$(k)$ with $F^* > 0$ has a choosable subcover with $m_{EC}(M) \ge \big(\sum_{j=1}^k 1/j\big) F^*$.

Milestones (proof of Theorem 6, pp. 271–272)

• the measure of the output is $|T| + \mathrm{OV}(F_1)$;
• if C2 may choose a set with $\mathrm{Ratio}(S') \ge y$, then $(y+1)\,|\mathrm{UNCOV}| \le F^*$;
• with $a = F^*/|T|$ and $x = |T - \mathrm{UNCOV}|/|T|$, the next chosen set has $\mathrm{Ratio}(S') \le a/(1-x) - 1$;
• on EC$(k)$, $\mathrm{OV}(F_1) \le |T|\,(a[\ln(k) + 1] - 1)$;
• the analytic inequality $1 + \ln(k) \le \sum_{j=1}^k 1/j + 1/2$;
• the lower-bound input (Fig. 1 of the paper with every set of $F_1$ filled out to exactly $k$ points) on which C2 may pay $\sum_{j=1}^k 1/j$ times the optimum.

Significance

The result. The measure $\sum|S|$ penalizes overlap, and the paper notes (without proof, p. 270) that an algorithm returning an optimal cover for the cardinality measure can be a factor $k$ from optimal for this one. Theorem 6 shows that the ratio rule C2 is within $1 + \ln k$ of the least-overlap cover, and the lower bound shows that no analysis of C2 can beat $\sum_{j=1}^k 1/j$. The two bounds differ by less than $1/2$ for every $k$. The theorem was an early instance of a logarithmic guarantee for a weighted covering problem, where each set's cost is its size.

Formalizing it. The theorem has been proved since 1974; no machine-checked proof is known to exist. The mission produces a formal model of the EC problem and of C2 as a nondeterministic process, the overlap identity, and the discrete form of the paper's area-under-a-curve estimate. The last is the part the paper argues informally, through a step function and an integral.

Difficulty

The cardinality argument for C1 counts the sets chosen; here the sets have different sizes, so it does not apply. The overlap C2 pays per newly covered point is not bounded by a constant: early choices can be free and late ones cost up to $k-1$ per point, and the bound on the cumulative overlap must hold against the whole run, for every sequence of tie-breaks.

In a formal proof the integral must be replaced by a sum. Covered points arrive in blocks (one block per chosen set), the charge is constant on a block but the bound depends on the covered fraction at the start of the block, and the sum has to be compared with a logarithm. The terms $a\ln a$ and $a/k$ that the paper drops using $1 \le a \le k$ must be controlled as well, and the relation $1 \le a \le k$ must itself be proved from optimality. The lower bound needs an explicit run of C2 through ties on an input with $k \cdot k!$ points, checking at every stage that the intended set is a ratio minimizer.

Formalization scope

• Inputs. A family is p : ℕ with S : Fin p → Finset α (0-based, repetitions allowed; a repeated set counts twice in the measure if both copies are chosen, which C2 never does). Subcovers and SUB, LEFT are index sets. $F^*$ is a minimum over the finite, nonempty set of subcovers (Finset.inf'), never a junk value.
• Algorithm. C2 is a step relation on states (SUB, LEFT, UNCOV). The choice at Step 3 is existential over all minimizers, so every result quantifies over every choosable output (the paper's WORST). $\mathrm{Ratio}(S)$ is $+\infty$ when $S \cap \mathrm{UNCOV} = \emptyset$; the formal rule requires the chosen set to meet UNCOV and compares ratios by cross-multiplication, with no division.
• Overlap. The cumulative overlap depends on the run, not on the output alone, so it is carried by an inductive run relation RunOV.
• No problem size. The paper's $R[A, P](n)$ maximizes over inputs of size at most $n$ in an unspecified encoding. Upper bounds are stated for every input and every choosable output; the lower bound exhibits one input and one choosable output. Given monotonicity of $R$ in $n$, these are equivalent to the paper's claims. Ratios are stated multiplicatively, so $F^* = 0$ does not create a vacuous bound.
• Numbers. Measures are natural numbers cast to $\mathbb R$; $\ln$ is Real.log; $\sum_{j=1}^k 1/j$ is Mathlib's harmonic k.
• Ruled out. A deterministic tie-break would prove a weaker upper bound and could not realize the lower-bound run, and a ratio with $x/0 = 0$ would make disjoint-from-UNCOV sets the most attractive choice. The formalization uses neither.

The overlap identity and the discrete integral comparison are reusable for any greedy covering analysis that charges cost per newly covered point. Contributions welcome include proofs of the milestones, the invariants of the C2 run relation (SUB and LEFT partition the indices; UNCOV $= T - \bigcup$ SUB), and the explicit lower-bound run.

Selected references

• D. S. Johnson, Approximation algorithms for combinatorial problems, J. Comput. System Sci. 9 (1974), 256–278. https://doi.org/10.1016/S0022-0000(74)80044-9
• R. M. Karp, Reducibility among combinatorial problems, in Complexity of Computer Computations, Plenum, 1972, 85–103. https://doi.org/10.1007/978-1-4684-2001-2_9
• V. Chvátal, A greedy heuristic for the set-covering problem, Math. Oper. Res. 4 (1979), 233–235. https://doi.org/10.1287/moor.4.3.233

Motivation

Many problems in signal processing, statistics and function interpolation ask for a solution of a linear system $Ax\approx b$ that uses as few columns of $A$ as possible: a sparse approximate solution. Natarajan's 1995 paper Sparse Approximate Solutions to Linear Systems (SIAM J. Comput. 24(2):227–234) was motivated by radial basis interpolation, where each column corresponds to a basis function and fewer columns mean a cheaper interpolant. The paper does two things. It proves that finding the sparsest approximate solution is computationally hard (§2, Theorem 1), and it analyses a greedy column-selection algorithm whose number of chosen columns is within a factor, depending on the conditioning of $A$, of the optimum (§3, Theorem 2; a separate mission of this series).

The hardness theorem is the reason the second half of the paper exists: once exact minimization is ruled out, one settles for approximation guarantees. It is cited throughout the compressed-sensing literature as the canonical statement that $\ell_0$-minimization under an $\ell_2$ error constraint is NP-hard, and is the starting point for the later theory of when convex relaxations recover sparse solutions.

The argument follows the classical reduction from Exact Cover by 3-sets (X3C) to minimum-weight solutions of linear systems in Garey and Johnson (1979), pp. 221 and 246, adapted to an approximate right-hand side.

Setting

Sparse approximate solution (SAS). Given a matrix $A\in\mathbb R^{m\times n}$, a vector $b\in\mathbb R^m$ and a tolerance $\varepsilon>0$, find a vector $x\in\mathbb R^n$ with $\|Ax-b\|_2\le\varepsilon$ whose number of nonzero entries, written $\|x\|_0=|\{j : x_j\neq0\}|$, is as small as possible. Here $\|\cdot\|_2$ is the Euclidean norm.

Exact Cover by 3-sets (X3C). An instance is a ground set $S=\{s_1,\dots,s_m\}$ and a list $C=c_1,\dots,c_n$ of subsets of $S$, each with exactly three elements. An exact cover is a sub-collection $\hat C=\{c_j : j\in J\}$, $J\subseteq\{1,\dots,n\}$, such that every element of $S$ occurs in exactly one set of $\hat C$.

The transformation. From an X3C instance build the SAS instance with

• the incidence matrix $A\in\mathbb R^{m\times n}$: $A_{ij}=1$ if $s_i\in c_j$ and $A_{ij}=0$ otherwise, so column $j$ is the characteristic vector of $c_j$;
• the all-ones vector $b=(1,1,\dots,1)\in\mathbb R^m$;
• the tolerance $\varepsilon=\tfrac12$.

In Lean, $S$ is Fin m, the collection is C : Fin n → Finset (Fin m) with hC : ∀ j, (C j).card = 3, an exact cover is an index set J with IsExactCover C J, the matrix is incidence C, the vector $b$ is onesVec m, $Ax$ is Matrix.toEuclideanLin (incidence C) x, and $\|x\|_0$ is nnz x.

Formalization targets

Goal: correctness of the reduction

For every X3C instance $(S,C)$ as above,

$$\bigl(\exists J,\ \{c_j\}_{j\in J}\text{ is an exact cover of }S\bigr)\iff\bigl(\exists x\in\mathbb R^n,\ \|Ax-b\|_2\le\tfrac12\ \text{and}\ 3\,\|x\|_0\le m\bigr).$$

This is the sentence the proof of Theorem 1 (p. 228) establishes: "the constructed instance of SAS has a solution with $m/3$ or fewer entries if and only if the given instance of X3C has a solution."

Milestones

1. Forward direction. If $\{c_j\}_{j\in J}$ is an exact cover, the indicator vector $x=\mathbf 1_J$ satisfies $Ax=b$ and $3\|x\|_0=m$.
2. Entry bounds. For every $x$ with $\|Ax-b\|_2\le\frac12$, each entry of $Ax$ lies in $[\frac12,\frac32]$.
3. Lower bound on sparsity. For every such $x$, $m\le3\|x\|_0$.
4. Exact cover from a sparse solution. If moreover $3\|x\|_0\le m$, the sets $c_j$ with $x_j\neq0$ form an exact cover.

Significance

The result. The equivalence shows that deciding whether a sparse approximate solution with a prescribed number of nonzeros exists is at least as hard as X3C, which is NP-complete. Consequently no polynomial-time algorithm computes the optimum of SAS unless P = NP, and approximation algorithms such as the greedy method of §3 are the natural object of study. The same instance shows hardness persists for 0/1 matrices, a right-hand side of all ones and a constant tolerance, so the difficulty does not come from ill-conditioned data or from vanishing precision.

Formalizing it. The reduction is proved in the paper; this mission produces a machine-checked proof of its correctness, the combinatorial core of every NP-hardness claim for $\ell_0$-constrained least squares. No prior machine-checked version is known to exist, on the platform or elsewhere. The complexity-theoretic wrapper is out of scope (see below).

Difficulty

The forward direction is a direct computation. The converse contains the only real step, which the paper passes over with "it is clear". From $\|Ax-b\|_2\le\frac12$ one gets only that every entry of $Ax$ is in $[\frac12,\frac32]$; the entries of $x$ themselves are arbitrary reals, possibly negative or not equal to $1$, so $x$ need not be an indicator vector and $Ax=b$ need not hold. The exact-cover property must therefore be extracted from support sizes alone: every element is covered by some column in the support, the support has at most $m/3$ columns of three elements each, and a counting argument forces the chosen sets to be pairwise disjoint. Reading off a cover from the values of $x$ (for instance, taking the $j$ with $x_j=1$) does not work.

Formalization scope

• Vectors live in EuclideanSpace ℝ (Fin m) and EuclideanSpace ℝ (Fin n), so ‖·‖ is the paper's $\|\cdot\|_2$. Using the sup norm of Fin m → ℝ would give a different statement.
• The tolerance is exactly $\varepsilon=\frac12$, as printed.
• The collection is indexed, C : Fin n → Finset (Fin m): repeated sets are allowed and are distinct indices; an exact cover is a set of indices, and on the SAS side one nonzero entry is counted per index, so both sides treat duplicates consistently.
• "$m/3$ or fewer" is written $3\|x\|_0\le m$, never with natural-number division. With this form the equivalence holds for every $m$ (both sides are false when $3\nmid m$), which absorbs the paper's "without loss of generality $m$ is a multiple of 3"; no divisibility hypothesis is assumed. For $m=0$ both sides are true.
• The hypothesis that every set has exactly three elements is essential for the converse (with $m=9$, $O=\{s_3,\dots,s_9\}$, $c_1=\{s_1\}\cup O$, $c_2=\{s_2\}\cup O$, $c_3=O$, the vector $x=(1,1,-1)$ solves $Ax=b$ with $3\|x\|_0=m$, yet $C$ has no exact cover since $c_1$ and $c_2$ must both be chosen) and is kept as hC.
• Not formalized: the infinite-precision RAM machine model, polynomial-time many-one reductions, polynomial-time computability of the transformation (evident: an $m\times n$ 0/1 matrix), and the NP-completeness of X3C (cited by the paper from Garey–Johnson). The goal is therefore the correctness of the transformation, not a statement titled "SAS is NP-hard". A statement that only records the forward direction, or that fixes $x$ to be a 0/1 vector on the SAS side, would trivialize the converse and is not the target.
• Tools a solver will need are in Mathlib: coordinate bounds for the Euclidean norm (PiLp.norm_apply_le), Finset.card_biUnion_le, and Finset.card_biUnion for disjoint unions. Contributions of a reusable exact-cover API, or of a polynomial-time reduction framework that could later wrap this equivalence into an NP-hardness theorem, are welcome.

Selected references

• B. K. Natarajan, Sparse Approximate Solutions to Linear Systems, SIAM Journal on Computing 24(2):227–234, 1995. https://doi.org/10.1137/s0097539792240406
• M. R. Garey and D. S. Johnson, Computers and Intractability: A Guide to the Theory of NP-Completeness, W. H. Freeman, 1979 (X3C: problem [SP2], p. 221; minimum weight solution to linear equations: [MP5], p. 246).
• F. P. Preparata and M. I. Shamos, Computational Geometry: An Introduction, Springer, 1985 (the real RAM model). https://doi.org/10.1007/978-1-4612-1098-6

Motivation

The nearest common ancestor (nca) problem asks, for a fixed rooted tree and a sequence of vertex pairs $(v, w)$, for the deepest vertex that is an ancestor of both. It is a basic step in suffix-tree string algorithms and is equivalent to range-minimum queries (Bender, Farach-Colton, 2000). Harel and Tarjan, Fast Algorithms for Finding Nearest Common Ancestors, SIAM J. Comput. 13 (1984) 338–355, gave the first algorithm answering each query on a static tree in constant time on a random-access machine after linear preprocessing.

Their construction reduces the general problem to the case of a complete binary tree, where §3 of the paper shows that nca queries can be answered "by direct calculation" on vertex numbers: multiplication, division, powers of two, the base-two logarithm and bitwise exclusive or. The later simplification of Schieber and Vishkin (1988) is built on the same in-order numbering of a complete binary tree. This mission formalizes that arithmetic core.

Timeline, as reviewed in the paper's §1 (pp. 338–340):

• 1976: Aho, Hopcroft and Ullman (SIAM J. Comput. 5) give an $O(n + m\alpha(m+n, n))$-time off-line algorithm on a pointer machine, and for static trees a random-access algorithm with $O(n \log\log n)$ preprocessing and $O(\log\log n)$ time per query.
• 1976: van Leeuwen (unpublished report) gives an $O(n + m \log\log n)$-time algorithm for linking roots and static trees that runs on a pointer machine in $O(n)$ space.
• 1980: Harel (Proc. 21st FOCS) gives a preliminary version of the paper's results.
• 1984: Harel and Tarjan prove that pointer machines need $\Omega(\log\log n)$ time per query on static trees (Theorem 1), and give the $O(n)$-preprocessing, $O(1)$-query random-access algorithm whose base case is the subject of this mission.

Setting

Fix $d \ge 0$ and let $T$ be the complete binary tree of depth $d$. A vertex is identified with the path from the root to it, a word of at most $d$ left or right turns; the root is the empty word and $T$ has $n = 2^{d+1} - 1$ vertices. Following the paper's Appendix (pp. 354–355):

• $w$ is an ancestor of $v$ ($v$ a descendant of $w$) if the word $w$ is a prefix of the word $v$; every vertex is its own ancestor. $v$ and $w$ are unrelated if neither is an ancestor of the other.
• The depth of $v$ is its distance to the root; its height $h(v)$ is the length of the longest path from a leaf to $v$, which in $T$ is $d - \operatorname{depth}(v)$.
• $\operatorname{nca}(v, w)$ is the vertex of greatest depth that is an ancestor of both: the longest common prefix.

The vertices of $T$ are numbered from $1$ to $n$ in symmetric order (in-order): at every vertex, first the left subtree, then the vertex, then the right subtree. $\mathrm{sym}(v)$ is the number of $v$ and $\mathrm{sym}^{-1}(i)$ the vertex numbered $i$. For $d = 4$ (Fig. 1 of the paper) the root is $16$, its children $8$ and $24$, and the leaves $1, 3, 5, \dots, 31$. $i \oplus j$ denotes bitwise exclusive or and $\lg$ the base-two logarithm.

Two procedures of §3 use only numbers, heights and $d$:

• the nca depth algorithm: return $d - h(v)$ if $\mathrm{sym}(w) \in [\mathrm{sym}(v) - 2^{h(v)} + 1, \mathrm{sym}(v) + 2^{h(v)} - 1]$; else $d - h(w)$ if the same holds with $v, w$ exchanged; else $d - \lfloor \lg(\mathrm{sym}(v) \oplus \mathrm{sym}(w)) \rfloor$;
• the depth algorithm: given $v$ and a depth $d_2 \le \operatorname{depth}(v)$, with $h = d - d_2$, return $\mathrm{sym}^{-1}\bigl(2^{h+1}\lfloor \mathrm{sym}(v)/2^{h+1}\rfloor + 2^h\bigr)$.

Formalization targets

Goal: the nca algorithm is correct

The algorithm to compute $\operatorname{nca}(v,w)$ (p. 342) runs the nca depth algorithm to obtain $d_0$ and then the depth algorithm on $(v, d_0)$. The goal states that it returns the nearest common ancestor: for all vertices $v, w$ of $T$, with $d_0$ the output of the nca depth algorithm and $h = d - d_0$,

$$\mathrm{sym}(\operatorname{nca}(v,w)) = 2^{h+1}\left\lfloor \frac{\mathrm{sym}(v)}{2^{h+1}} \right\rfloor + 2^h .$$

Milestones

In the order the paper uses them:

1. Numbers at height $h$ (p. 341): the vertices of height $h$ are numbered $2^h, 3\cdot 2^h, 5\cdot 2^h, \dots$ from left to right.
2. Lemma 1: $h(v)$ is the largest $h$ with $2^h \mid \mathrm{sym}(v)$.
3. Lemma 2: the descendants of $v$ are the vertices numbered in $[\mathrm{sym}(v) - 2^{h(v)} + 1, \mathrm{sym}(v) + 2^{h(v)} - 1]$.
4. Lemma 3: for a height $h \ge h(v)$, the height-$h$ ancestor of $v$ has number $2^{h+1}\lfloor \mathrm{sym}(v)/2^{h+1}\rfloor + 2^h$.
5. Lemma 4: for unrelated $v, w$,

$$h(\operatorname{nca}(v,w)) = \lfloor \lg(\mathrm{sym}(v) \oplus \mathrm{sym}(w)) \rfloor .$$

6. The nca depth algorithm returns $\operatorname{depth}(\operatorname{nca}(v,w))$.
7. The depth algorithm returns the number of the depth-$d_2$ ancestor of $v$.

Two supporting statements pin the definitions to the paper: $\mathrm{sym}$ is a bijection onto $\{1, \dots, 2^{d+1} - 1\}$, and the longest common prefix is the deepest common ancestor.

Significance

The constant-time nca computation on complete binary trees is the base case of the whole paper: §§4–5 embed an arbitrary tree into a moderately sized complete binary tree through a compressed tree and a balanced binary tree, and every query ends with the arithmetic of §3. The same idea, that in-order numbers encode ancestry in their low-order bits, underlies the Schieber–Vishkin algorithm. Lemma 1 identifies the height with the 2-adic valuation of the number, and Lemma 4 identifies the nca height with the position of the highest differing bit.

The results are proved in the paper, with the proofs left as "easy to verify". No machine-checked version of this numbering or of these four lemmas is known to exist in Mathlib or on this platform. A formal development supplies proofs of the four lemmas and the two algorithms, and a reusable library connecting in-order ranks of a complete binary tree to binary arithmetic (Nat.log, bitwise xor, 2-adic valuation).

Difficulty

The numbering is defined by a traversal order, while the lemmas speak about divisibility, floor division and exclusive or. The work lies in connecting the rank of a vertex in symmetric order to its closed form $(2j+1)\cdot 2^{h(v)}$, where $j$ is its left-to-right position. That counting argument sums the sizes of the subtrees that precede $v$ and is where most of the effort goes. Lemma 4 then needs the observation that two unrelated numbers agree in all bits above the height of their nca and differ in the bit at that height. This is a statement about Nat.testBit of the exclusive or, and it fails for related vertices. The algorithm statements add a case analysis whose first two cases overlap when $v = w$.

Formalization scope

• A vertex of the tree of depth $d$ is a List Bool of length at most $d$ (false = left). Ancestry is the prefix relation, $\operatorname{nca}$ the longest common prefix, depth the length, and height $d - \text{length}$. None of these structural notions uses the numbering.
• $\mathrm{sym}(v)$ is the number of vertices whose in-order sort key is lexicographically at most that of $v$. The key is the path with left $\mapsto 0$, right $\mapsto 2$, followed by $1$. The numbering is not defined by the closed form or by a recursion on numbers: a definition of that kind would make the height-$h$ numbering and Lemma 1 immediate and move the content of the mission into an uncheckable definition.
• $\lfloor \lg x \rfloor$ is Nat.log 2 x, which agrees for $x \ge 1$. $\oplus$ is ^^^ on $\mathbb N$, and floor division is / on $\mathbb N$.
• Interval tests $a \in [b - c + 1, b + c - 1]$ are written additively as $b + 1 \le a + c$ and $a + 1 \le b + c$. The subtractions $d - h(v)$ and $d - d_2$ never truncate for heights and depths of vertices.
• Lemma 3 states explicitly that $h \le d$ ("$h$ is a height") and that the ancestor exists. The depth algorithm assumes $d_2 \le \operatorname{depth}(v)$, as printed.
• $\mathrm{sym}^{-1}$ is not defined as a function. The goal and the depth algorithm state that a vertex has the computed number if and only if it is the nearest common ancestor (respectively the ancestor at depth $d_2$), which says that $\mathrm{sym}^{-1}$ of that number is that vertex.
• The $O(1)$ time bounds are not formalized, since the random-access machine model is out of scope.

Proofs of any milestone are welcome.

Selected references

• D. Harel and R. E. Tarjan, Fast Algorithms for Finding Nearest Common Ancestors, SIAM J. Comput. 13(2) (1984), 338–355. https://doi.org/10.1137/0213024
• A. V. Aho, J. E. Hopcroft and J. D. Ullman, On Finding Lowest Common Ancestors in Trees, SIAM J. Comput. 5(1) (1976), 115–132. https://doi.org/10.1137/0205011
• B. Schieber and U. Vishkin, On Finding Lowest Common Ancestors: Simplification and Parallelization, SIAM J. Comput. 17(6) (1988), 1253–1262. https://doi.org/10.1137/0217079
• M. A. Bender and M. Farach-Colton, The LCA Problem Revisited, LATIN 2000, LNCS 1776, 88–94. https://doi.org/10.1007/10719839_9

Motivation

The nearest common ancestor problem asks, for a rooted tree and two of its vertices $x$ and $y$, for the deepest vertex that is an ancestor of both, written $\operatorname{nca}(x,y)$. It appears as a subroutine in string algorithms (suffix trees), in graph algorithms (path queries, dominators) and in the analysis of set-union structures. Aho, Hopcroft and Ullman (On finding lowest common ancestors in trees, SIAM J. Comput. 5, 1976) posed it in several versions, differing in how much the tree changes while the queries are answered.

Harel and Tarjan (Fast Algorithms for Finding Nearest Common Ancestors, SIAM J. Comput. 13, 1984) study how the answer depends on the machine model. On a random-access machine, where addresses can be computed arithmetically, they preprocess a static tree in linear time and then answer each query in constant time. On a pointer machine, where memory can only be traversed by following pointers, their §2 shows that no representation of the tree allows constant-time queries: $\Omega(\log\log n)$ steps are needed in the worst case. This mission formalizes that lower bound.

Timeline.

• 1976: Aho, Hopcroft and Ullman give an $O(\log\log n)$-per-query random-access algorithm for static trees.
• 1976: van Leeuwen (Finding lowest common ancestors in less than logarithmic time, unpublished report, reference [14] of Harel–Tarjan) gives an $O(n + m\log\log n)$ algorithm for static trees that runs on a pointer machine.
• 1984: Harel and Tarjan prove Theorem 1, the matching $\Omega(\log\log n)$ lower bound for pointer machines, and the $O(1)$-per-query random-access algorithm.

Setting

A pointer machine stores its data as a collection of nodes. Each node has a fixed number of fields, and a pointer field holds either a node or nil. The machine can follow a pointer from a node it holds, but it cannot compute an address. Following Harel and Tarjan (p. 340), a static tree is represented by a list structure: each tree vertex $v$ is represented by a single node $\mathrm{rep}(v)$, distinct vertices by distinct nodes, and the structure may contain further nodes that represent no vertex. Each node has two pointer fields; the paper reduces any fixed number of pointers to two "without loss of generality". To answer a query on $x$ and $y$, the machine is given pointers to $\mathrm{rep}(x)$ and $\mathrm{rep}(y)$ and must return a pointer to $\mathrm{rep}(\operatorname{nca}(x,y))$.

The node $b$ is accessible from $a$ in $j$ steps or less if it can be reached from $a$ by following at most $j$ pointers. Write $\mathrm{acc}_j(a)$ for the set of such nodes. A run of $t$ steps from input nodes $a$ and $b$ is a sequence $n_1,\dots,n_t$ in which each $n_s$ is the content of a pointer field of a node among $a, b, n_1, \dots, n_{s-1}$. A query with answer $c$ is answered in $k$ steps if some run of at most $k$ steps holds $c$.

The tree is the complete binary tree $T$ of height $h$, with $n = 2^h$ leaves. Its vertices are the words $w \in \{0,1\}^{\le h}$ (the root-to-vertex path, $0$ = left), the ancestors of $v$ are its prefixes, the depth of $w$ is $|w|$ and its height is $h - |w|$. Then $\operatorname{nca}(x,y)$ is the longest common prefix of $x$ and $y$. Logarithms are binary: $\lg = \log_2$.

Formalization targets

Goal: Theorem 1 in the explicit form of its proof

For every $h$, every node type, every list structure with two pointers per node and every injective representation $\mathrm{rep}$ of the complete binary tree with $n = 2^h$ leaves: if every nca query on two leaves is answered in $k$ steps, then

$$k > \lg\lg n - 2 .$$

This is the last display of the proof (p. 341), which is what the paper's $\Omega(\log\log n)$ means. The representation is arbitrary and is quantified before the query bound, so the bound holds for every representation.

Milestones: the claims of the proof

1. A query answered in $k$ steps reaches only nodes in $\mathrm{acc}_k(\mathrm{rep}(x)) \cup \mathrm{acc}_k(\mathrm{rep}(y))$.
2. $|\mathrm{acc}_j(a)| \le 2^{j+1} - 1$ for every node $a$.
3. With $A_x$ the set of vertices whose nodes are accessible from $\mathrm{rep}(x)$ in $k$ steps or less: for a nonleaf $w$ with children $u, v$, either $w \in A_x$ for every leaf $x$ below $u$, or $w \in A_y$ for every leaf $y$ below $v$.
4. A vertex of height $i \ge 1$ lies in $A_x$ for at least $2^{i-1}$ leaves $x$.

$$\sum_{x \in L} |A_x| \ge \frac{n}{2}\lg n,$$

where $L$ is the set of leaves.

Significance

The result. Theorem 1 shows that van Leeuwen's pointer-machine algorithm for static trees is optimal up to a constant factor, and that the constant-time queries of the paper's §§3–5 depend on address arithmetic. It is an early nontrivial lower bound for pointer machines on a natural problem; the paper compares it with Tarjan's lower bound for disjoint-set union on a pointer machine (J. Comput. System Sci. 18, 1979).

Formalizing it. The theorem is proved in the paper; as far as could be determined no machine-checked version exists, and Mathlib has no pointer-machine model. The mission produces an explicit, reusable definition of pointer-machine runs and accessibility together with a complete proof of the explicit bound. A formal model of this kind is the precondition for stating any other pointer-machine lower bound.

Difficulty

The statement must hold for every representation, including structures with many auxiliary nodes and arbitrary pointers between tree nodes. Arguing about one natural representation, such as parent pointers, where a leaf is far from its ancestors, says nothing about other representations: a structure with shortcut pointers or auxiliary nodes may bring some ancestors close to some leaves, and the bound must survive every such choice. In the formal setting the counting also has to handle overlaps: nodes reachable from several leaves, nodes that represent no vertex, and pointer cycles.

Formalization scope

• Model. Nodes form an arbitrary type N, not necessarily finite. ptr : N → Fin 2 → Option N gives the two pointer fields (none = nil), and rep : Vertex h → N is required to be injective. acc ptr j a is defined recursively. Run ptr a b t held is an inductive predicate for runs of $t$ steps, AnsweredIn asks for some run of at most $k$ steps holding the answer, and AnswersLeafQueriesIn ptr rep k requires this for every pair of leaves.
• Conventions.
  • Two pointer fields per node, as the paper's "without loss of generality" reduction allows; the reduction itself is not formalized.
  • Only queries on two leaves are assumed answerable. This is weaker than all queries, so the theorem is at least as strong as the paper's.
  • Time is counted as pointer-following steps. Mutation of the structure during a query and non-pointer fields are not modelled: neither lets the machine hold a node it has not reached by following pointers. The clause "the algorithm remembers nothing between queries" is built into the static structure.
  • Vertex h is {s : List Bool // s.length ≤ h}, nca is the longest common prefix, and a separate theorem identifies it with the Appendix's deepest common ancestor. $n = 2^h$ counts leaves, not vertices.
  • $\lg$ is Real.logb 2. For $h = 0$ Lean's $\log_2 0 = 0$ gives the true statement $k > -2$; for $h \ge 1$, $\lg\lg n = \log_2 h$.
  • Cardinalities in the milestones are Set.encard in $\mathbb N \cup \{\infty\}$, so finiteness is part of each claim. Divisions are cleared: $h\,2^h \le 2\sum_x |A_x|$.
• Ruling out trivial formalizations. The hypothesis AnswersLeafQueriesIn is satisfiable: the parent-pointer representation answers every leaf query in $h$ steps. If rep were not injective, a constant rep would answer every query in zero steps, so injectivity is kept in the goal. The milestones do not need it and do not assume it.
• Infrastructure. The goal needs finite-set counting over the leaves of the complete binary tree and a double count over heights. The run and accessibility definitions are reusable for other pointer-machine arguments. Proofs of the milestones, and of the ℕ form $h < 2^{k+2}$ that the goal reduces to, are welcome.

Selected references

• D. Harel, R. E. Tarjan, Fast Algorithms for Finding Nearest Common Ancestors, SIAM J. Comput. 13(2):338–355, 1984. https://doi.org/10.1137/0213024
• A. V. Aho, J. E. Hopcroft, J. D. Ullman, On finding lowest common ancestors in trees, SIAM J. Comput. 5(1):115–132, 1976. https://doi.org/10.1137/0205011
• A. Schönhage, Storage modification machines, SIAM J. Comput. 9(3):490–508, 1980. https://doi.org/10.1137/0209036
• R. E. Tarjan, A class of algorithms which require nonlinear time to maintain disjoint sets, J. Comput. System Sci. 18(2):110–127, 1979. https://doi.org/10.1016/0022-0000(79)90042-4

Motivation

The maximum-flow problem asks how much of a commodity can be sent from a source to a sink through a network whose edges carry capacities. It is a basic model of operations research. Transportation, scheduling, bipartite matching and image segmentation reduce to it, and it is the inner step of many combinatorial algorithms.

Goldberg and Tarjan introduced the push–relabel (preflow) method in A New Approach to the Maximum-Flow Problem (J. ACM 35(4), 1988). Ford–Fulkerson-type algorithms augment along whole source–sink paths. The push–relabel method instead moves excess flow across single edges, guided by integer distance labels on the vertices. Whatever order its local operations are applied in, it is correct and performs $O(n^2 m)$ of them (§3 of the paper). Section 4 shows that one particular order, processing the active vertices first-in, first-out, cuts the dominant term, the number of nonsaturating pushes, to $O(n^3)$. The method and its FIFO and highest-label variants are the standard practical maximum-flow codes.

Timeline:

• 1956: Ford and Fulkerson, augmenting paths and max-flow min-cut.
• 1970–72: Dinic, and Edmonds and Karp, give polynomial augmenting-path bounds.
• 1974: Karzanov introduces preflows and obtains $O(n^3)$.
• 1982: Shiloach and Vishkin give a parallel $O(n^2 \log n)$ preflow algorithm with a first-in, first-out flavour.
• 1988: Goldberg and Tarjan, the generic push–relabel method, the FIFO bound of this mission, and $O(nm \log(n^2/m))$ with dynamic trees.

Setting

A flow network has a finite vertex set $V$ with $n = |V|$, a capacity $c(v,w) \ge 0$ on every ordered pair, a source $s$ and a sink $t \neq s$. The edges are the pairs with $c(v,w) > 0$, and there are no loops. A preflow is a function $f$ on vertex pairs with $f(v,w) \le c(v,w)$ and $f(v,w) = -f(w,v)$. Its excess $e(v) = \sum_u f(u,v)$ must be nonnegative at every $v \neq s$. The residual capacity is $r_f(v,w) = c(v,w) - f(v,w)$. A labeling $d$ assigns each vertex a value in $\mathbb{N} \cup \{\infty\}$. A vertex $v \notin \{s,t\}$ is active if $d(v) < \infty$ and $e(v) > 0$.

The two basic operations (Fig. 1 of the paper) are:

• push$(v,w)$, applicable when $v$ is active, $r_f(v,w) > 0$ and $d(v) = d(w)+1$. It sends $\delta = \min(e(v), r_f(v,w))$ from $v$ to $w$. The push is saturating if $r_f(v,w) = 0$ afterwards and nonsaturating otherwise.
• relabel$(v)$, applicable when $v$ is active and $d(v) \le d(w)$ for every residual edge $(v,w)$. It sets $d(v) \leftarrow \min\{d(w)+1 : r_f(v,w) > 0\}$.

The algorithm starts by saturating every edge leaving $s$, with $d(s) = n$ and $d(v) = 0$ for $v \neq s$.

In the first-in, first-out algorithm (§4), each vertex $v$ scans a fixed list $L(v)$ of its neighbours through a current edge. The push/relabel$(v)$ operation pushes through the current edge if possible. Otherwise it advances the current edge, or, at the end of the list, returns to the first edge and relabels $v$. Active vertices wait in a queue $Q$, initially $\{v \in V - \{s,t\} : c(s,v) > 0\}$. The discharge operation removes the front vertex $v$ and repeats push/relabel$(v)$ until $e(v) = 0$ or $d(v)$ increases. Every vertex that becomes active meanwhile is appended to $Q$, and $v$ is appended too if it is still active. Passes over the queue are defined inductively. Pass 1 consists of the discharges of the initially queued vertices. Pass $i+1$ consists of the discharges of vertices added during pass $i$.

Formalization targets

Goal: Corollary 4.4 (p. 931)

For every network, every edge-list order, every initial queue order, and every run of the FIFO algorithm,

$$\#\{\text{nonsaturating pushes}\} \le 4n^3 .$$

The constant is the printed one.

Milestones

• Lemma 4.1 (p. 929): the push/relabel operation relabels only when relabeling is applicable.
• Lemma 3.5 (p. 926): from any vertex with positive excess, the source is reachable in the residual graph.
• Lemma 3.7 (p. 927): at any time, $d(v) \le 2n-1$ for every vertex.
• Lemma 3.8 (p. 927): at most $2n-1$ relabelings per vertex and at most $(2n-1)(n-2) < 2n^2$ in total.
• Lemma 4.3 (p. 930): at most $4n^2$ passes over the queue.

Significance

Corollary 4.4 is the combinatorial core of Theorem 4.5, which states that the FIFO algorithm runs in $O(n^3)$ time. Theorem 4.2 shows that the remaining work of the implementation is $O(nm)$ plus constant time per nonsaturating push. The bound of Corollary 4.4 is therefore what separates the $O(n^3)$ FIFO method from the $O(n^2 m)$ bound of the generic method, which matters on dense networks. The same pass-counting argument is reused for the parallel algorithm of §6 and underlies later analyses of highest-label and wave variants.

The results are proved in the paper. Formalizing them adds an analysis of a push–relabel algorithm, which the platform does not yet have. Its existing network-flow material states max-flow min-cut and Ford–Fulkerson termination in an arc-based model with nonnegative flows (the Introduction to Linear Optimization missions). The mission builds a precise operational model of the FIFO implementation, with edge lists, current edges and a queue carrying pass numbers, and states an explicit operation count for it. A companion mission in this series treats the generic algorithm's correctness and its $(2n-1)(n-2) + 2nm + 4n^2m$ operation bound.

Difficulty

The obvious argument is the potential-function count of §3, over the sum of the labels of active vertices. It yields only $4n^2 m$ and does not use the queue discipline at all. The $4n^3$ bound has to charge nonsaturating pushes to passes over the queue, and then bound the number of passes by the total growth of the labels. Neither step is visible in the generic algorithm, because both depend on the order in which vertices are processed.

Making this rigorous requires invariants of the implementation that the paper uses silently:

• a vertex is in the queue exactly when it is active, and at most once;
• pass numbers are nondecreasing along the queue;
• current edges only move forward between relabelings.

Lemma 4.1 in particular depends on the current-edge scan: an edge passed over earlier stays inadmissible until $v$ is relabeled.

Formalization scope

The Lean development works in namespace GoldbergTarjan.FIFO. Vertices form a type V with [Fintype V] [DecidableEq V], and $n$ is Fintype.card V. Capacities are c : V → V → ℝ with c ≥ 0 and c v v = 0. Flows are antisymmetric real functions on all ordered pairs, not nonnegative arc flows. Excess is computed from the flow. Labels are in ℕ∞, and the empty minimum in relabel is ⊤.

The state of the algorithm consists of the flow, the labels, the current-edge index cur v into the edge list L v, and the queue Q : List (V × ℕ), each entry tagged with its pass number. Push/relabel (Fig. 3) is a total function, and a discharge (Fig. 4) is a relation carrying the number of push/relabel operations it performs. A run consists of the states S 0, …, S K with S 0 the initial state and consecutive states related by one discharge. The printed variant of Fig. 4, which stops as soon as $v$ is relabeled, is the one formalized. Counts are natural numbers over all push/relabel operations of all discharges. The number of passes is the largest pass tag of a discharged entry.

All constants are explicit, exactly as printed:

• $2n-1$ (Lemmas 3.7, 3.8);
• $(2n-1)(n-2)$ and $2n^2$ (Lemma 3.8);
• $4n^2$ (Lemma 4.3);
• $4n^3$ (Corollary 4.4).

No asymptotic notation is used, and no $m \ge n-1$ assumption is made.

A model without current edges, where relabeling happens whenever no push applies, would make Lemma 4.1 vacuous and change the algorithm. Pass numbers that are not propagated by the "added during pass $i$" rule would make the pass count arbitrary. Both are ruled out by the definitions. A sorry-free check, outside the proposal, exhibits a three-vertex network with two legal discharges, two passes and no nonsaturating push, so the run hypotheses are satisfiable.

Reusable beyond this mission are the network, preflow, push and relabel definitions and Lemma 3.5, which is about an arbitrary preflow. Contributions welcome: invariants of FIFO runs (preflow, valid labeling, queue = active set, cur within bounds), proofs of the milestones, and the reduction of Corollary 4.4 to Lemma 4.3.

Selected references

• A. V. Goldberg, R. E. Tarjan, A New Approach to the Maximum-Flow Problem, Journal of the ACM 35(4):921–940, 1988. https://doi.org/10.1145/48014.61051
• A. V. Karzanov, Determining the maximal flow in a network by the method of preflows, Soviet Math. Doklady 15:434–437, 1974.
• Y. Shiloach, U. Vishkin, An O(n² log n) parallel max-flow algorithm, Journal of Algorithms 3(2):128–146, 1982. https://doi.org/10.1016/0196-6774(82)90013-X
• L. R. Ford, D. R. Fulkerson, Maximal flow through a network, Canadian Journal of Mathematics 8:399–404, 1956. https://doi.org/10.4153/CJM-1956-045-5
• J. Edmonds, R. M. Karp, Theoretical improvements in algorithmic efficiency for network flow problems, Journal of the ACM 19(2):248–264, 1972. https://doi.org/10.1145/321694.321699

Motivation

The maximum-flow problem asks how much of a commodity can be sent from a source to a sink through a network whose edges have capacities. It is a basic model in operations research (transportation, scheduling, bipartite matching) and a standard subroutine in combinatorial optimization.

Classical algorithms, from Ford and Fulkerson (1956) through Edmonds–Karp and Dinic (1970–1972) and Karzanov (1974), increase a feasible flow along augmenting paths or blocking flows. Goldberg and Tarjan, A New Approach to the Maximum-Flow Problem (J. ACM 35(4), 1988, doi:10.1145/48014.61051), replaced this global view by a local one: the push-relabel method maintains a preflow, which may violate conservation at intermediate vertices, and moves excess along edges toward vertices with smaller distance labels. The generic method, with the basic operations applied in any order, is the starting point of the FIFO, highest-label and dynamic-tree implementations analysed later in the same paper, and push-relabel codes remain among the fastest practical maximum-flow solvers.

This mission formalizes §2–§3 of the paper: the generic algorithm is correct, and it stops after a number of basic operations bounded by an explicit polynomial in the numbers of vertices and edges, whatever order of operations is chosen.

Setting

A flow network has a finite vertex set $V$ with $n = |V|$, a source $s$ and a sink $t \ne s$, and a capacity $c(v,w) \ge 0$ for every ordered pair of vertices, positive exactly on the edges $E = \{(v,w) : c(v,w) > 0\}$; $m = |E|$, and there are no loops, $c(v,v) = 0$.

Flows are real functions on all vertex pairs. A function $f$ satisfies the capacity constraint if $f(v,w) \le c(v,w)$ and antisymmetry if $f(v,w) = -f(w,v)$ for all pairs. The excess of $v$ is $e(v) = \sum_{u} f(u,v)$. A flow also has $e(v) = 0$ for $v \notin \{s,t\}$; a preflow only $e(v) \ge 0$ for $v \ne s$. The value of a flow is $|f| = \sum_v f(v,t)$, and a maximum flow is a flow of maximum value.

The residual capacity is $r_f(v,w) = c(v,w) - f(v,w)$; pairs with $r_f(v,w) > 0$ are the edges of the residual graph $G_f$. A valid labeling is $d : V \to \mathbb{N} \cup \{\infty\}$ with $d(s) = n$, $d(t) = 0$ and $d(v) \le d(w) + 1$ on every residual edge. A vertex $v$ is active if $v \notin \{s,t\}$, $d(v) < \infty$ and $e(v) > 0$.

The two basic operations (Fig. 1 of the paper) are:

• Push$(v,w)$, applicable when $v$ is active, $r_f(v,w) > 0$ and $d(v) = d(w)+1$: send $\delta = \min(e(v), r_f(v,w))$, i.e. $f(v,w) \mathrel{+}= \delta$, $f(w,v) \mathrel{-}= \delta$. It is saturating if $r_f(v,w) = 0$ afterwards and nonsaturating otherwise.
• Relabel$(v)$, applicable when $v$ is active and $d(v) \le d(w)$ for every residual edge $(v,w)$: set $d(v) \leftarrow \min\{d(w)+1 : (v,w) \in E_f\}$ ($\infty$ if there is none).

The generic algorithm (Fig. 2) starts from the preflow that saturates every edge leaving $s$ and is zero elsewhere, with the simple labeling $d(s) = n$, $d(v) = 0$ otherwise, and applies applicable basic operations in any order while one exists. An execution with $K$ basic operations is a sequence of states $(f_0,d_0),\dots,(f_K,d_K)$ from the initial state, each obtained from the previous one by one applicable operation.

Formalization targets

Goal: Theorems 3.11 and 3.4

Assume the paper's standing assumption $m \ge n-1$. For every execution with $K$ basic operations,

$$K \le (2n-1)(n-2) + 2nm + 4n^2 m,$$

and if no basic operation applies in the final state, then $f_K$ is a maximum flow. The paper states the bound as $O(n^2m)$ and proves it as "immediate from Lemmas 3.8, 3.9, and 3.10"; the goal states the sum of those three printed bounds. Since every execution is this short, no order of operations runs forever.

Milestones

In the order the proof uses them: Lemma 2.1 (at an active vertex a push or a relabel applies); Lemma 3.1 (the labeling stays valid); Theorem 3.2 (Ford–Fulkerson: a flow is maximum iff $t$ is unreachable from $s$ in $G_f$); Lemma 3.3 (under a valid labeling $t$ is unreachable from $s$); Lemma 3.5 (from any vertex with positive excess, $s$ is reachable); Lemma 3.6 (labels never decrease; a relabeling increases the label); Lemma 3.7 ($d(v) \le 2n-1$ throughout); Theorem 3.4 (termination with finite labels gives a maximum flow); Lemma 3.8 ($\le 2n-1$ relabelings per vertex, $\le (2n-1)(n-2) < 2n^2$ in total); Lemma 3.9 ($\le 2nm$ saturating pushes); Lemma 3.10 ($\le 4n^2m$ nonsaturating pushes, under $m \ge n-1$). A further, non-milestone item states the unnumbered invariant that every $f_k$ is a preflow.

Significance

The generic bound shows that push-relabel terminates in a polynomial number of steps without any rule for choosing the next operation; the specific orderings of §4–§5 of the paper (first-in first-out, $O(n^3)$; dynamic trees, $O(nm\log(n^2/m))$) refine only the count of nonsaturating pushes, and reuse Lemmas 3.1–3.9 unchanged. The correctness argument, a valid labeling excludes augmenting paths, is the template for the push-relabel minimum-cost flow and assignment algorithms that followed.

These results are proved in the paper and are textbook material. Their machine-checked counterparts are, as far as is known here, not on the Prove2Me platform: the platform's network-flow statements (from Introduction to Linear Optimization, e.g. LinearOptimization.max_flow_min_cut) use a different model, with arc-indexed nonnegative flows and extended-real capacities, and contain nothing about preflows, labels or operation counts. This mission produces a formal account of the antisymmetric-flow model, of Ford–Fulkerson in that model, and of the amortized counting arguments, with the constants the paper prints.

Difficulty

The correctness half is short once the invariants are in place; the difficulty is in the counting. The label bound (Lemma 3.7) is a statement about the whole execution, and it depends on a structural fact about preflows (Lemma 3.5) whose truth rests on antisymmetry and on the nonnegativity of excesses. The obvious first idea for the push counts, bounding pushes per edge or per vertex locally, fails for nonsaturating pushes: flow pushed across a pair can be pushed back later, and nothing local limits how often this happens, so Lemma 3.10 holds only as an amortized statement over the entire execution and depends on both earlier counts. Saturating pushes on a pair can also recur, in both directions, and Lemma 3.9 has to control the interaction between the two directions.

Formally, all of this is reasoning about arbitrary interleavings of operations, with labels in $\mathbb{N} \cup \{\infty\}$ and real-valued flows.

Formalization scope

• Vertices form a finite type with decidable equality; $n$ is its cardinality, $s \ne t$, so $n \ge 2$ and the natural-number subtractions $2n-1$ and $n-2$ are exact. Capacities are a real function on all pairs, nonnegative, zero on the diagonal; $E$ is its support and $m$ its cardinality.
• Flows and preflows are antisymmetric real functions on all pairs (not nonnegative arc flows); the excess is computed from $f$, never stored. A maximum flow is a flow whose value is at least that of every flow.
• Labels live in ℕ∞, with $\infty + 1 = \infty$; the relabel value is an infimum, which is $\infty$ on the empty set.
• An execution is a sequence of states σ : ℕ → State V with a length $K$, starting at the Fig. 2 state with the simple labeling (the paper's own assumption for its proofs), each step an applicable push or relabel. "Terminates" means that no basic operation applies, the loop guard of Fig. 2. The three counts are cardinalities of the sets of step indices of each kind.
• Explicit constants: $2n-1$ per-vertex relabelings, $(2n-1)(n-2) < 2n^2$ total relabelings, $2nm$ saturating pushes, $4n^2m$ nonsaturating pushes, label bound $2n-1$, and the total $(2n-1)(n-2)+2nm+4n^2m$. The standing assumption $m \ge n-1$ appears only on Lemma 3.10 and the goal.
• A trivializing formalization is ruled out: the step relation fixes the pushed amount $\delta = \min(e(v), r_f(v,w))$ and the new label exactly as in Fig. 1, termination is the loop guard rather than "the result is a flow", and a sorry-free check exhibits a concrete network $s \to a \to t$ with a two-step execution (relabel $a$, then push $(a,t)$), so the run hypotheses are satisfiable.

Welcome contributions: proofs of the invariants (preflow, valid labeling, label monotonicity), of Ford–Fulkerson for antisymmetric flows (reusable beyond this mission), and of the counting lemmas. The FIFO bound of §4 is the subject of a companion mission.

Selected references

• A. V. Goldberg, R. E. Tarjan, A New Approach to the Maximum-Flow Problem, Journal of the ACM 35(4):921–940, 1988. doi:10.1145/48014.61051
• L. R. Ford, D. R. Fulkerson, Flows in Networks, Princeton University Press, 1962.
• J. Edmonds, R. M. Karp, Theoretical improvements in algorithmic efficiency for network flow problems, Journal of the ACM 19(2):248–264, 1972. doi:10.1145/321694.321699
• R. K. Ahuja, T. L. Magnanti, J. B. Orlin, Network Flows: Theory, Algorithms, and Applications, Prentice Hall, 1993.