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Operations Research

1,339 missions · 674 completed

The discipline of applying mathematical analysis to complex decision problems in operations: allocating scarce resources, scheduling, routing, inventory, and the design of service and production systems. Drawing on mathematical programming, stochastic modeling, queueing, simulation, and game-theoretic reasoning, it seeks policies that perform provably well in systems shaped by constraints, congestion, and uncertainty.

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Dynamic ProgrammingLinear OptimizationMarkov Chain·Captain: mikedeng1

Linear Programming and Sequential Decisions: An Optimal Solution of the Equilibrium LP Yields a Stationary Decision Rule of Least Expected Monthly CostResearch Paper

Motivation

Alan S. Manne's Linear Programming and Sequential Decisions (Management Science 6(3), 1960, pp. 259–267) is the first formulation of an infinite-horizon, average-cost sequential decision problem as a linear program. The illustration is a single-item inventory problem, but the construction, with unknowns indexed by a state and a decision and constraints expressing statistical equilibrium, became the standard state–action frequency linear program of Markov decision processes. Later LP approaches to average-cost Markov decision processes, including constrained ones, build on it.

Timeline of the LP approach to average-cost problems:

  • 1960 — Manne: the inventory model as a linear program in the joint probabilities of (stock level, production quantity); mixed strategies allowed; the decision rule is read off as a conditional probability.
  • 1960 — H. M. Wagner, in a companion note in the same issue, shows that an optimal solution consisting of pure strategies exists.
  • 1962 — C. Derman (Management Science 9(1), 1962) gives the general finite-state, finite-action version, assuming every stationary randomized rule yields an irreducible chain.
  • 1960 — F. d'Epenoux (Revue Française de Recherche Opérationnelle 4, No. 14; English translation 1963) treats the discounted criterion by linear programming, as Manne's closing note records.

Setting

A positive integer TTT bounds inventory accumulation; the stock levels are 0,1,…,T0,1,\dots,T0,1,…,T. At the start of a month the initial stock iii is observed and a production quantity jjj is chosen; the available stock is k=i+jk=i+jk=i+j. The month's demand n∈{0,1,2,… }n\in\{0,1,2,\dots\}n∈{0,1,2,…} is independent of everything else and has law pnp_npn​. Backlogs are excluded, so the terminal stock is t=max⁡(0,k−n)t=\max(0,k-n)t=max(0,k−n), which becomes the next initial stock. A finite set AAA of admissible pairs (i,j)(i,j)(i,j), all with i+j≤Ti+j\le Ti+j≤T and containing (i,0)(i,0)(i,0) for every i≤Ti\le Ti≤T, lists the decisions available at each stock level. Costs are three arbitrary real functions: C1(i)C_1(i)C1​(i) of the initial stock, C2(j)C_2(j)C2​(j) of the production quantity, and C3(n−k)C_3(n-k)C3​(n−k) of the shortage level.

A stationary randomized decision rule q(j∣i)q(j\mid i)q(j∣i) is a conditional probability of producing jjj at stock iii, supported on admissible pairs. It makes the initial stock a Markov chain. A statistical equilibrium of qqq is a stationary distribution y=(y0,…,yT)y=(y_0,\dots,y_T)y=(y0​,…,yT​) of that chain: the law y′y'y′ of the terminal stock equals the law yyy of the initial stock, (2). The expected monthly cost (1) of qqq in the equilibrium yyy is

EC1(i)+EC2(j)+EC3(n−k),\mathcal EC_1(i)+\mathcal EC_2(j)+\mathcal EC_3(n-k),EC1​(i)+EC2​(j)+EC3​(n−k),

the expectation taken under the joint law yi q(j∣i) pny_i\,q(j\mid i)\,p_nyi​q(j∣i)pn​ of (initial stock, production, demand).

The linear program has one unknown xijx_{ij}xij​ per admissible pair, the joint probability of (initial stock iii, production jjj). Its constraints are xij≥0x_{ij}\ge0xij​≥0, (4) ∑i,jxij=1\sum_{i,j}x_{ij}=1∑i,j​xij​=1, and the equilibrium equations

(8.t)∑jxtj=∑i,j,n:i+j−n=tpnxij(t=1,…,T),\text{(8.t)}\qquad \sum_j x_{tj}=\sum_{\substack{i,j,n:\\ i+j-n=t}}p_nx_{ij}\qquad(t=1,\dots,T),(8.t)j∑​xtj​=i,j,n:i+j−n=t​∑​pn​xij​(t=1,…,T),

and its objective (9) is ∑i,jcijxij\sum_{i,j}c_{ij}x_{ij}∑i,j​cij​xij​ with the cost coefficients (10)

cij=C1(i)+C2(j)+∑npnC3(n−i−j).c_{ij}=C_1(i)+C_2(j)+\sum_np_nC_3(n-i-j).cij​=C1​(i)+C2​(j)+n∑​pn​C3​(n−i−j).

The companion equation (8.0) for t=0t=0t=0 has right-hand side ∑i+j−n≤0pnxij\sum_{i+j-n\le0}p_nx_{ij}∑i+j−n≤0​pn​xij​ and is omitted from the constraints. A feasible xxx is decoded into yi=∑jxijy_i=\sum_jx_{ij}yi​=∑j​xij​ and q(j∣i)=xij/yiq(j\mid i)=x_{ij}/y_iq(j∣i)=xij​/yi​.

Formalization targets

The paper labels no theorem or lemma. The goal is assembled from §1 (third paragraph), §3 (N.B.), §4 (last two paragraphs), §5 and §7 (3), and every milestone is cited by section, display, table or footnote.

Goal: an LP optimum gives an optimal stationary rule

Assume ∑npn∣C3(n−i−j)∣<∞\sum_np_n|C_3(n-i-j)|<\infty∑n​pn​∣C3​(n−i−j)∣<∞ for every admissible pair. Then the linear program has an optimal solution, and for every optimal solution x∗x^*x∗, with decoding (q∗,y∗)(q^*,y^*)(q∗,y∗), q∗q^*q∗ is a stationary randomized rule, y∗y^*y∗ is a statistical equilibrium of q∗q^*q∗, and

Cost(q∗,y∗)=∑i,jcijxij∗≤Cost(q,y)\mathrm{Cost}(q^*,y^*)=\sum_{i,j}c_{ij}x^*_{ij}\le \mathrm{Cost}(q,y)Cost(q∗,y∗)=i,j∑​cij​xij∗​≤Cost(q,y)

for every stationary randomized rule qqq and every statistical equilibrium yyy of qqq.

Milestones

  1. (7): under a rule in a distribution yyy, the law of the terminal stock is the right-hand side of (7)/(8) evaluated at xij=yiq(j∣i)x_{ij}=y_iq(j\mid i)xij​=yi​q(j∣i).
  2. (8.0)–(8.T): a rule in statistical equilibrium yields a point satisfying x≥0x\ge0x≥0, (4) and all of (8.0)–(8.T).
  3. (8.0) is redundant: (4) and (8.1)–(8.T) imply (8.0).
  4. §3, N.B.: every feasible xxx equals yiq(j∣i)y_iq(j\mid i)yi​q(j∣i) for its decoding (q,y)(q,y)(q,y), with yyy an equilibrium of qqq.
  5. (10): the expected monthly cost (1) under the joint law xijpnx_{ij}p_nxij​pn​ equals ∑cijxij\sum c_{ij}x_{ij}∑cij​xij​.
  6. Table 1: the cost coefficients of the §6 example (T=3T=3T=3, p=(2/3,0,1/3)p=(2/3,0,1/3)p=(2/3,0,1/3), C1(i)=iC_1(i)=iC1​(i)=i, C2(j)=3jC_2(j)=3jC2​(j)=3j, C3(m)=max⁡[0,6m]C_3(m)=\max[0,6m]C3​(m)=max[0,6m], j∈{0,1}j\in\{0,1\}j∈{0,1}) are 4,5,3,4,2,5,34,5,3,4,2,5,34,5,3,4,2,5,3.
  7. Table 2, footnote 3: x01=1/3x_{01}=1/3x01​=1/3, x11=2/9x_{11}=2/9x11​=2/9, x20=4/9x_{20}=4/9x20​=4/9 is optimal with cost 31/931/931/9; the do-nothing solution costs 444.
  8. Footnote 5: the implicit prices −7/3,−13/3,−11/3-7/3,-13/3,-11/3−7/3,−13/3,−11/3 of (8.1)–(8.3), with 31/931/931/9 on (4), are an optimal dual solution.

Significance

The result turns an infinite-horizon control problem into a finite linear program. Equilibrium joint laws of (state, decision) under stationary randomized rules are exactly the feasible points of a polytope, and the average cost is linear on it. Consequences include computability by the simplex method; an economic reading of the dual variables (Manne's footnote 5 interprets them as the relative advantage of starting at a given stock level, related to Bellman's functional equation); and, in later work, the treatment of side constraints, which dynamic programming handles poorly.

The result is classical and proved in the paper (largely by inspection of the definitions). It has not been formalized. The formalization adds three things. First, a precise statement of what is optimized when the chain of a rule is not irreducible: the paper's §7 (3) concedes that a "decomposable" optimum makes the equilibrium depend on initial conditions, and the goal resolves this by optimizing over (rule, equilibrium) pairs. Second, an explicit treatment of the stock levels the equilibrium never visits, where the paper's quotient xij/∑jxijx_{ij}/\sum_jx_{ij}xij​/∑j​xij​ is undefined. Third, a machine-checked numerical example whose LP is derived from the general definitions, not entered by hand. Derman's later irreducible-case version exists on the platform as a separate open statement; this mission covers Manne's irreducibility-free version on state-dependent action sets.

Difficulty

Each step is elementary; the work is bookkeeping across three descriptions of the same object. The equilibrium is defined through the transition kernel of the controlled chain, the LP through the displayed sums over (i,j,n)(i,j,n)(i,j,n) with conditions i+j−n≤0i+j-n\le0i+j−n≤0 and i+j−n=ti+j-n=ti+j−n=t, and the cost through the joint law of three variables. Identifying them needs a reindexing of the admissible pairs by stock level, the interchange of a finite sum with an infinite sum over demands, and ∑npn=1\sum_np_n=1∑n​pn​=1. The naive argument "the LP constraints are the equilibrium equations, so the LP optimum is the optimal rule" skips two points. The constraints omit (8.0), so equilibrium at stock level 000 must be recovered from (4) and the bound i+j≤Ti+j\le Ti+j≤T. And the decoding fails at unvisited stock levels unless a default action is supplied. Existence of an LP optimum requires compactness of the feasible polytope, not just its nonemptiness.

Formalization scope

All statements live in the namespace ManneLP.Equilibrium. Conventions:

  • Stock levels and production quantities are natural numbers; a model carries T>0T>0T>0, the admissible set AAA (a Finset (ℕ × ℕ) with i+j≤Ti+j\le Ti+j≤T and every (i,0)∈A(i,0)\in A(i,0)∈A), a demand law p:N→Rp:\mathbb N\to\mathbb Rp:N→R with pn≥0p_n\ge0pn​≥0 and HasSum p 1, and costs C1,C2:N→RC_1,C_2:\mathbb N\to\mathbb RC1​,C2​:N→R, C3:Z→RC_3:\mathbb Z\to\mathbb RC3​:Z→R. The shortage level n−i−jn-i-jn−i−j is an integer; no convexity, sign or monotonicity of the costs is assumed.
  • The admissible set is a parameter: §6 imposes a capacity limit j≤1j\le1j≤1. Reading of the page: i+j≤Ti+j\le Ti+j≤T is how §2's requirement max⁡(0,k−n)≤T\max(0,k-n)\le Tmax(0,k−n)≤T holds whatever the demand; (i,0)∈A(i,0)\in A(i,0)∈A (producing nothing is possible) is implicit in §2.
  • The terminal stock is computed by truncated subtraction in N\mathbb NN, which equals max⁡(0,k−n)\max(0,k-n)max(0,k−n). Sums over demands are tsums; the demand is not assumed bounded. The goal and the cost identity assume the expected shortage cost at each admissible pair is finite (absolutely summable), which the page takes for granted.
  • A statistical equilibrium is a stationary distribution of the chain of the rule, defined from the transition probabilities, not from (8). The expected monthly cost is defined from the joint law of (stock, production, demand), not as ∑cijxij\sum c_{ij}x_{ij}∑cij​xij​. The LP constraint set omits (8.0), exactly as the page does.
  • The decoded rule uses the default action j=0j=0j=0 at stock levels with ∑jxij=0\sum_jx_{ij}=0∑j​xij​=0; any admissible default would do.
  • Table 2's x30=εx_{30}=\varepsilonx30​=ε is the paper's device against degeneracy; the example's solution has x30=0x_{30}=0x30​=0. Footnote 5 prints no price for (4); 31/931/931/9 is the price forced by equal objectives, and the dual optimum is not claimed unique.

Trivializing formalizations are ruled out: equilibrium is not defined as (8) (which would make milestone 2 vacuous), (8.0) is not a constraint (which would make milestone 3 vacuous), the cost is not defined as ∑cijxij\sum c_{ij}x_{ij}∑cij​xij​ (which would make milestone 5 and the goal's cost clause definitional), and the optimality comparison ranges over all rules and all of their equilibria, not over irreducible chains or pure rules.

Contributions welcome: proofs of the milestones and the goal; computations of the §6 example from the general definitions; reusable lemmas on stationary distributions of finite stochastic matrices and on the existence of LP optima over compact polytopes.

Selected references

  • A. S. Manne, Linear Programming and Sequential Decisions, Management Science 6(3), 259–267, 1960. https://doi.org/10.1287/mnsc.6.3.259
  • H. M. Wagner, On the Optimality of Pure Strategies, Management Science 6(3), 268–269, 1960. https://doi.org/10.1287/mnsc.6.3.268
  • C. Derman, On Sequential Decisions and Markov Chains, Management Science 9(1), 16–24, 1962. https://doi.org/10.1287/mnsc.9.1.16
  • F. d'Epenoux, A Probabilistic Production and Inventory Problem, Management Science 10(1), 98–108, 1963 (translation of the 1960 French paper). https://doi.org/10.1287/mnsc.10.1.98
13 thms1 active userReviewed
OptimizationProbabilityStatistics·Captain: mikedeng1

Asymptotic Theory for Solutions in Statistical Estimation and Stochastic Programming: Generalized M-Estimates Converge in Distribution to the Inverse Contingent Derivative at a GaussianResearch Paper

Motivation

Maximum likelihood estimates, least-squares fits and sample-average approximations of stochastic programs solve 0=fˉν(x)0 = \bar f^\nu(x)0=fˉ​ν(x), where fˉν\bar f^\nufˉ​ν averages a random integrand over ν\nuν observations. Their classical asymptotic theory rests on the implicit function theorem and needs a smooth, unconstrained problem.

When the estimate is constrained to a set, for example a nonnegativity constraint, a simplex or a polyhedron, the first-order conditions become a generalized equation

0∈f(z,x)+N(x),0 \in f(z, x) + N(x),0∈f(z,x)+N(x),

where NNN is a multifunction such as the normal cone of the constraint set. The same form describes optimality conditions of stochastic programs and variational inequalities. Aitchison and Silvey (1958) treated equality-constrained maximum likelihood. Huber (1967) allowed nonsmooth estimating functions but required an open parameter domain. Dupačová and Wets (1988) and Shapiro (1989) derived limit laws for solutions of stochastic programs under smoothness of the expected gradient. King and Rockafellar (1993) gave a general theory that needs neither smoothness of the expected map nor single-valuedness of NNN: the limit law of the normalized error is the image of a Gaussian under a contingent derivative, a positively homogeneous and generally nonlinear map, so the limit is generally not normal.

Setting

Let ZZZ be a separable Banach space with norm ∥⋅∥\|\cdot\|∥⋅∥, and let ∣⋅∣|\cdot|∣⋅∣ be the Euclidean norm on Rn\mathbb R^nRn and Rm\mathbb R^mRm. A multifunction G:Z⇉RnG : Z \rightrightarrows \mathbb R^nG:Z⇉Rn assigns a set G(z)⊆RnG(z) \subseteq \mathbb R^nG(z)⊆Rn to each zzz. Its graph is gph⁡G\operatorname{gph} GgphG and its inverse is G−1(x)={z∣x∈G(z)}G^{-1}(x) = \{z \mid x \in G(z)\}G−1(x)={z∣x∈G(z)}.

For sets AtA_tAt​ indexed by t↓0t \downarrow 0t↓0, the upper limit lim sup⁡At\limsup A_tlimsupAt​ consists of the points xxx with x=lim⁡xkx = \lim x_kx=limxk​, xk∈Atkx_k \in A_{t_k}xk​∈Atk​​ for some tk↓0t_k \downarrow 0tk​↓0, and the lower limit of the points reachable along every such sequence. The contingent derivative of GGG at (z,x)∈gph⁡G(z, x) \in \operatorname{gph} G(z,x)∈gphG is the multifunction DG(z∣x)DG(z|x)DG(z∣x) with

gph⁡DG(z∣x)=lim sup⁡t↓0t−1[gph⁡G−(z,x)].\operatorname{gph} DG(z|x) = \limsup_{t \downarrow 0} t^{-1}\big[\operatorname{gph} G - (z,x)\big].gphDG(z∣x)=t↓0limsup​t−1[gphG−(z,x)].

GGG is proto-differentiable when this upper limit equals the lower limit, and semi-differentiable when t−1[G(z+tw′)−x]→DG(z∣x)(w)t^{-1}[G(z + t w') - x] \to DG(z|x)(w)t−1[G(z+tw′)−x]→DG(z∣x)(w) as t↓0t \downarrow 0t↓0 and w′→ww' \to ww′→w. A single-valued ggg is B-differentiable at zzz when t−1[g(z+tw′)−g(z)]→Dg(z)(w)t^{-1}[g(z + tw') - g(z)] \to Dg(z)(w)t−1[g(z+tw′)−g(z)]→Dg(z)(w) in the same sense.

The deterministic problem is 0∈f(z,x)+N(x)0 \in f(z, x) + N(x)0∈f(z,x)+N(x) with f:Z×Rn→Rmf : Z \times \mathbb R^n \to \mathbb R^mf:Z×Rn→Rm, data zzz and solution map J(z)J(z)J(z). At a reference pair (z∗,x∗)(z^*, x^*)(z∗,x∗) set F=f(z∗,⋅)+NF = f(z^*, \cdot) + NF=f(z∗,⋅)+N. The analytical assumptions M.1–M.4 are as follows. fff is jointly continuous and B-differentiable in each variable, with the zzz-derivative Dzf(z∗,x∗)D_z f(z^*, x^*)Dz​f(z∗,x∗) strong (uniform in xxx near x∗x^*x∗). NNN is closed and proto-differentiable. FFF is subinvertible: 0∈F(x∗)0 \in F(x^*)0∈F(x∗), and a closed-graph, convex-valued selection of F−1F^{-1}F−1 near 000 passes through x∗x^*x∗. The contingent derivative DF−1(0∣x∗)DF^{-1}(0|x^*)DF−1(0∣x∗) is at most single-valued.

The statistical problem has i.i.d. random elements s1,s2,…s_1, s_2, \dotss1​,s2​,… of a measurable space SSS and an integrand f:U×S→Rmf : U \times S \to \mathbb R^mf:U×S→Rm on a compact neighborhood UUU of x∗x^*x∗. It satisfies the probabilistic assumptions P.1–P.4: continuity in xxx, measurability in sss, a finite second moment at one point, and a Lipschitz bound ∣f(x1,s)−f(x2,s)∣≤a(s)∣x1−x2∣|f(x_1,s) - f(x_2,s)| \le a(s)|x_1 - x_2|∣f(x1​,s)−f(x2​,s)∣≤a(s)∣x1​−x2​∣ with Ea(s1)2<∞E a(s_1)^2 < \inftyEa(s1​)2<∞. The M-estimate xνx^\nuxν is a measurable solution of

0∈fˉν(x)+N(x),fˉν(x)=1ν∑i=1νf(x,si),0 \in \bar f^\nu(x) + N(x), \qquad \bar f^\nu(x) = \frac1\nu\sum_{i=1}^\nu f(x, s_i),0∈fˉ​ν(x)+N(x),fˉ​ν(x)=ν1​i=1∑ν​f(x,si​),

and the true equation is 0∈Ef(x)+N(x)0 \in Ef(x) + N(x)0∈Ef(x)+N(x) with F=Ef+NF = Ef + NF=Ef+N.

Formalization targets

Goal: Theorem 2.7 (asymptotic distribution of M-estimates)

Under P.1–P.4 on a compact neighborhood UUU of x∗x^*x∗, B-differentiability of EfEfEf at x∗x^*x∗, and M.2–M.4 for F=Ef+NF = Ef + NF=Ef+N, every sequence of measurable solutions xνx^\nuxν of (2.5) with xν→x∗x^\nu \to x^*xν→x∗ almost surely satisfies

ν [xν−x∗]→ D DF−1(0∣x∗)(−w∗),w∗∼N(0,cov⁡f(x∗,s1)).\sqrt\nu\,[x^\nu - x^*] \xrightarrow{\ \mathcal D\ } DF^{-1}(0|x^*)(-w^*), \qquad w^* \sim \mathcal N\big(0, \operatorname{cov} f(x^*, s_1)\big).ν​[xν−x∗] D ​DF−1(0∣x∗)(−w∗),w∗∼N(0,covf(x∗,s1​)).

The goal fixes the limit law completely: the map is the contingent derivative of F−1F^{-1}F−1, and the Gaussian has the covariance of the integrand at x∗x^*x∗.

Milestones

  • Theorem 2.4 gives bounds in probability, P{∣xν−x∗∣>δ}≤P{αλ∥zν−z∗∥>δ}P\{|x^\nu - x^*| > \delta\} \le P\{\alpha\lambda\|z^\nu - z^*\| > \delta\}P{∣xν−x∗∣>δ}≤P{αλ∥zν−z∗∥>δ}. Its proof uses the upper-Lipschitz property U∩F−1(y)⊆x∗+λ∣y∣BU \cap F^{-1}(y) \subseteq x^* + \lambda|y|BU∩F−1(y)⊆x∗+λ∣y∣B (a display of the proof).
  • Theorem 2.6 is the abstract limit theorem: if τν−1[zν−z∗]→Dw\tau_\nu^{-1}[z^\nu - z^*] \to_{\mathcal D} wτν−1​[zν−z∗]→D​w, then τν−1[xν−x∗]→DDF−1(0∣x∗)(−Dzf(z∗,x∗)(w))\tau_\nu^{-1}[x^\nu - x^*] \to_{\mathcal D} DF^{-1}(0|x^*)(-D_z f(z^*,x^*)(w))τν−1​[xν−x∗]→D​DF−1(0∣x∗)(−Dz​f(z∗,x∗)(w)). Its proof uses two displays: semi-differentiability of the localized solution map, with DJ(z∗∣x∗)(w)=DF−1(0∣x∗)(−Dzf(z∗,x∗)(w))DJ(z^*|x^*)(w) = DF^{-1}(0|x^*)(-D_z f(z^*,x^*)(w))DJ(z∗∣x∗)(w)=DF−1(0∣x∗)(−Dz​f(z∗,x∗)(w)), and a Lipschitz bound ∣x−x∗∣≤λ∥z−z∗∥|x - x^*| \le \lambda\|z - z^*\|∣x−x∗∣≤λ∥z−z∗∥ on U∩J(z)U \cap J(z)U∩J(z).
  • Proposition A1, Corollary A2 and Theorem A3 concern the space Cm(U)C_m(U)Cm​(U) under P.1–P.4. The integrand and the empirical means are random elements of Cm(U)C_m(U)Cm​(U), and ν(fˉν−Ef)\sqrt\nu(\bar f^\nu - Ef)ν​(fˉ​ν−Ef) converges in distribution to a Gaussian element of Cm(U)C_m(U)Cm​(U).

The three displays (Theorem 2.4's upper-Lipschitz inclusion, and Theorem 2.6's semi-differentiability and Lipschitz bound) are statements the paper cites from King and Rockafellar, Sensitivity analysis for nonsmooth generalized equations ([12]: Proposition 2.1, Theorem 4.1, Remark 4.3). They are cited results, not this paper's own, and are milestones because the proofs of Theorems 2.4 and 2.6 rest on them.

Significance

Theorem 2.7 gives the limit law of constrained and nonsmooth M-estimates in a form that can be computed. When NNN is the normal cone of a polyhedron, DF−1(0∣x∗)DF^{-1}(0|x^*)DF−1(0∣x∗) is piecewise linear, and the limit is the solution of a random linear complementarity or quadratic problem driven by a Gaussian vector. This underlies the asymptotic theory of sample-average approximation in stochastic programming, where the paper applies it to stochastic programs (§3) and to piecewise linear-quadratic tracking problems (§4). Theorem 2.6 separates the deterministic sensitivity analysis from the probability, so any data sequence with a known limit law yields a limit law for the solutions.

The result is proved in the paper modulo the cited theorems of [12] and [11], but none of it is machine-checked. Mathlib has the real-valued i.i.d. central limit theorem, Gaussian measures on Banach spaces and convergence in distribution. It has no multivariate or Banach-space central limit theorem, no contingent derivatives and no set-valued implicit function theorem. Formalizing the mission produces these, along with a checked version of the cited sensitivity results.

Difficulty

The classical argument linearizes FFF at x∗x^*x∗, inverts the Jacobian and applies the delta method. Here FFF is set-valued and its derivative is only positively homogeneous. There is no Jacobian to invert, and the solution map need not be differentiable or even single-valued away from x∗x^*x∗. The replacement for the implicit function theorem is the semi-differentiability of the localized solution map under M.1–M.4. Proving it means controlling both the upper and the lower set limits of difference quotients of solution sets, and subinvertibility is what supplies existence of nearby solutions.

The probabilistic side cannot work coordinate by coordinate either. The estimate solves an equation in the whole function fˉν\bar f^\nufˉ​ν, so convergence of fˉν\bar f^\nufˉ​ν at finitely many points is not enough. The central limit theorem must hold in the sup norm on Cm(U)C_m(U)Cm​(U), which requires tightness of the empirical process, and only then can the deterministic sensitivity result be composed with it.

Formalization scope

Points live in EuclideanSpace ℝ (Fin n), and ZZZ is a real normed space with [CompleteSpace Z] [SeparableSpace Z] where the paper says "separable Banach". Set limits are Kuratowski limits along filters (t↓0t \downarrow 0t↓0 is 𝓝[>] 0, and (t,w′)→(0+,w)(t,w') \to (0^+,w)(t,w′)→(0+,w) is the product filter). The contingent derivative is defined by (2.2) alone. M.4's printed sum formula equals it under M.1, and this is not assumed. "B-differentiable" is read as the limit (2.4). Products carry Lean's max norm. s1s_1s1​ is s 0, empirical means sum over Finset.range ν, and (2.5) is required for ν≥1\nu \ge 1ν≥1. F=Ef+NF = Ef + NF=Ef+N is empty off UUU. Convergence in distribution is Mathlib's TendstoInDistribution, with the limit on its own probability space. The law of w∗w^*w∗ is fixed through linear functionals: ⟨ℓ,w∗⟩∼N(0,Var⁡⟨ℓ,f(x∗,s1)⟩)\langle\ell, w^*\rangle \sim \mathcal N(0, \operatorname{Var}\langle\ell, f(x^*,s_1)\rangle)⟨ℓ,w∗⟩∼N(0,Var⟨ℓ,f(x∗,s1​)⟩). In Appendix A1–A3 the integrand is S → C(↥U, Rn m) with the Borel σ-algebra, and "Gaussian" is IsGaussian.

Explicit choices relative to the printed text:

  1. The paper states that an almost surely convergent sequence of solutions "converges to the point x∗x^*x∗" (Theorems 2.6 and 2.7). This is false when the true equation has a second solution: f(z,x)=x2−x−zf(z,x) = x^2 - x - zf(z,x)=x2−x−z, N≡{0}N \equiv \{0\}N≡{0}, z∗=x∗=0z^* = x^* = 0z∗=x∗=0 satisfies M.1–M.4 with J(0)={0,1}J(0) = \{0, 1\}J(0)={0,1}. The formalization assumes xν→x∗x^\nu \to x^*xν→x∗ almost surely instead.
  2. 0∈F(x∗)0 \in F(x^*)0∈F(x∗) (presupposed by DF−1(0∣x∗)DF^{-1}(0|x^*)DF−1(0∣x∗)) is explicit in Theorem 2.4 and the upper-Lipschitz display.
  3. The threshold for "all sufficiently small δ\deltaδ" in Theorem 2.4 is chosen with UUU and λ\lambdaλ, before the random elements.
  4. Proposition A1 and Corollary A2 carry P.1–P.4, as stated or inherited on the Appendix page, though their measurability conclusions use only P.1.
  5. In the limit theorems, the single-valued map DF−1(0∣x∗)DF^{-1}(0|x^*)DF−1(0∣x∗) is a function LLL whose values lie in the contingent derivative at every point.

The conclusion of the goal names its limit: the image of the stated Gaussian under a map LLL whose values lie in the contingent derivative of F−1F^{-1}F−1. A statement asserting only that ν(xν−x∗)\sqrt\nu(x^\nu - x^*)ν​(xν−x∗) converges in distribution to some limit, or replacing DF−1(0∣x∗)DF^{-1}(0|x^*)DF−1(0∣x∗) by a linear map, is a different and weaker theorem. A sorry-free check confirms that the goal's hypotheses can all be met (a degenerate instance with f(x,s)=xf(x,s) = xf(x,s)=x, N≡{0}N \equiv \{0\}N≡{0}).

A complete development needs Kuratowski set convergence and contingent derivatives (reusable across set-valued analysis), a central limit theorem in C(K)C(K)C(K) for Lipschitz-indexed processes (reusable for empirical-process theory), and a continuous-mapping argument for random closed sets. Contributions to any of these layers are welcome, as are proofs of the cited [12] statements.

Selected references

  • A. J. King and R. T. Rockafellar, Asymptotic theory for solutions in statistical estimation and stochastic programming, Mathematics of Operations Research 18(1) (1993). https://doi.org/10.1287/moor.18.1.148
  • A. J. King and R. T. Rockafellar, Sensitivity analysis for nonsmooth generalized equations, Mathematical Programming 55 (1992) 193–212. https://doi.org/10.1007/BF01581199
  • A. J. King, Generalized delta theorems for multivalued mappings and measurable selections, Mathematics of Operations Research 14(4) (1989) 720–736. https://doi.org/10.1287/moor.14.4.720
  • J. Dupačová and R. J.-B. Wets, Asymptotic behavior of statistical estimators and of optimal solutions of stochastic optimization problems, Annals of Statistics 16(4) (1988) 1517–1549. https://doi.org/10.1214/aos/1176351052
  • A. Shapiro, Asymptotic properties of statistical estimators in stochastic programming, Annals of Statistics 17(2) (1989) 841–858. https://doi.org/10.1214/aos/1176347146
  • P. J. Huber, The behavior of maximum likelihood estimates under nonstandard conditions, Proc. Fifth Berkeley Symp. Math. Statist. Probab. 1 (1967) 221–233. https://projecteuclid.org/euclid.bsmsp/1200512988
  • A. Araujo and E. Giné, The Central Limit Theorem for Real and Banach Valued Random Variables, Wiley, 1980.
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Dynamic ProgrammingOptimization·Captain: mikedeng1

A Functional Equation and Its Application to Resource Allocation and Sequencing Problems 1: Equation (1) Computes the Minimum Total Loss of a Fixed Job Order with Two Processing Modes per JobResearch Paper

Motivation

Many single-machine scheduling problems have a structure that is easy to describe: the jobs must be processed in an order that is already known, and the only decisions are how and when to process each one. Lawler and Moore, in A Functional Equation and Its Application to Resource Allocation and Sequencing Problems (Management Science 16(1), 1969, 77–84), isolate this situation in its simplest form: each job has two processing modes, and each mode has its own processing time and loss as a function of the completion time. They solve it with a single functional equation, Eq. (1), closely related to the recursion for the knapsack problem.

The point of the paper is that Eq. (1) solves more than this fixed-order problem. Later sections apply it to resource allocation in critical-path scheduling and to several sequencing problems with deadlines, including minimization of the weighted number of tardy jobs, often written 1∥∑wjUj1\|\sum w_jU_j1∥∑wj​Uj​. Everything in those applications rests on the claim checked in this mission: that the recursion computes the minimum loss it is said to compute.

Setting

There are nnn jobs, performed one at a time in the fixed order 1,2,…,n1, 2, \dots, n1,2,…,n. Job jjj can be performed in one of two modes. In the first mode it takes aja_jaj​ time units, and a loss αj(t)\alpha_j(t)αj​(t) is incurred if it is completed at time ttt. In the other mode it takes bjb_jbj​ time units, with loss βj(t)\beta_j(t)βj​(t). Times are nonnegative integers.

A mode assignment mmm picks a mode for each job, which fixes its processing time pj∈{aj,bj}p_j \in \{a_j, b_j\}pj​∈{aj​,bj​}. A feasible timing of the first jjj jobs is a vector of completion times c1,…,cjc_1, \dots, c_jc1​,…,cj​ with

ci−1+pi≤ci(i=1,…,j),c0=0,c_{i-1} + p_i \le c_i \qquad (i = 1, \dots, j), \qquad c_0 = 0,ci−1​+pi​≤ci​(i=1,…,j),c0​=0,

so each job starts at time 000 or later and after its predecessor has finished; idle time between jobs is allowed. The total loss Lj(m,c)L_j(m, c)Lj​(m,c) is the sum of αi(ci)\alpha_i(c_i)αi​(ci​) over the jobs i≤ji \le ji≤j in the first mode and of βi(ci)\beta_i(c_i)βi​(ci​) over the others. The problem asks for the mode assignment and timing of all nnn jobs that minimize Ln(m,c)L_n(m, c)Ln​(m,c).

The paper's recursion is defined for j=0,…,nj = 0, \dots, nj=0,…,n and integer ttt, with values in R∪{+∞}\mathbb R \cup \{+\infty\}R∪{+∞}:

f(0,t)=0 (t≥0),f(j,t)=+∞ (t<0),f(0, t) = 0 \ (t \ge 0), \qquad f(j, t) = +\infty \ (t < 0),f(0,t)=0 (t≥0),f(j,t)=+∞ (t<0), f(j,t)=min⁡{f(j,t−1), αj(t)+f(j−1,t−aj), βj(t)+f(j−1,t−bj)}(j≥1, t≥0).(1)f(j, t) = \min\{f(j, t-1),\ \alpha_j(t) + f(j-1, t-a_j),\ \beta_j(t) + f(j-1, t-b_j)\} \quad (j \ge 1,\ t \ge 0). \tag{1}f(j,t)=min{f(j,t−1), αj​(t)+f(j−1,t−aj​), βj​(t)+f(j−1,t−bj​)}(j≥1, t≥0).(1)

Formalization targets

Milestone: Eq. (1) computes the constrained minimum

For 0≤j≤n0 \le j \le n0≤j≤n and every integer ttt, let L(j,t)\mathcal L(j, t)L(j,t) be the set of total losses of the first jjj jobs over all mode assignments and feasible timings with cj≤tc_j \le tcj​≤t. Then

f(j,t)=min⁡L(j,t),f(j, t) = \min \mathcal L(j, t),f(j,t)=minL(j,t),

meaning: f(j,t)=+∞f(j, t) = +\inftyf(j,t)=+∞ exactly when L(j,t)\mathcal L(j, t)L(j,t) is empty, and otherwise f(j,t)f(j,t)f(j,t) is an element of L(j,t)\mathcal L(j,t)L(j,t) and a lower bound for it. No assumption is made on the losses.

Goal: f(n,T)f(n, T)f(n,T) solves the problem

If every αj\alpha_jαj​ and βj\beta_jβj​ is monotone nondecreasing and

T=∑j=1nmax⁡{aj,bj},T = \sum_{j=1}^n \max\{a_j, b_j\},T=j=1∑n​max{aj​,bj​},

then f(n,T)f(n, T)f(n,T) is finite and equals the minimum of Ln(m,c)L_n(m, c)Ln​(m,c) over all mode assignments and all feasible timings of the nnn jobs, with no deadline. This is the paper's sentence "The problem is solved by the calculation of f(n,T)f(n, T)f(n,T), where TTT is a sufficiently large number. For example, if all of the αj\alpha_jαj​'s and βj\beta_jβj​'s are monotone nondecreasing, we may choose T=∑j=1nmax⁡{aj,bj}T = \sum_{j=1}^n \max\{a_j, b_j\}T=∑j=1n​max{aj​,bj​}."

Significance

The milestone is the correctness of a dynamic program: the value table computed by (1) coincides with the optimum of a combinatorial problem whose feasible set is infinite (timings are unbounded because of idle time). The goal turns that into a finite certificate for the unconstrained problem, with an explicit horizon. Downstream, the same equation with bj=0b_j = 0bj​=0 and suitable losses is the paper's algorithm for the weighted number of tardy jobs (§5–6), and the other sequencing applications are specializations of (1) as well; a formal version of this mission is the base on which those reductions can be stated.

The result is classical and its proof is short on paper. It has not been formalized: no item on Prove2Me states a two-mode fixed-order recursion. The work this mission asks for is a machine-checked proof of the two statements above, from the paper's own definitions.

Difficulty

The paper gives no argument beyond "the usual dynamic programming argumentation". The recursion is in two variables: f(j,⋅)f(j, \cdot)f(j,⋅) refers to itself at t−1t-1t−1 as well as to f(j−1,⋅)f(j-1, \cdot)f(j−1,⋅), so the correspondence with timings is not a plain stage-by-stage principle of optimality. The edge cases are where a careless reading fails: t=0t = 0t=0, where f(j,−1)=+∞f(j, -1) = +\inftyf(j,−1)=+∞; zero processing times aj=0a_j = 0aj​=0 or bj=0b_j = 0bj​=0, where f(j−1,t−aj)f(j-1, t-a_j)f(j−1,t−aj​) is evaluated at the same ttt; and j=0j = 0j=0, where the deadline constraint degenerates to 0≤t0 \le t0≤t. The goal is harder than the milestone: the problem's feasible set is unbounded, and the horizon TTT is only sufficient under monotone losses. Without monotonicity the goal is false: one job with a=b=1a = b = 1a=b=1, α(1)=5\alpha(1) = 5α(1)=5, α(t)=0\alpha(t) = 0α(t)=0 for t≥2t \ge 2t≥2 and β≡5\beta \equiv 5β≡5 has optimum 000, attained only at completion time 2>T=12 > T = 12>T=1, while f(1,1)=5f(1, 1) = 5f(1,1)=5.

Formalization scope

  • Jobs are Fin n and 0-based: Lean job i is the paper's job i+1i+1i+1. The recursion f takes the number j∈{0,…,n}j \in \{0,\dots,n\}j∈{0,…,n} of jobs done; for j>nj > nj>n its value is +∞+\infty+∞ and is never used.
  • Modes are Bool, with true the first mode (aja_jaj​, αj\alpha_jαj​). Processing times are natural numbers and may be 000. Losses are real-valued functions of the natural-number completion time.
  • Integer time is an explicit reading: the paper's recursion steps ttt by one, so time is discrete.
  • +∞+\infty+∞ is ⊤ : WithTop ℝ, so x+(+∞)=+∞x + (+\infty) = +\inftyx+(+∞)=+∞ as in the paper.
  • Idle time is allowed, and the first job starts at time 000 or later.
  • "Minimum total loss" is read as an attained least element of the set of total losses, with +∞+\infty+∞ exactly when that set is empty. "The problem is solved by f(n,T)f(n, T)f(n,T)" is read as: f(n,T)f(n, T)f(n,T) is finite, at most every feasible total loss, and attained.
  • "Monotone nondecreasing" is Monotone on the completion time; it is a hypothesis of the goal only.

The optimum is defined from the scheduling problem itself (lossesBy, IsFeasible, totalLoss), never as fff; statements in which the optimum is fff again, timings forced to have no idle time, an R\mathbb RR-valued fff with an arbitrary value for t<0t < 0t<0, or monotone losses assumed in the milestone are all ruled out by this choice of definitions.

Related platform items: GilmoreGomory61.CuttingStock.knapsack_dp_recursion (an unbounded-knapsack recursion) and the CriticalPath.CostCurve items (Kelley's continuous time–cost model) treat neighbouring recursions and models; neither states this problem. Proofs of the milestone and the goal, and reusable lemmas about the value function, are welcome.

Selected references

  • E. L. Lawler, J. M. Moore, A Functional Equation and Its Application to Resource Allocation and Sequencing Problems, Management Science 16(1) (1969) 77–84. https://doi.org/10.1287/mnsc.16.1.77
  • R. Bellman, Dynamic Programming, Princeton University Press, 1957. https://doi.org/10.1515/9781400835386
  • J. M. Moore, An n Job, One Machine Sequencing Algorithm for Minimizing the Number of Late Jobs, Management Science 15(1) (1968) 102–109. https://doi.org/10.1287/mnsc.15.1.102
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Dynamic ProgrammingMarkov ChainProbability·Captain: mikedeng1

Some Monotonicity Results for Partially Observed Markov Decision Processes: MLR-Monotone Optimal Values and the Myopic Policy as a Lower Bound on the Optimal PolicyResearch Paper

Motivation

A partially observed Markov decision process (POMDP) models a controller that cannot see the state of the system it controls. It sees only noisy observations, and so it acts on a belief: a probability vector over the hidden states. Machine maintenance, medical screening, quality control and search problems all have this form. The dynamic program of a POMDP lives on the simplex of beliefs, a continuum, so computing exact optimal policies is expensive even when the state, action and observation sets are small. Structural results reduce that cost. A value function that is monotone in the belief, or a policy known to dominate a cheap reference policy, shrinks the space a computation must search. Such results also explain the model: they say when one belief is "better" than another.

W. S. Lovejoy, Some Monotonicity Results for Partially Observed Markov Decision Processes (Operations Research 35(5):736–743, 1987), provides such results by ordering beliefs with the monotone likelihood ratio (MLR) order instead of first-order stochastic dominance.

Timeline. Smallwood and Sondik (1973) set out the finite POMDP and its piecewise-linear value functions. White (1979, 1980) obtained monotone policies and values for machine replacement, for single-stage problems, and for the completely observed and completely unobserved extremes, all under first-order stochastic dominance. Albright (1979) treated the two-state case, where the usual orders coincide. Whitt (1979, 1982) developed the likelihood-ratio orders and showed that they are preserved by Bayesian updating. Lovejoy (1987) combined these into general monotonicity results for finite POMDPs, and the MLR order has since become the standard tool for structural results in POMDPs.

Setting

Let S={1,…,n}S=\{1,\dots,n\}S={1,…,n} (states) and O={1,…,m}O=\{1,\dots,m\}O={1,…,m} (observations) carry their natural orders, and let AAA be a finite, completely ordered action set. For a finite chain XXX, Π(X)\Pi(X)Π(X) is the set of probability vectors on XXX. For π,π′∈Π(X)\pi,\pi'\in\Pi(X)π,π′∈Π(X), π≥sπ′\pi\ge_s\pi'π≥s​π′ (first-order stochastic dominance) means ∑i≥qπi≥∑i≥qπi′\sum_{i\ge q}\pi_i\ge\sum_{i\ge q}\pi'_i∑i≥q​πi​≥∑i≥q​πi′​ for every qqq. π≥rπ′\pi\ge_r\pi'π≥r​π′ (MLR order) means πiπi′′≥πi′πi′\pi_i\pi'_{i'}\ge\pi_{i'}\pi'_iπi​πi′′​≥πi′​πi′​ whenever i≥i′i\ge i'i≥i′. For matrices f,gf,gf,g on X×YX\times YX×Y, f≥tpgf\ge_{tp}gf≥tp​g means f(x∨x′,y∨y′) g(x∧x′,y∧y′)≥f(x,y) g(x′,y′)f(x\vee x',y\vee y')\,g(x\wedge x',y\wedge y')\ge f(x,y)\,g(x',y')f(x∨x′,y∨y′)g(x∧x′,y∧y′)≥f(x,y)g(x′,y′) for all pairs, and fff is TP₂ if f≥tpff\ge_{tp}ff≥tp​f.

In each period the decision maker in state st=is_t=ist​=i chooses a∈Aa\in Aa∈A and receives the reward g(i,a)g(i,a)g(i,a). The state moves to jjj with probability pijap^a_{ij}pija​ (matrix PaP^aPa), and an observation kkk arrives with probability rjkar^a_{jk}rjka​ (matrix RaR^aRa, with row ra(j)∈Π(O)r^a(j)\in\Pi(O)ra(j)∈Π(O)), generated by the new state jjj and the action aaa. The paper assumes rjka>0r^a_{jk}>0rjka​>0 throughout. The discount factor is β≥0\beta\ge 0β≥0. From a belief π\piπ, the observation kkk has probability σ(k;π,a)=∑i,jπipijarjka\sigma(k;\pi,a)=\sum_{i,j}\pi_i p^a_{ij}r^a_{jk}σ(k;π,a)=∑i,j​πi​pija​rjka​, and the posterior is the Bayes update Tj(π,a,k)=∑iπipijarjka/σ(k;π,a)T_j(\pi,a,k)=\sum_i\pi_ip^a_{ij}r^a_{jk}/\sigma(k;\pi,a)Tj​(π,a,k)=∑i​πi​pija​rjka​/σ(k;π,a). With

h(π,a,V)=∑iπig(i,a)+β∑kσ(k;π,a) V(T(π,a,k)),h(\pi,a,V)=\sum_{i}\pi_i g(i,a)+\beta\sum_{k}\sigma(k;\pi,a)\,V(T(\pi,a,k)),h(π,a,V)=i∑​πi​g(i,a)+βk∑​σ(k;π,a)V(T(π,a,k)),

a finite horizon NNN with salvage value gsg_sgs​ gives the optimal values VN+1∗(π)=∑iπigs(i)V^*_{N+1}(\pi)=\sum_i\pi_ig_s(i)VN+1∗​(π)=∑i​πi​gs​(i) and Vt∗(π)=max⁡ah(π,a,Vt+1∗)V^*_t(\pi)=\max_a h(\pi,a,V^*_{t+1})Vt∗​(π)=maxa​h(π,a,Vt+1∗​). For N=∞N=\inftyN=∞ and 0<β<10<\beta<10<β<1, V∗V^*V∗ is the bounded solution of V∗(π)=max⁡ah(π,a,V∗)V^*(\pi)=\max_a h(\pi,a,V^*)V∗(π)=maxa​h(π,a,V∗). The myopic actions are α(π)=argmax⁡a∑iπig(i,a)\alpha(\pi)=\operatorname{argmax}_a\sum_i\pi_ig(i,a)α(π)=argmaxa​∑i​πi​g(i,a).

Formalization targets

Goal: Proposition 2 (myopic lower bound)

Under (a) gsg_sgs​ nondecreasing, (b) g(⋅,a)g(\cdot,a)g(⋅,a) nondecreasing, (c) Pa≥tpPa′P^a\ge_{tp}P^{a'}Pa≥tp​Pa′ for a≥a′a\ge a'a≥a′, (d) ra(j)≥rra(j′)r^a(j)\ge_r r^a(j')ra(j)≥r​ra(j′) for j≥j′j\ge j'j≥j′, (e) ra(j)≥sra′(j)r^a(j)\ge_s r^{a'}(j)ra(j)≥s​ra′(j) for a≥a′a\ge a'a≥a′, and (f) rjkarj′ka′≥rjka′rj′kar^a_{jk}r^{a'}_{j'k}\ge r^{a'}_{jk}r^a_{j'k}rjka​rj′ka′​≥rjka′​rj′ka​ for a≥a′a\ge a'a≥a′, j≥j′j\ge j'j≥j′: for every t≤Nt\le Nt≤N (finite horizon), or for the infinite horizon with 0<β<10<\beta<10<β<1, and every π∈Π(S)\pi\in\Pi(S)π∈Π(S),

∀ δ∗(π) ∃ α(π)≤δ∗(π),∀ α(π) ∃ δ∗(π)≥α(π),\forall\,\delta^*(\pi)\ \exists\,\alpha(\pi)\le\delta^*(\pi),\qquad \forall\,\alpha(\pi)\ \exists\,\delta^*(\pi)\ge\alpha(\pi),∀δ∗(π) ∃α(π)≤δ∗(π),∀α(π) ∃δ∗(π)≥α(π),

where δ∗(π)\delta^*(\pi)δ∗(π) ranges over the maximizers of a↦h(π,a,Vt+1∗)a\mapsto h(\pi,a,V^*_{t+1})a↦h(π,a,Vt+1∗​) (resp. h(π,a,V∗)h(\pi,a,V^*)h(π,a,V∗)).

Milestone: Proposition 1 (MLR-monotone values)

Under (a)–(d) with every PaP^aPa TP₂: π≥rπ′\pi\ge_r\pi'π≥r​π′ in Π(S)\Pi(S)Π(S) implies Vt∗(π)≥Vt∗(π′)V^*_t(\pi)\ge V^*_t(\pi')Vt∗​(π)≥Vt∗​(π′) for t=1,…,N+1t=1,\dots,N+1t=1,…,N+1, and, without (a), V∗(π)≥V∗(π′)V^*(\pi)\ge V^*(\pi')V∗(π)≥V∗(π′) for N=∞N=\inftyN=∞.

Supporting milestones

The ordering facts behind both propositions: MLR implies stochastic dominance (§1), Lemma 1.1 (characterization of ≥s\ge_s≥s​), Lemma 1.3 (TP₂ prediction preserves ≥r\ge_r≥r​), Lemma 1.2 (the Bayes update is MLR-monotone in the observation, the prior and the action), the stochastic monotonicity of σ\sigmaσ in the belief (proof of Proposition 1) and in the action (Lemma 2.3), the comparison of hhh-increments with myopic increments (proof of Proposition 2), and Lemma 2.2 (dominated increments order maximizer sets).

Significance

The result. Proposition 2 makes the myopic policy, which solves a one-stage problem, a lower bound on an optimal policy for every belief and every period. In a search over policies, actions below α(π)\alpha(\pi)α(π) can be discarded. When ggg also has isotone differences, α\alphaα is nondecreasing, and the optimal policy is bounded below by a monotone function that is easy to compute. Proposition 1 gives MLR-monotone value functions, the input to many later structural results for POMDPs. Lemma 1.2 records the fact behind it: Bayesian updating respects the MLR order, while first-order stochastic dominance does not survive conditioning.

Formalizing it. These results are proved on paper. This mission produces machine-checked proofs, together with a reusable finite-POMDP layer (belief update, observation probabilities, Bellman operator, finite- and infinite-horizon values) and a library of the stochastic orders on finite chains. One statement in the paper is wrong: the printed "only if" direction of Lemma 1.2(1) is false. The mission states only the direction that is true and used.

Difficulty

The obvious induction on ttt for Proposition 1 needs k↦V(T(π,a,k))k\mapsto V(T(\pi,a,k))k↦V(T(π,a,k)) to be nondecreasing and σ(π,a)\sigma(\pi,a)σ(π,a) to increase with π\piπ. Both need a belief order that conditioning preserves. Under first-order stochastic dominance the posterior is not monotone in the prior, and the paper's counterexample (p. 740) shows that the induction then fails. The MLR order repairs this, but proving that the prediction step preserves it (Lemma 1.3) requires a total-positivity composition argument (Karlin–Rinott, Theorem 2.4) on product lattices. For Proposition 2 the difficulty is to compare continuation values across actions: the observation distribution and the posterior both change with the action, and two separate orderings (Lemma 2.3 and Lemma 1.2(3)) must be combined before the maximizer comparison applies. The infinite-horizon parts additionally need the Bellman fixed point characterized well enough to pass monotonicity to the limit.

Formalization scope

Everything is finite, so all probabilities and expectations are finite sums and no measure theory is involved. States, observations and actions are finite nonempty types with a LinearOrder (any finite chain is isomorphic to {1,…,n}\{1,\dots,n\}{1,…,n}). Π(X)\Pi(X)Π(X) is Mathlib's stdSimplex ℝ X. The orders ≥s,≥r,≥tp\ge_s,\ge_r,\ge_{tp}≥s​,≥r​,≥tp​ are plain relations (StochGE, MLRGE, TPGE) with the larger argument first, and every statement assumes simplex membership explicitly. The standing assumptions (stochastic rows of PaP^aPa and RaR^aRa, rjka>0r^a_{jk}>0rjka​>0, β≥0\beta\ge0β≥0) are fields of the structure POMDP. Vt∗V^*_tVt∗​ is computed by recursion (3) counted in steps to go, Vstar gs N t = valueToGo gs (N+1-t). The infinite-horizon V∗V^*V∗ is any function bounded on Π(S)\Pi(S)Π(S) that solves the Bellman equation there. For 0<β<10<\beta<10<β<1 such a function exists and is unique on Π(S)\Pi(S)Π(S), by contraction. The equivalence between recursion (3) and the optimum over history-dependent strategies is cited by the paper from the literature and is not part of this mission. Maximizer sets (argmaxSet) carry the "for all δ∗\delta^*δ∗ / there exists α\alphaα" quantifiers. Both halves of each part of Proposition 2 are ∀∃\forall\exists∀∃ statements.

The goal is not to be read with Vt+1∗V^*_{t+1}Vt+1∗​ or V∗V^*V∗ replaced by an arbitrary, or an arbitrary nondecreasing, value function. That reading would reduce Proposition 2 to Lemma 2.2 plus a hypothesis. The goal quantifies only over the value functions of recursion (3) and over bounded Bellman solutions.

A complete development needs finite total-positivity composition (Mathlib's four functions theorem is the natural starting point), Abel summation for Lemma 1.1, and a contraction argument for the infinite horizon. The order library and the finite POMDP layer are reusable beyond this paper. Proofs of any milestone are welcome, as are alternative arguments for Lemma 1.3.

Selected references

  • W. S. Lovejoy, Some Monotonicity Results for Partially Observed Markov Decision Processes, Operations Research 35(5):736–743, 1987. https://doi.org/10.1287/opre.35.5.736
  • R. D. Smallwood and E. J. Sondik, The Optimal Control of Partially Observable Markov Processes over a Finite Horizon, Operations Research 21(5):1071–1088, 1973. https://doi.org/10.1287/opre.21.5.1071
  • W. Whitt, A Note on the Influence of the Sample on the Posterior Distribution, Journal of the American Statistical Association 74:424–426, 1979.
  • W. Whitt, Multivariate Monotone Likelihood Ratio and Uniform Conditional Stochastic Order, Journal of Applied Probability 19:695–701, 1982.
  • S. Karlin and Y. Rinott, Classes of Orderings of Measures and Related Correlation Inequalities. I. Multivariate Totally Positive Distributions, Journal of Multivariate Analysis 10(4):467–498, 1980. https://doi.org/10.1016/0047-259X(80)90065-2
  • C. White, Optimal Control-limit Strategies for a Partially Observed Replacement Problem, International Journal of Systems Science 10:321–331, 1979 (the machine-replacement model of §5).
  • S. C. Albright, Structural Results for Partially Observable Markov Decision Processes, Operations Research 27(5):1041–1053, 1979. https://doi.org/10.1287/opre.27.5.1041
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CombinatoricsOptimization·Captain: mikedeng1

Application of the Branch and Bound Technique to Some Flow-Shop Scheduling Problems 1: Branch and Bound with the Three-Machine Lower Bound Finds a Minimum-Makespan SequenceResearch Paper

Motivation

In a flow shop every job passes through the same machines in the same order. Choosing the order in which jobs enter the shop so that the last job finishes as early as possible (minimizing the makespan) is one of the oldest problems of scheduling theory. For two machines Johnson (1954) gave an exact sorting rule; for three machines his rule is exact only in special cases, and the general problem was later shown to be NP-hard (Garey, Johnson and Sethi, 1976). Exact methods for three or more machines therefore enumerate sequences, and the question is how to enumerate as few of them as possible.

Ignall and Schrage (1965), at the same time as Lomnicki (Operational Research Quarterly, 1965), introduced branch and bound for the three-machine permutation flow shop. Their lower bound, which adds to the time each machine becomes free the work that remains on it and the least work that must follow on later machines, became the template for the machine-based bounds used in flow-shop branch and bound ever since.

Timeline. 1954: Johnson proves the two-machine rule and shows that for three machines a common job order on all machines loses nothing for the makespan. 1960: Land and Doig introduce branch and bound for integer programs. 1965: Ignall–Schrage and Lomnicki apply it to the three-machine flow shop. 1976: Garey, Johnson and Sethi prove the three-machine makespan problem NP-hard, so exact enumeration with good bounds remains the method for exact solutions.

Setting

There are n≥1n\ge1n≥1 jobs and three machines AAA, BBB, CCC. Job iii needs time aia_iai​ on AAA, then bib_ibi​ on BBB, then cic_ici​ on CCC. Following Johnson, only permutation schedules are considered: a full sequence is a permutation σ\sigmaσ of the jobs, processed in that order on all three machines, each operation starting as early as possible. Its makespan is the time machine CCC finishes the last job.

A node JrJ_rJr​ is a partial sequence: an ordered list of rrr distinct jobs. Its unscheduled set Jˉr\bar J_rJˉr​ is the set of the other n−rn-rn−r jobs. The attributes TIMEA(Jr)\mathrm{TIMEA}(J_r)TIMEA(Jr​), TIMEB(Jr)\mathrm{TIMEB}(J_r)TIMEB(Jr​), TIMEC(Jr)\mathrm{TIMEC}(J_r)TIMEC(Jr​) are the times at which machines AAA, BBB, CCC finish the jobs of JrJ_rJr​ processed in order. A full sequence begins with JrJ_rJr​ when its first rrr positions are JrJ_rJr​. The lower bound of a node with Jˉr≠∅\bar J_r\neq\emptysetJˉr​=∅ is

LB(Jr)=max⁡[TIMEA(Jr)+∑Jˉrai+min⁡Jˉr(bi+ci), TIMEB(Jr)+∑Jˉrbi+min⁡Jˉrci, TIMEC(Jr)+∑Jˉrci].LB(J_r)=\max\Big[\mathrm{TIMEA}(J_r)+\sum_{\bar J_r}a_i+\min_{\bar J_r}(b_i+c_i),\ \mathrm{TIMEB}(J_r)+\sum_{\bar J_r}b_i+\min_{\bar J_r}c_i,\ \mathrm{TIMEC}(J_r)+\sum_{\bar J_r}c_i\Big].LB(Jr​)=max[TIMEA(Jr​)+Jˉr​∑​ai​+Jˉr​min​(bi​+ci​), TIMEB(Jr​)+Jˉr​∑​bi​+Jˉr​min​ci​, TIMEC(Jr​)+Jˉr​∑​ci​].

The branch-and-bound procedure keeps a list of nodes ranked by LBLBLB, starting from the root (no job scheduled). At each step it removes the first node, creates one child per unscheduled job by appending that job, and inserts the children ranked into the list, a new node going before old nodes of equal bound. It stops when the first node of the list is terminal, i.e. has scheduled n−1n-1n−1 jobs, so that its last job is forced.

Formalization targets

Goal: the stopping rule certifies an optimal sequence

∃k: the first node of run(k) is terminal,andfirst node P terminal ⟹ makespan(σP)≤makespan(σ)  ∀σ,\exists k:\ \text{the first node of } \mathrm{run}(k) \text{ is terminal},\qquad\text{and}\qquad \text{first node } P \text{ terminal} \ \Longrightarrow\ \mathrm{makespan}(\sigma_P)\le\mathrm{makespan}(\sigma)\ \ \forall\sigma ,∃k: the first node of run(k) is terminal,andfirst node P terminal ⟹ makespan(σP​)≤makespan(σ)  ∀σ,

where σP\sigma_PσP​ is PPP followed by its one unscheduled job. This is the paper's "the problem is solved: that node's sequence is an optimal one" (p. 402), for any real processing times.

Milestones

  1. Validity of the bound (p. 401): LB(Jr)≤makespan(σ)LB(J_r)\le\mathrm{makespan}(\sigma)LB(Jr​)≤makespan(σ) for every σ\sigmaσ beginning with JrJ_rJr​, r<nr<nr<n.
  2. Exactness at depth n−1n-1n−1 (general form of an observation on p. 403): for r=n−1r=n-1r=n−1, LB(Jn−1)LB(J_{n-1})LB(Jn−1​) equals the makespan of the completed sequence.
  3. Node counts (p. 403): at least 12n(n+1)\tfrac12n(n+1)21​n(n+1) nodes are created before stopping, at most 1+n+n(n−1)+⋯+n!1+n+n(n-1)+\cdots+n!1+n+n(n−1)+⋯+n! are ever created, and the list never holds more than n!n!n! nodes.
  4. Dominance (pp. 403–404): if JrJ_rJr​ and IrI_rIr​ contain the same jobs, TIMEA(Jr)=TIMEA(Ir)\mathrm{TIMEA}(J_r)=\mathrm{TIMEA}(I_r)TIMEA(Jr​)=TIMEA(Ir​); if also TIMEB(Jr)≤TIMEB(Ir)\mathrm{TIMEB}(J_r)\le\mathrm{TIMEB}(I_r)TIMEB(Jr​)≤TIMEB(Ir​) and TIMEC(Jr)≤TIMEC(Ir)\mathrm{TIMEC}(J_r)\le\mathrm{TIMEC}(I_r)TIMEC(Jr​)≤TIMEC(Ir​), replacing IrI_rIr​ by JrJ_rJr​ at the beginning of any sequence does not increase its makespan, and LB(Jr)≤LB(Ir)LB(J_r)\le LB(I_r)LB(Jr​)≤LB(Ir​) (p. 404).
  5. The 4-job example (pp. 402–403): the run stops after 6 steps with node 231 first, whose bound 62 is the optimal makespan.
  6. The refined bound (p. 409): the strengthened bound used in the paper's computations is still a lower bound.

Significance

The goal is the correctness theorem of the Ignall–Schrage algorithm: best-first search over partial sequences, with a machine-based bound, returns an optimal permutation. Milestone 1 is the bound's validity, which every later machine-based flow-shop bound generalizes; milestone 4 is the dominance rule that justifies discarding nodes; milestone 3 brackets the effort of the search between the best case 12n(n+1)\tfrac12n(n+1)21​n(n+1) and full enumeration.

The paper's arguments are short and informal: validity of the bound is asserted with a one-line reason, and optimality at stopping is argued on the example. None of these results has a machine-checked proof. The formalization makes the procedure itself a precise object (list, insertion rule, stopping rule) and turns the paper's claims into statements about it, so that the correctness of the search is proved rather than illustrated.

Difficulty

Each inequality is elementary, but the goal is a statement about an iterated list-manipulating procedure. The obvious argument "the first node has the smallest bound and the bound is valid" needs an invariant that the paper does not state: at every step, every full sequence begins with some node on the list, and the list is sorted. Both must be proved to survive one step of removal and ranked insertion. Termination is likewise not stated: it needs a measure that strictly decreases, for instance the set of nodes that remain to be created, which requires showing that no node is created twice. A tempting shortcut, proving optimality only among the sequences whose nodes were created, is not the theorem.

Formalization scope

Jobs are Fin n (the paper's job iii is i−1i-1i−1), processing times are arbitrary reals with no sign condition, and full sequences are Equiv.Perm (Fin n). The makespan is the machine-CCC completion time of Johnson's as-soon-as-possible schedule, taken from the published definition JohnsonFlowShop.ThreeStage.asapSchedule. Nodes are lists of jobs, and the node attributes are the same recursion folded over the list. Minima over Jˉr\bar J_rJˉr​ are Finset.inf'; on an empty set the auxiliary minimum returns a placeholder 000 that no statement evaluates.

Explicit readings of the paper's phrases:

  • "is a lower bound … for any node that emanates from node PPP" is "≤\le≤ the makespan of every permutation whose first rrr positions are JrJ_rJr​";
  • "a node that has scheduled all nnn jobs" is read as a node with n−1n-1n−1 jobs, whose sequence is completed by the forced last job. The example (stop at node 231 of a 4-job problem), the node counts and the formula for LBLBLB all require this reading;
  • "the problem is solved: that node's sequence is an optimal one" is termination together with optimality over all n!n!n! permutations;
  • "cannot be hurt by replacing IrI_rIr​ with JrJ_rJr​" compares the completions of JrJ_rJr​ and IrI_rIr​ by the same ordering of the remaining jobs;
  • children are inserted one at a time in increasing job index, each before every node with an equal bound; this reproduces the paper's LIST table;
  • dominance discarding is not part of the procedure in the goal.

A formalization in which the optimality claim ranges only over sequences the run created, or one that asserts optimality without termination, would be trivial or vacuous and is ruled out by the goal's statement.

Not formalized: the reduction from general schedules to permutation schedules (Johnson's Lemma 3, proved on the platform as JohnsonFlowShop.ThreeStage.same_ordering_dominant), the dominance bookkeeping and its percentages, the two-machine mean-completion problem (a separate mission of this series), the computational tables, and the TIMEB/TIMEC columns of the example's LIST table. No platform item treats flow-shop lower bounds or branch and bound over sequences; the nearest branch-and-bound mission, BiconvexProg.BranchBound (Al-Khayyal and Falk), concerns biconvex programs and is unrelated.

The lemmas a solver will want (the invariant of the list, monotonicity of the completion recursion in its start times, validity of the bound) are reusable for the two-machine mission and for any other machine-based bound. Proofs of the milestones, and of the goal from them, are welcome in any order.

Selected references

  • E. Ignall and L. Schrage, Application of the branch and bound technique to some flow-shop scheduling problems, Operations Research 13(3), 400–412, 1965. https://doi.org/10.1287/opre.13.3.400
  • S. M. Johnson, Optimal two- and three-stage production schedules with setup times included, Naval Research Logistics Quarterly 1(1), 61–68, 1954. https://doi.org/10.1002/nav.3800010110
  • A. H. Land and A. G. Doig, An automatic method of solving discrete programming problems, Econometrica 28(3), 497–520, 1960. https://doi.org/10.2307/1910129
  • M. R. Garey, D. S. Johnson and R. Sethi, The complexity of flowshop and jobshop scheduling, Mathematics of Operations Research 1(2), 117–129, 1976. https://doi.org/10.1287/moor.1.2.117
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Algorithmic Game TheoryMechanism Design·Captain: mikedeng1

Job Matching, Coalition Formation, and Gross Substitutes 1: Under Gross Substitutes the Salary-Adjustment Process Reaches a Discrete Core Allocation in Finitely Many RoundsResearch Paper

Motivation

Labor markets match workers to firms, and the terms of each match (the salary) are negotiated along with the match itself. A firm's output depends on the whole team it hires, so a firm cares about sets of workers, not individual workers one at a time. Kelso and Crawford (1982) asked when such a market has a stable outcome, the core, in which no firm and group of workers can agree on terms that all of them prefer, and when a simple decentralized auction finds one.

Their answer is the gross-substitutes condition: raising some workers' salaries never makes a firm withdraw an offer from a worker whose salary has not risen. Under it, an ascending process in which firms make offers and workers reject all but their favorite reaches a core allocation. This condition became the standard hypothesis for the existence of Walrasian equilibrium with indivisible goods (Gul and Stacchetti 1999), it is the hypothesis behind matching with contracts (Hatfield and Milgrom 2005), and the process is the ancestor of ascending auction designs.

Timeline.

  • 1962: Gale and Shapley, deferred acceptance for one-to-one and many-to-one matching without money.
  • 1971: Shapley and Shubik, the assignment game: one-to-one matching with transferable utility; the core is nonempty and is the set of solutions of a dual linear program.
  • 1981: Crawford and Knoer, a salary-adjustment process for one-to-one matching with money, converging to the core.
  • 1982: Kelso and Crawford, many-to-one matching with money and production complementarities; existence of the core under gross substitutes via the salary-adjustment process (Theorem 1, the target of this mission).

Setting

There are finitely many workers i∈Wi \in Wi∈W and firms j∈Fj \in Fj∈F, with at least one firm. Worker iii's utility of working for firm jjj at salary sss is ui(j;s)u^i(j; s)ui(j;s), strictly increasing and continuous in sss. Firm jjj's gross product from hiring the set C⊆WC \subseteq WC⊆W is yj(C)y^j(C)yj(C), and its profit at the salary vector sj=(s1j,…,smj)s^j = (s_{1j},\dots,s_{mj})sj=(s1j​,…,smj​) is

πj(C;sj)=yj(C)−∑i∈Csij.\pi^j(C; s^j) = y^j(C) - \sum_{i \in C} s_{ij}.πj(C;sj)=yj(C)−i∈C∑​sij​.

Mj(sj)M^j(s^j)Mj(sj) is the set of profit-maximizing CCC. Each pair has a starting salary σij\sigma_{ij}σij​. The assumptions (p. 1486) are (MP) yj(C∪{i})−yj(C)−σij≥0y^j(C \cup \{i\}) - y^j(C) - \sigma_{ij} \ge 0yj(C∪{i})−yj(C)−σij​≥0 for i∉Ci \notin Ci∈/C; (NFL) yj(∅)=0y^j(\emptyset) = 0yj(∅)=0; and (GS): if C∈Mj(sj)C \in M^j(s^j)C∈Mj(sj) and s~j≥sj\tilde s^j \ge s^js~j≥sj, some C~∈Mj(s~j)\tilde C \in M^j(\tilde s^j)C~∈Mj(s~j) contains {i∈C:s~ij=sij}\{i \in C : \tilde s_{ij} = s_{ij}\}{i∈C:s~ij​=sij​}.

In the discrete market with unit δ>0\delta > 0δ>0, firm jjj may pay worker iii only σij+kδ\sigma_{ij} + k\deltaσij​+kδ, k=0,1,2,…k = 0, 1, 2, \dotsk=0,1,2,…. An allocation sends each worker iii to a firm f(i)f(i)f(i) at a salary sif(i)s_{if(i)}sif(i)​; it is individually rational (D1) if sif(i)≥σif(i)s_{if(i)} \ge \sigma_{if(i)}sif(i)​≥σif(i)​ and every firm's profit is nonnegative. It is a (discrete) core allocation (D3) if it is individually rational, pays permitted salaries, and no firm jjj, set CCC and permitted salaries rjr^jrj satisfy ui(j;rij)>ui(f(i);sif(i))u^i(j; r_{ij}) > u^i(f(i); s_{if(i)})ui(j;rij​)>ui(f(i);sif(i)​) for all i∈Ci \in Ci∈C and πj(C;rj)>πj(Cj;sj)\pi^j(C; r^j) > \pi^j(C^j; s^j)πj(C;rj)>πj(Cj;sj).

The salary-adjustment process (pp. 1488–1489), verbatim:

R1. Firms begin facing a set of permitted salaries sij(0)=σijs_{ij}(0) = \sigma_{ij}sij​(0)=σij​. Permitted salaries at round ttt, sij(t)s_{ij}(t)sij​(t), remain constant, except as noted below. In round zero, each firm makes offers to all workers; this is costless by (MP).

R2. On each round, each firm makes offers to the members of one of its favorite sets of workers, given the schedule of permitted salaries sj(t)≡[s1j(t),…,smj(t)]s^j(t) \equiv [s_{1j}(t), \dots, s_{mj}(t)]sj(t)≡[s1j​(t),…,smj​(t)]. That is, firm jjj makes offers to the members of Cj[sj(t)]C^j[s^j(t)]Cj[sj(t)], where Cj[sj(t)]C^j[s^j(t)]Cj[sj(t)] maximizes πj[C;sj(t)]\pi^j[C; s^j(t)]πj[C;sj(t)]. Firms may break ties between sets of workers however they like, with the following exception: Any offer made by firm jjj in round t−1t - 1t−1 that was not rejected must be repeated in round ttt. By (GS), the firm sacrifices no profits in doing this, since (by R4) other workers' permitted salaries cannot have fallen, and the salary of a worker who did not reject an offer remains constant.

R3. Each worker who receives one or more offers rejects all but his or her favorite (taking salaries into account), which he or she tentatively accepts. Workers may break ties at any time however they like.

R4. Offers not rejected in previous periods remain in force. If worker iii rejected an offer from firm jjj in round t−1t - 1t−1, sij(t)=sij(t−1)+1s_{ij}(t) = s_{ij}(t - 1) + 1sij​(t)=sij​(t−1)+1; otherwise sij(t)=sij(t−1)s_{ij}(t) = s_{ij}(t - 1)sij​(t)=sij​(t−1). Firms continue to make offers to their favorite sets of workers, taking into account their permitted salaries.

R5. The process stops when no rejections are issued in some period. Workers then accept the offers that remain in force from the firms they have not rejected.

Formalization targets

Goal: Theorem 1 (p. 1489)

"The salary-adjustment process R1–R5 converges in finite time to a discrete core allocation in the discrete market for which it is defined." Formally, under the assumptions above:

(∃ a run) ∧ ∀ρ run: (∃T: ρ issues no rejections in round T) ∧ (∀T such rounds, outcomeρ(T) is a discrete core allocation).\Big(\exists \text{ a run}\Big)\ \wedge\ \forall \rho \text{ run}:\ \Big(\exists T:\ \rho \text{ issues no rejections in round } T\Big)\ \wedge\ \Big(\forall T \text{ such rounds},\ \text{outcome}_\rho(T) \text{ is a discrete core allocation}\Big).(∃ a run) ∧ ∀ρ run: (∃T: ρ issues no rejections in round T) ∧ (∀T such rounds, outcomeρ​(T) is a discrete core allocation).

Milestones, in the paper's order

  1. R2 is well defined: a run exists (each firm has a favorite set containing its unrejected offers).
  2. Lemma 1: every worker has at least one offer in every period.
  3. Lemma 2: after finitely many rounds every worker has exactly one offer and the process stops.
  4. Lemma 3: the allocation at a stopping round is individually rational.
  5. Lemma 4: the allocation at a stopping round is a discrete core allocation.

Significance

Theorem 1 gives the existence of a core allocation in every discrete market satisfying (MP), (NFL) and (GS), with no convexity of production and arbitrary complementarity within the limits of (GS). It is the step from which the paper derives the existence of a strict core allocation of the continuous market (Theorem 2, by letting the unit shrink), and with additional no-ties assumptions the firm-optimality of the process's outcome (Theorem 4) and the comparative statics of entry and exit (Theorem 5).

The result is proved in the paper and has been reproved in more general settings; it has no machine-checked proof known to us. Nothing of this paper was on Prove2Me before this series. Related platform work, credited but not reused: the Gale–Shapley deferred-acceptance development (GS62CollegeAdmissions.*, proved), the no-money ancestor of this process; and the Shapley–Shubik assignment game (AssignmentGame.CoreLP), whose core the process approximates when production is additively separable. Neither states anything about this process. This mission is mission 1 of a series of seven on the paper: 2 (strict core of the continuous market), 3 (one-sided coalition formation), 4 (firm-optimality), 5 (comparative statics), 6 (gross substitutes and decreasing returns), 7 (a market without a core). Each states its own model.

Difficulty

The process is quantified over all tie-breakings, so no single computation settles it; the statements are about every sequence of rounds consistent with R1–R5. Two points carry the content. First, R2 imposes a constraint that may not be satisfiable: a firm must repeat its unrejected offers and choose a profit-maximizing set. Without (GS) no such set need exist and the process is not defined. Second, the core property at the stopping round compares the outcome with coalitions using salaries the process never reached; the comparison is with salaries on the discrete grid only, and it is false if coalitions may use salaries below σij\sigma_{ij}σij​ or off the grid. Termination is not automatic either: the process may continue after a round without rejections, salaries are real numbers, and no bound on the number of rounds is given.

Formalization scope

Lean representation, in namespace KelsoCrawford.Process:

  • Market W F carries u, y, σ; workers and firms are finite types, with [Nonempty F] (the paper's n≥1n \ge 1n≥1; with no firms and some worker no run exists).
  • Allocation assigns every worker to a firm (no unemployment, as in the paper's fff).
  • (GS) is GrossSubstitutesOn (M.y j) (M.gridVectors δ j) for each firm: the discrete (GS), since the paper notes that a discrete market may satisfy (GS) while its continuous version does not.
  • The core is D3 with permitted salaries M.grid δ ={σij+kδ:k∈N}= \{\sigma_{ij} + k\delta : k \in \mathbb N\}={σij​+kδ:k∈N}.
  • A Run records salaries, offers and tentative choices per round; IsRun is R1–R4, Stopped is R5's "no rejections", outcome is R5's allocation.

Explicit readings of the paper's phrases:

  • "converges in finite time" = every run has a round without rejections; no bound on that round is claimed;
  • "to a discrete core allocation" = at every round without rejections, the allocation read off is in the discrete core;
  • the unit 111 of R4 is a parameter δ>0\delta > 0δ>0 (same theorem in rescaled units);
  • σij\sigma_{ij}σij​ is data; its defining relation ui(j;σij)=ui(0;0)u^i(j;\sigma_{ij}) = u^i(0;0)ui(j;σij​)=ui(0;0) is not assumed (a more general statement).

The existence of a run is part of the goal: without it, statements about all runs would be vacuous. Fixing a tie-breaking rule would turn the statements into claims about a single run and is ruled out; so are the strict core D2 (false here because of ties at the grid) and an integer grid below σij\sigma_{ij}σij​ (false for improving coalitions). Proofs of the milestones and reusable infrastructure for ascending processes are welcome.

Selected references

  • A. S. Kelso, Jr. and V. P. Crawford, Job matching, coalition formation, and gross substitutes, Econometrica 50(6), 1982, 1483–1504. https://doi.org/10.2307/1913392
  • V. P. Crawford and E. M. Knoer, Job matching with heterogeneous firms and workers, Econometrica 49(2), 1981, 437–450. https://doi.org/10.2307/1913320
  • D. Gale and L. S. Shapley, College admissions and the stability of marriage, American Mathematical Monthly 69(1), 1962, 9–15. https://doi.org/10.2307/2312726
  • L. S. Shapley and M. Shubik, The assignment game I: The core, International Journal of Game Theory 1, 1971, 111–130. https://doi.org/10.1007/BF01753437
  • F. Gul and E. Stacchetti, Walrasian equilibrium with gross substitutes, Journal of Economic Theory 87(1), 1999, 95–124. https://doi.org/10.1006/jeth.1999.2531
  • J. W. Hatfield and P. R. Milgrom, Matching with contracts, American Economic Review 95(4), 2005, 913–935. https://doi.org/10.1257/0002828054825466
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Optimization·Captain: mikedeng1

The Fritz John Necessary Optimality Conditions in the Presence of Equality and Inequality Constraints: Every Minimizer of a C¹ Program Admits Multipliers (ū₀, ū, v̄) ≠ 0 with ū ≥ 0Research Paper

Motivation

For problems with inequality constraints only, F. John (1948) showed that every minimizer admits nonnegative multipliers (uˉ0,uˉ1,…,uˉm)≠0(\bar u_0, \bar u_1, \dots, \bar u_m) \ne 0(uˉ0​,uˉ1​,…,uˉm​)=0, one of which belongs to the objective. The Kuhn–Tucker conditions (1951) are the stronger statement in which the objective's multiplier can be taken equal to one; they need a constraint qualification.

Problems arising in practice mix equalities and inequalities. Fritz John's theorem does not cover them, and the obvious reduction, writing each equality hj(x)=0h_j(x) = 0hj​(x)=0 as the two inequalities hj(x)≤0h_j(x) \le 0hj​(x)≤0 and −hj(x)≤0-h_j(x) \le 0−hj​(x)≤0, destroys the content of the conditions: every feasible point then satisfies them with uˉ0=0\bar u_0 = 0uˉ0​=0. O. L. Mangasarian and S. Fromovitz (1967) proved a version of Fritz John's conditions that treats equalities directly and stays informative, and from it derived the constraint qualification now known as the Mangasarian–Fromovitz constraint qualification (MFCQ). MFCQ is the standard regularity assumption in the convergence theory of sequential quadratic programming, interior-point and augmented-Lagrangian methods, and in the stability theory of parametric programs.

Timeline.

  • 1939, W. Karush (master's thesis) and 1951, H. W. Kuhn and A. W. Tucker: multiplier conditions with uˉ0=1\bar u_0 = 1uˉ0​=1 for inequality constraints, under a constraint qualification.
  • 1948, F. John: the multiplier rule with uˉ0≥0\bar u_0 \ge 0uˉ0​≥0 for inequality constraints, no qualification needed.
  • 1967, Mangasarian and Fromovitz (this paper): the multiplier rule for equalities and inequalities together, and the qualification (3.4)–(3.6).

Setting

Let EnE^nEn be nnn-dimensional Euclidean space, and let θ,g1,…,gm,h1,…,hk:En→R\theta, g_1, \dots, g_m, h_1, \dots, h_k : E^n \to \mathbb Rθ,g1​,…,gm​,h1​,…,hk​:En→R be functions with continuous first partial derivatives on EnE^nEn. The program is

minimize θ(x)subject togi(x)≤0, i∈M={1,…,m},hj(x)=0, j∈K={1,…,k}.(1.1)\text{minimize } \theta(x) \quad \text{subject to} \quad g_i(x) \le 0,\ i \in M = \{1,\dots,m\}, \qquad h_j(x) = 0,\ j \in K = \{1,\dots,k\}. \tag{1.1}minimize θ(x)subject togi​(x)≤0, i∈M={1,…,m},hj​(x)=0, j∈K={1,…,k}.(1.1)

The feasible set is S={x∈En:gi(x)≤0, i∈M, hj(x)=0, j∈K}S = \{x \in E^n : g_i(x) \le 0,\ i \in M,\ h_j(x) = 0,\ j \in K\}S={x∈En:gi​(x)≤0, i∈M, hj​(x)=0, j∈K}. A point xˉ\bar xxˉ is a solution of (1.1) if xˉ∈S\bar x \in Sxˉ∈S and θ(xˉ)≤θ(x)\theta(\bar x) \le \theta(x)θ(xˉ)≤θ(x) for all x∈Sx \in Sx∈S. The active set at xˉ\bar xxˉ is Mˉ={i∈M:gi(xˉ)=0}\bar M = \{i \in M : g_i(\bar x) = 0\}Mˉ={i∈M:gi​(xˉ)=0}. The gradient of fff at xˉ\bar xxˉ is ∇f(xˉ)\nabla f(\bar x)∇f(xˉ), and y′zy'zy′z denotes the inner product.

The generalized Fritz John conditions hold at xˉ\bar xxˉ if there are uˉ=(uˉ0,uˉ1,…,uˉm)\bar u = (\bar u_0, \bar u_1, \dots, \bar u_m)uˉ=(uˉ0​,uˉ1​,…,uˉm​) and vˉ=(vˉ1,…,vˉk)\bar v = (\bar v_1, \dots, \bar v_k)vˉ=(vˉ1​,…,vˉk​) with

uˉ0∇θ(xˉ)+∑i=1muˉi∇gi(xˉ)+∑j=1kvˉj∇hj(xˉ)=0,∑i=1muˉigi(xˉ)=0,uˉ≥0,(uˉ,vˉ)≠0.\bar u_0 \nabla\theta(\bar x) + \sum_{i=1}^m \bar u_i \nabla g_i(\bar x) + \sum_{j=1}^k \bar v_j \nabla h_j(\bar x) = 0, \qquad \sum_{i=1}^m \bar u_i g_i(\bar x) = 0, \qquad \bar u \ge 0, \qquad (\bar u, \bar v) \ne 0 .uˉ0​∇θ(xˉ)+i=1∑m​uˉi​∇gi​(xˉ)+j=1∑k​vˉj​∇hj​(xˉ)=0,i=1∑m​uˉi​gi​(xˉ)=0,uˉ≥0,(uˉ,vˉ)=0.

The Kuhn–Tucker conditions are the same system with uˉ0=1\bar u_0 = 1uˉ0​=1 and no nontriviality requirement.

Formalization targets

Goal: the generalized Fritz John necessary conditions (p. 41)

If xˉ\bar xxˉ is a solution of (1.1), then there exist uˉ∈Em+1\bar u \in E^{m+1}uˉ∈Em+1 and vˉ∈Ek\bar v \in E^kvˉ∈Ek with

uˉ0∇θ(xˉ)+∑i=1muˉi∇gi(xˉ)+∑j=1kvˉj∇hj(xˉ)=0,∑i=1muˉigi(xˉ)=0,uˉ≥0,(uˉ,vˉ)≠0.(2.9–2.12)\bar u_0 \nabla\theta(\bar x) + \sum_{i=1}^m \bar u_i \nabla g_i(\bar x) + \sum_{j=1}^k \bar v_j \nabla h_j(\bar x) = 0, \quad \sum_{i=1}^m \bar u_i g_i(\bar x) = 0, \quad \bar u \ge 0, \quad (\bar u, \bar v) \ne 0. \tag{2.9–2.12}uˉ0​∇θ(xˉ)+i=1∑m​uˉi​∇gi​(xˉ)+j=1∑k​vˉj​∇hj​(xˉ)=0,i=1∑m​uˉi​gi​(xˉ)=0,uˉ≥0,(uˉ,vˉ)=0.(2.9–2.12)

No regularity of the constraints is assumed. The nontriviality requirement covers uˉ0\bar u_0uˉ0​, the uˉi\bar u_iuˉi​ and the vˉj\bar v_jvˉj​ together.

Milestones

  1. Motzkin's transposition theorem (p. 39): for real matrices A,B,CA, B, CA,B,C with AAA nonempty, exactly one of y′A<0, y′B≤0, y′C=0y'A < 0,\ y'B \le 0,\ y'C = 0y′A<0, y′B≤0, y′C=0 and Az1+Bz2+Cz3=0, z1≥0, z1≠0, z2≥0Az_1 + Bz_2 + Cz_3 = 0,\ z_1 \ge 0,\ z_1 \ne 0,\ z_2 \ge 0Az1​+Bz2​+Cz3​=0, z1​≥0, z1​=0, z2​≥0 is solvable.
  2. Lemma 1 (pp. 39–40): if fi(xˉ)=0f_i(\bar x) = 0fi​(xˉ)=0, hj(xˉ)=0h_j(\bar x) = 0hj​(xˉ)=0 at some xˉ\bar xxˉ in an open set DDD, no x∈Dx \in Dx∈D has fi(x)<0f_i(x) < 0fi​(x)<0 for all iii and hj(x)=0h_j(x) = 0hj​(x)=0 for all jjj, and the ∇hj(xˉ)\nabla h_j(\bar x)∇hj​(xˉ) are linearly independent, then no yyy has y′∇fi(xˉ)<0y'\nabla f_i(\bar x) < 0y′∇fi​(xˉ)<0 and y′∇hj(xˉ)=0y'\nabla h_j(\bar x) = 0y′∇hj​(xˉ)=0.
  3. Lemma 2 (p. 40): under the same assumptions, without independence, there are rˉ≥0\bar r \ge 0rˉ≥0 and sˉ\bar ssˉ, not both zero, with ∑rˉi∇fi(xˉ)+∑sˉj∇hj(xˉ)=0\sum \bar r_i \nabla f_i(\bar x) + \sum \bar s_j \nabla h_j(\bar x) = 0∑rˉi​∇fi​(xˉ)+∑sˉj​∇hj​(xˉ)=0.
  4. DDD is open (p. 41): D={x:gi(x)<0, i∈M∖Mˉ}D = \{x : g_i(x) < 0,\ i \in M \setminus \bar M\}D={x:gi​(x)<0, i∈M∖Mˉ} is open.
  5. The reduction (pp. 41–42): at a solution xˉ\bar xxˉ of (1.1), xˉ∈D\bar x \in Dxˉ∈D and the system θ(x)−θ(xˉ)<0\theta(x) - \theta(\bar x) < 0θ(x)−θ(xˉ)<0, gi(x)<0g_i(x) < 0gi​(x)<0 (i∈Mˉi \in \bar Mi∈Mˉ), hj(x)=0h_j(x) = 0hj​(x)=0 has no solution in DDD.

Companion results

  • Corollary (p. 43): the generalized Fritz John conditions hold at any feasible point satisfying (2.27) or (2.28).
  • The generalized constraint qualification (pp. 43–44): at a solution, yˉ′∇gi(xˉ)<0\bar y'\nabla g_i(\bar x) < 0yˉ​′∇gi​(xˉ)<0 (i∈Mˉi \in \bar Mi∈Mˉ), yˉ′∇hj(xˉ)=0\bar y'\nabla h_j(\bar x) = 0yˉ​′∇hj​(xˉ)=0 and independent ∇hj(xˉ)\nabla h_j(\bar x)∇hj​(xˉ) imply the Kuhn–Tucker conditions.
  • The splitting remark (p. 38): after splitting equalities, every feasible point satisfies Fritz John's original conditions.

Significance

The theorem is a multiplier rule for smooth programs with both kinds of constraints and no assumption on the constraints. It has two direct consequences in the paper. First, it yields MFCQ, the condition (3.4)–(3.6) under which the Kuhn–Tucker conditions hold at every solution. MFCQ is weaker than linear independence of all active gradients (LICQ), and it is equivalent to boundedness of the Kuhn–Tucker multiplier set (Gauvin, 1977). Second, the corollary identifies feasible non-minimizers at which the conditions hold anyway.

The result is classical and its proof is in every nonlinear-programming textbook. Mathlib has the equality-constrained Lagrange multiplier rule for a local extremum (IsLocalExtrOn.exists_multipliers_of_hasStrictFDerivAt) and the implicit function theorem, and Prove2Me has a formalized Fritz John theorem for inequality constraints only. To our knowledge no machine-checked proof of the mixed equality–inequality Fritz John rule, of MFCQ, or of Motzkin's transposition theorem in this form exists. The formalization would provide the standard multiplier rule and qualification on which a formal theory of nonlinear programming builds.

Difficulty

With inequalities alone, John's theorem follows from the observation that if xˉ\bar xxˉ is a minimizer, no direction yyy strictly decreases θ\thetaθ and every active gig_igi​ to first order; Motzkin's (or Gordan's) theorem then produces the multipliers. With equalities the first step fails: a direction with y′∇hj(xˉ)=0y'\nabla h_j(\bar x) = 0y′∇hj​(xˉ)=0 is tangent to the equality manifold but generally leaves it, so a first-order descent direction does not yield a feasible point with smaller objective. Lemma 1 is precisely the claim that it does when the ∇hj(xˉ)\nabla h_j(\bar x)∇hj​(xˉ) are independent, and it needs a curve inside {h=0}\{h = 0\}{h=0} along which the strict inequalities persist: the implicit function theorem, applied on an open set, with care that the curve stays in DDD. The linearly dependent case must be handled separately, and it is the only place the multipliers vˉ\bar vvˉ can be nonzero with uˉ=0\bar u = 0uˉ=0.

Formalization scope

EnE^nEn is EuclideanSpace ℝ (Fin n), the inner product y′zy'zy′z is inner ℝ y z, and ∇f(xˉ)\nabla f(\bar x)∇f(xˉ) is Mathlib's gradient f xbar. "Continuous first partial derivatives on EnE^nEn", the standing assumption of §1, is ContDiff ℝ 1 (equivalent in finite dimension) and is a hypothesis of the goal, the corollary and the constraint qualification; Lemma 1 and Lemma 2 use ContDiffOn ℝ 1 · D on an open set DDD, as on the page. Indices i∈Mi \in Mi∈M and j∈Kj \in Kj∈K are Fin m and Fin k, 0-based; m=0m = 0m=0 and k=0k = 0k=0 are allowed. The multiplier vector uˉ∈Em+1\bar u \in E^{m+1}uˉ∈Em+1 is split into u0 : ℝ and u : Fin m → ℝ, and (uˉ,vˉ)≠0(\bar u, \bar v) \ne 0(uˉ,vˉ)=0 is "u0 ≠ 0, or some u i ≠ 0, or some v j ≠ 0". A solution of (1.1) is a global minimizer over SSS, as the proof on p. 42 uses. In Motzkin's theorem a matrix is the family of its columns, and "either … or …, but never both" is Xor. In Lemma 2, "the assumptions of Lemma 1" exclude the proviso (2.5) of linear independence, which the proof of Lemma 2 treats separately.

The goal does not assume linear independence of the ∇hj(xˉ)\nabla h_j(\bar x)∇hj​(xˉ), does not mention DDD or Lemma 1, and requires (uˉ,vˉ)≠0(\bar u, \bar v) \ne 0(uˉ,vˉ)=0 with uˉ0\bar u_0uˉ0​ included; a formalization that drops uˉ0\bar u_0uˉ0​ from the nontriviality condition is false at m=k=0m = k = 0m=k=0, and one that requires uˉ0≠0\bar u_0 \ne 0uˉ0​=0 is the Kuhn–Tucker statement, false without a qualification.

A complete development needs Motzkin's (or Gordan's) theorem of the alternative, which is reusable well beyond this mission, and the implicit function theorem on open sets in Euclidean space with a C1C^1C1 curve argument. Proofs of any milestone are welcome, as is a proof of the goal by another route.

Selected references

  • O. L. Mangasarian and S. Fromovitz, The Fritz John necessary optimality conditions in the presence of equality and inequality constraints, J. Math. Anal. Appl. 17 (1967), 37–47. https://doi.org/10.1016/0022-247X(67)90163-1
  • F. John, Extremum problems with inequalities as subsidiary conditions, in Studies and Essays Presented to R. Courant on his 60th Birthday, Interscience, New York, 1948, 187–204.
  • H. W. Kuhn and A. W. Tucker, Nonlinear programming, Proc. Second Berkeley Symposium on Mathematical Statistics and Probability, University of California Press, 1951, 481–492.
  • J. Gauvin, A necessary and sufficient regularity condition to have bounded multipliers in nonconvex programming, Math. Programming 12 (1977), 136–138. https://doi.org/10.1007/BF01593777
  • O. L. Mangasarian, Nonlinear Programming, McGraw-Hill, 1969; reprinted SIAM Classics in Applied Mathematics 10, 1994. https://doi.org/10.1137/1.9781611971255
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Control TheoryDynamical Systems·Captain: mikedeng1

Dynamic Instabilities and Stabilization Methods in Distributed Real-Time Scheduling of Manufacturing Systems 1: Clearing Policies Are Unstable on a Re-Entrant Two-Machine Line, Even Without Set-UpsResearch Paper

Motivation

A flexible manufacturing system is a set of machines through which parts of several types travel along fixed routes; each machine serves several buffers and must pay a set-up time whenever it switches from one buffer to another. Real-time scheduling decides, as the system evolves, which buffer each machine works on. Perkins and Kumar (IEEE Trans. Automat. Control 34, 1989) introduced simple distributed policies for this problem, of which the most natural is the clearing policy: a machine keeps working on a buffer until it is empty, and only then switches. They proved that every clear-a-fraction policy, a subclass of clearing policies, keeps every buffer bounded on acyclic systems whenever each machine has spare capacity, and left open whether clear-a-fraction policies stabilize all systems in which material flows around cycles.

Kumar and Seidman (IEEE Trans. Automat. Control 35(3), 1990, doi:10.1109/9.50339) answered no. Their Example 1 is a single part type that visits two machines in the order 1, 2, 2, 1. Every machine has spare capacity, yet under the clearing policy the buffer levels grow without bound, and they do so even when all set-up times are zero, so the instability comes from machines starving each other rather than from time lost to set-ups. Until then, instability had been suspected to require positive set-up times. Shortly afterwards Lu and Kumar exhibited instability of a static buffer-priority rule in a re-entrant network (IEEE Trans. Automat. Control 36, 1991); together these examples started the study of stability of multiclass queueing networks.

Setting

A manufacturing system has part types ppp arriving at rates dp>0d_p > 0dp​>0. Parts of type ppp follow a route of length npn_pnp​: their iii-th operation is at machine μp,i\mu_{p,i}μp,i​, and they wait for it in buffer bp,ib_{p,i}bp,i​, where each part needs processing time τp,i>0\tau_{p,i} > 0τp,i​>0. Machine mmm serves the buffers Bm={b:μb=m}B_m = \{b : \mu_b = m\}Bm​={b:μb​=m}, and switching from bbb to b′b'b′ costs set-up time δb,b′≥0\delta_{b,b'} \ge 0δb,b′​≥0.

Flows are continuous (fluid). The level of buffer bbb at time t≥0t \ge 0t≥0 is xb(t)=xb(0)+ub(t)−yb(t)≥0x_b(t) = x_b(0) + u_b(t) - y_b(t) \ge 0xb​(t)=xb​(0)+ub​(t)−yb​(t)≥0, where yb(t)y_b(t)yb​(t) is its cumulative output and ub(t)u_b(t)ub​(t) its cumulative input: dptd_p tdp​t for the first buffer of a route, and the output of the preceding buffer otherwise. Each machine works in runs: run kkk is a set-up phase of length δβk−1,βk\delta_{\beta_{k-1},\beta_k}δβk−1​,βk​​ followed by a processing phase on buffer βk\beta_kβk​, during which the buffer is drained at rate 1/τb1/\tau_b1/τb​ while it is nonempty and passed through at its inflow rate when it is empty. The system is stable if sup⁡0≤t<∞xb(t)<∞\sup_{0 \le t < \infty} x_b(t) < \inftysup0≤t<∞​xb​(t)<∞ for every buffer.

A clearing policy (Definition 1) is one in which a machine processing bbb continues until the first time that bbb is empty and some other buffer of the same machine is nonempty, and then commences a set-up for one of the nonempty buffers.

Example 1. One part type arrives at rate d=1d = 1d=1 and visits machine 1, machine 2, machine 2 and machine 1; its buffers are 1,2,3,41, 2, 3, 41,2,3,4, so B1={1,4}B_1 = \{1, 4\}B1​={1,4} and B2={2,3}B_2 = \{2, 3\}B2​={2,3}. Processing times are τ1,…,τ4>0\tau_1, \dots, \tau_4 > 0τ1​,…,τ4​>0, and δk\delta_kδk​ is the time to set up to buffer kkk. The parameters satisfy the critical condition and the capacity condition

τ2+τ4>1,τ1+τ4<1,τ2+τ3<1.(3–5)\tau_2 + \tau_4 > 1, \qquad \tau_1 + \tau_4 < 1, \qquad \tau_2 + \tau_3 < 1. \tag{3–5}τ2​+τ4​>1,τ1​+τ4​<1,τ2​+τ3​<1.(3–5)

The initial state is x(0)=(ξ,0,0,0)x(0) = (\xi, 0, 0, 0)x(0)=(ξ,0,0,0), with machine 1 set up for buffer 4 and machine 2 set up for buffer 3. Write

λ=τ41−τ2>1,α=(τ4+1)(δ1+δ2)1−τ2+δ3(τ4+1)+δ4,β=τ4(δ1+δ2)1−τ2+τ4δ3+δ4.\lambda = \frac{\tau_4}{1-\tau_2} > 1, \quad \alpha = \frac{(\tau_4+1)(\delta_1+\delta_2)}{1-\tau_2} + \delta_3(\tau_4+1) + \delta_4, \quad \beta = \frac{\tau_4(\delta_1+\delta_2)}{1-\tau_2} + \tau_4\delta_3 + \delta_4.λ=1−τ2​τ4​​>1,α=1−τ2​(τ4​+1)(δ1​+δ2​)​+δ3​(τ4​+1)+δ4​,β=1−τ2​τ4​(δ1​+δ2​)​+τ4​δ3​+δ4​.

Formalization targets

Goal: Example 1, both cases

Assume (3)–(5).

  1. If δ1,…,δ4>0\delta_1, \dots, \delta_4 > 0δ1​,…,δ4​>0, there is ξ0\xi_0ξ0​ such that for every ξ≥ξ0\xi \ge \xi_0ξ≥ξ0​ (ξ>0\xi > 0ξ>0) a clearing trajectory from (ξ,0,0,0)(\xi, 0, 0, 0)(ξ,0,0,0) exists, and every such trajectory has
sup⁡0≤t<∞x1(t)=+∞.\sup_{0 \le t < \infty} x_1(t) = +\infty.0≤t<∞sup​x1​(t)=+∞.
  1. If δ1=⋯=δ4=0\delta_1 = \dots = \delta_4 = 0δ1​=⋯=δ4​=0, the same holds for every ξ>0\xi > 0ξ>0.

Milestone: the Case 1 cycle map

For ξ\xiξ large enough, every clearing trajectory from (ξ,0,0,0)(\xi, 0, 0, 0)(ξ,0,0,0) reaches, at T1=(λ+τ2/(1−τ2))ξ+αT_1 = (\lambda + \tau_2/(1-\tau_2))\xi + \alphaT1​=(λ+τ2​/(1−τ2​))ξ+α,

x(T1)=(λξ+β,0,0,0),x(T_1) = (\lambda\xi + \beta, 0, 0, 0),x(T1​)=(λξ+β,0,0,0),

with machines 1 and 2 again set up for buffers 4 and 3.

Milestone: the Case 2 magnification

With zero set-up times and any ξ>0\xi > 0ξ>0, every clearing trajectory reaches, at t5=(τ2+τ4)ξ/(1−τ2)t_5 = (\tau_2+\tau_4)\xi/(1-\tau_2)t5​=(τ2​+τ4​)ξ/(1−τ2​),

x(t5)=(λξ,0,0,0),x(t_5) = (\lambda\xi, 0, 0, 0),x(t5​)=(λξ,0,0,0),

with machines 1 and 2 again set up for buffers 4 and 3.

Significance

The example shows that the condition ρm<1\rho_m < 1ρm​<1 on every machine, which is necessary for stability and sufficient for the existence of some stabilizing policy, does not make natural distributed policies stable once material flows around a cycle. The throughput of the line falls to 1/(τ2+τ4)<11/(\tau_2+\tau_4) < 11/(τ2​+τ4​)<1 part per unit time although each machine could handle the demand. This motivates the paper's two positive results: sufficient conditions under which clear-a-fraction policies are stable (Theorem 1), and a supervisory mechanism that stabilizes any policy (Theorem 2), which are the subjects of the other missions of this series. The example is also an early instance of the phenomenon later studied as instability of multiclass fluid networks under work-conserving policies.

The paper's argument is a stage-by-stage computation of piecewise linear trajectories. No machine-checked version of it exists. A formal proof has to make precise what the paper leaves to the reader: that the clearing rule determines the trajectory, that the stage formulas are what that trajectory does, and that the cycle can be restarted. The formal model of runs, set-ups and the clearing rule built here is the same as in the other missions of the series.

Difficulty

The arithmetic of each cycle is routine once the trajectory is known. The difficulty is in the universal quantifier: the claim covers every clearing trajectory, and the clearing rule is defined implicitly, through "the first time thereafter" at which a buffer is empty and another one is nonempty. At several switching instants the buffer a machine switches to is empty and only starts to fill at that instant, and with zero set-up times a machine may begin a run at an instant where the switching condition already holds. Showing that each switch happens exactly when the paper says, and that the fluid levels then follow the printed formulas (including the reduced rate of a machine working on an empty buffer), is a uniqueness argument for a hybrid system, not a simulation. The existence half asks for the converse: an explicit trajectory, defined for all time, with infinitely many runs whose start times tend to infinity.

Formalization scope

  • Time is real (t≥0t \ge 0t≥0); flows are fluid; there are no transport delays or assembly.
  • The system is a general structure (part types Fin P, machines Fin M, buffers ⟨p, i⟩ with i : Fin (n p), paper index iii = Lean index i+1i+1i+1), instantiated as Example 1 with d=1d = 1d=1, route (1,2,2,1)(1,2,2,1)(1,2,2,1), and δb,b′=δb′\delta_{b,b'} = \delta_{b'}δb,b′​=δb′​ for b≠b′b \ne b'b=b′; staying on a buffer costs nothing.
  • A trajectory is a schedule of runs per machine (possibly finitely many, the last lasting forever), with only finitely many run starts in any bounded interval. Processing obeys a rate cap (yby_byb​ grows at most at rate 1/τb1/\tau_b1/τb​, and only while machine μb\mu_bμb​ is in a processing phase of bbb) and runs at full rate while the buffer is nonempty.
  • In Definition 1, a target buffer counts as "nonempty" when it is demanding: positive level, or inflow starting at that instant. The no-early-exit condition is imposed on the open processing interval. Under the literal reading (positive level) or a closed interval, the paper's own trajectories are not clearing, and the goal would hold vacuously; the existence clause in the goal rules out that trivialization.
  • "Set up for buffer bbb at time TTT" means the run in force on (sk,sk+1](s_k, s_{k+1}](sk​,sk+1​].
  • Unboundedness is stated for buffer 1: for every CCC there is t≥0t \ge 0t≥0 with x1(t)>Cx_1(t) > Cx1​(t)>C; "ξ\xiξ large enough" is ∃ξ0,∀ξ≥ξ0\exists \xi_0, \forall \xi \ge \xi_0∃ξ0​,∀ξ≥ξ0​.

Useful contributions include lemmas about fluid trajectories that do not depend on the example (continuity of levels, the pass-through rate on an empty buffer, restarting a trajectory at a run boundary), which also serve the other missions of the series.

Selected references

  • P. R. Kumar and T. I. Seidman, Dynamic instabilities and stabilization methods in distributed real-time scheduling of manufacturing systems, IEEE Trans. Automat. Control 35(3), 289–298, 1990. https://doi.org/10.1109/9.50339
  • J. R. Perkins and P. R. Kumar, Stable, distributed, real-time scheduling of flexible manufacturing/assembly/disassembly systems, IEEE Trans. Automat. Control 34, 139–148, 1989 (reference [18] of the paper).
  • S. H. Lu and P. R. Kumar, Distributed scheduling based on due dates and buffer priorities, IEEE Trans. Automat. Control 36, 1991.
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ProbabilityStochastic Systems·Captain: mikedeng1

Stochastic Inequalities on Partially Ordered Spaces 1: Stochastically Ordered Initial Laws and Kernels Give Coupled Random Sequences with Xₙ ≤ Yₙ for All n Almost SurelyResearch Paper

Motivation

Comparison theorems for stochastic processes answer a practical question: if one system starts "lower" and moves "upward" less aggressively than another, does it stay below the other one for all time? Queueing, reliability and inventory models use such statements to order performance measures of two systems without computing either distribution. Results of this kind for real-valued Markov chains go back to Kalmykov (1962) and Daley (1968); O'Brien (1975) proved a comparison theorem for real sequences with general (non-Markov) dependence on the past.

Kamae, Krengel and O'Brien (1977) placed these results on a common foundation: an arbitrary partially ordered Polish space, where the state can be a vector, a path, a configuration or a measure. Their tool is a characterization of the stochastic order through monotone couplings, which goes back to Strassen (1965).

Timeline.

  • 1962, Kalmykov: comparison of real Markov chains with stochastically monotone kernels.
  • 1965, Strassen: existence of probability measures with given marginals; a coupling on a closed set K⊆E×EK \subseteq E \times EK⊆E×E exists iff the marginals satisfy the matching inequalities.
  • 1968, Daley: stochastically monotone Markov chains on R\mathbb RR.
  • 1975, O'Brien: comparison theorem for real random sequences with history-dependent kernels.
  • 1977, Kamae–Krengel–O'Brien: the order on a partially ordered Polish space, Theorem 1 (six characterizations), and the comparison theorem for sequences in products of such spaces (Theorem 2).

Setting

Let EEE be a complete separable metric space with a closed partial order ≤\le≤ (the set {(x,y):x≤y}\{(x,y) : x \le y\}{(x,y):x≤y} is closed in E×EE \times EE×E) and its Borel σ\sigmaσ-algebra. Write M(E)\mathcal M(E)M(E) for the probability measures on EEE. A set A⊆EA \subseteq EA⊆E is increasing if x∈Ax \in Ax∈A and x≤yx \le yx≤y imply y∈Ay \in Ay∈A; a function fff is increasing if x≤yx \le yx≤y implies f(x)≤f(y)f(x) \le f(y)f(x)≤f(y). For P1,P2∈M(E)P_1, P_2 \in \mathcal M(E)P1​,P2​∈M(E), P1P_1P1​ is stochastically smaller than P2P_2P2​, written P1≺P2P_1 \prec P_2P1​≺P2​, if

∫f dP1≤∫f dP2for every bounded measurable increasing f:E→R.\int f\,dP_1 \le \int f\,dP_2 \quad\text{for every bounded measurable increasing } f : E \to \mathbb R .∫fdP1​≤∫fdP2​for every bounded measurable increasing f:E→R.

A stochastic kernel kkk from E1E_1E1​ to E2E_2E2​ assigns to every x∈E1x \in E_1x∈E1​ a probability measure k(x,⋅)k(x,\cdot)k(x,⋅) on E2E_2E2​, measurably in xxx. For P1∈M(E1)P_1 \in \mathcal M(E_1)P1​∈M(E1​), P1∗kP_1 * kP1​∗k is the measure on E1×E2E_1 \times E_2E1​×E2​ with (P1∗k)(A1×A2)=∫A1k(x,A2) P1(dx)(P_1 * k)(A_1 \times A_2) = \int_{A_1} k(x, A_2)\,P_1(dx)(P1​∗k)(A1​×A2​)=∫A1​​k(x,A2​)P1​(dx), and P1kP_1^{k}P1k​ is its second marginal. A kernel kkk on E×EE \times EE×E is upward if k(x,⋅)k(x,\cdot)k(x,⋅) is concentrated on {y:y≥x}\{y : y \ge x\}{y:y≥x} for every xxx.

For partially ordered Polish spaces E1,E2,…E_1, E_2, \dotsE1​,E2​,…, the products En=E1×⋯×EnE^{n} = E_1 \times \cdots \times E_nEn=E1​×⋯×En​ and E∞=∏iEiE^\infty = \prod_i E_iE∞=∏i​Ei​ carry the product topology and the coordinatewise order, and are again partially ordered Polish spaces. Given P1∈M(E1)P_1 \in \mathcal M(E_1)P1​∈M(E1​) and kernels pnp_npn​ from En−1E^{n-1}En−1 to EnE_nEn​ (n≥2n \ge 2n≥2), the measure P1∗p2∗⋯∗pnP_1 * p_2 * \cdots * p_nP1​∗p2​∗⋯∗pn​ on EnE^{n}En is the law of (X1,…,Xn)(X_1, \dots, X_n)(X1​,…,Xn​) when X1∼P1X_1 \sim P_1X1​∼P1​ and XnX_nXn​ is drawn from pn(X1,…,Xn−1,⋅)p_n(X_1, \dots, X_{n-1}, \cdot)pn​(X1​,…,Xn−1​,⋅); its projective limit is the law of the whole sequence.

Formalization targets

Goal: Theorem 2 (the discrete-time comparison theorem)

Let P1,Q1∈M(E1)P_1, Q_1 \in \mathcal M(E_1)P1​,Q1​∈M(E1​) and let pn,qnp_n, q_npn​,qn​ be stochastic kernels from En−1E^{n-1}En−1 to EnE_nEn​, n≥2n \ge 2n≥2. If P1≺Q1P_1 \prec Q_1P1​≺Q1​ and

pn(xn−1,⋅)≺qn(yn−1,⋅)whenever xn−1≤yn−1,p_n(x^{n-1},\cdot) \prec q_n(y^{n-1},\cdot) \quad\text{whenever } x^{n-1} \le y^{n-1},pn​(xn−1,⋅)≺qn​(yn−1,⋅)whenever xn−1≤yn−1,

then there are random sequences (Xn)(X_n)(Xn​), (Yn)(Y_n)(Yn​) on one probability space, with initial laws P1P_1P1​, Q1Q_1Q1​ and conditional laws pnp_npn​, qnq_nqn​, such that

P(Xi≤Yi, i=1,2,… )=1.P(X_i \le Y_i,\ i = 1, 2, \dots) = 1 .P(Xi​≤Yi​, i=1,2,…)=1.

Milestones

  1. Theorem 1: for P1,P2∈M(E)P_1, P_2 \in \mathcal M(E)P1​,P2​∈M(E), P1≺P2P_1 \prec P_2P1​≺P2​ is equivalent to each of: a coupling supported on {x≤y}\{x \le y\}{x≤y}; a representation f(Z)∼P1f(Z) \sim P_1f(Z)∼P1​, g(Z)∼P2g(Z) \sim P_2g(Z)∼P2​ with f≤gf \le gf≤g and ZZZ real; random variables X1≤X2X_1 \le X_2X1​≤X2​ a.s. with laws P1,P2P_1, P_2P1​,P2​; P2=P1kP_2 = P_1^{k}P2​=P1k​ for an upward kernel kkk; and P1(B)≤P2(B)P_1(B) \le P_2(B)P1​(B)≤P2​(B) for every closed increasing BBB.
  2. Proposition 1: under the hypotheses of Theorem 2, P1∗p2∗⋯∗pn≺Q1∗q2∗⋯∗qnP_1 * p_2 * \cdots * p_n \prec Q_1 * q_2 * \cdots * q_nP1​∗p2​∗⋯∗pn​≺Q1​∗q2​∗⋯∗qn​ for every nnn.
  3. Proposition 2: on E∞E^\inftyE∞, if all finite-dimensional marginals satisfy P(i)≺Q(i)P^{(i)} \prec Q^{(i)}P(i)≺Q(i), then P≺QP \prec QP≺Q.
  4. Proposition 3: ≺\prec≺ is preserved under weak convergence.
  5. Proposition 4: P1≺P2≺⋯P_1 \prec P_2 \prec \cdotsP1​≺P2​≺⋯ iff there are random elements X1≤X2≤⋯X_1 \le X_2 \le \cdotsX1​≤X2​≤⋯ a.s. with Xi∼PiX_i \sim P_iXi​∼Pi​, iff there are ZZZ real and f1≤f2≤⋯f_1 \le f_2 \le \cdotsf1​≤f2​≤⋯ with fi(Z)∼Pif_i(Z) \sim P_ifi​(Z)∼Pi​.
  6. Corollary 1 (i)–(iii), consequences of Theorem 2 when E1=E2=⋯E_1 = E_2 = \cdotsE1​=E2​=⋯: for an increasing set AAA, P(Xn∈A)≤P(Yn∈A)P(X_n \in A) \le P(Y_n \in A)P(Xn​∈A)≤P(Yn​∈A); first entrance times into AAA satisfy P(Nx<n)≤P(Ny<n)P(N_x < n) \le P(N_y < n)P(Nx​<n)≤P(Ny​<n); and Ef(Xn)≤Ef(Yn)E f(X_n) \le E f(Y_n)Ef(Xn​)≤Ef(Yn​) for nondecreasing fff whenever the expectations exist.
  7. Theorem 3: in a partially ordered Polish space with a compatible vector structure, kernels ordered through the increments they produce yield processes (Sn)(S_n)(Sn​), (Tn)(T_n)(Tn​) with S1≤T1S_1 \le T_1S1​≤T1​ and Sn+1−Sn≤Tn+1−TnS_{n+1} - S_n \le T_{n+1} - T_nSn+1​−Sn​≤Tn+1​−Tn​ for all nnn, almost surely.

Significance

The result. Theorem 2 turns an inequality between one-step transition laws into a pathwise inequality between whole trajectories. Every functional that is increasing in the path, such as hitting times of increasing sets, maxima, occupation counts or cumulative costs, is then ordered between the two processes, as Corollary 1 illustrates. Because the state space is any partially ordered Polish space, the theorem covers vector-valued queue lengths, networks, and processes whose state is itself a sequence, not just real chains. Theorem 1 is the basic tool for working with the order on such spaces: it converts between integrals, couplings, kernels and closed increasing sets.

Formalizing it. All results are proved in the paper (1977); none is machine-checked. The platform has the special case of Theorem 1 (i) ⇔\Leftrightarrow⇔ (iv) for E=RnE = \mathbb R^nE=Rn as an open statement (PalmQueueing.Ordering.strassen_st); the general partially ordered Polish case, and the sequence comparison, are new. A formal development yields a reusable library for the stochastic order on general ordered spaces and its interaction with Mathlib's Ionescu-Tulcea construction of process laws (Kernel.trajMeasure).

Difficulty

The obvious argument for Theorem 2 couples the processes step by step: couple X1≤Y1X_1 \le Y_1X1​≤Y1​, then, given the two histories, couple X2≤Y2X_2 \le Y_2X2​≤Y2​, and so on. Each step needs a monotone coupling of pn(xn−1,⋅)p_n(x^{n-1},\cdot)pn​(xn−1,⋅) and qn(yn−1,⋅)q_n(y^{n-1},\cdot)qn​(yn−1,⋅) chosen measurably in the pair of histories; the existence of one coupling for each fixed pair (Theorem 1) does not by itself give a kernel.

Theorem 1's central implication, from the integral inequality to a coupling supported on the closed set {x≤y}\{x \le y\}{x≤y}, is Strassen's theorem. On R\mathbb RR it follows from quantile functions; on a general ordered space there is no such formula. Passing from finite horizons to the infinite sequence is a further step, since an ordering of every finite-dimensional marginal must be turned into an ordering of the infinite-dimensional laws.

Formalization scope

  • A partially ordered Polish space is [TopologicalSpace E] [PolishSpace E] [MeasurableSpace E] [BorelSpace E] [PartialOrder E] [OrderClosedTopology E]. This is the paper's standing assumption (Sec. 1, p. 899) and is carried by every theorem, also where a statement says only "Polish space". All sets and functions the paper quantifies over are measurable, also by the standing assumption.
  • StochLE P₁ P₂ is defined through bounded measurable monotone functions, exactly as on p. 899. It is not defined through increasing sets, closed increasing sets or couplings, so none of the milestones holds by definition.
  • Sequences are 0-based: the paper's EiE_iEi​, XiX_iXi​ are Lean's E (i - 1), X (i - 1), and the paper's kernel pnp_npn​ (n≥2n \ge 2n≥2) is p (n - 2) : Kernel (Π i : Iic (n - 2), E i) (E (n - 1)), the signature of Mathlib's Ionescu-Tulcea API.
  • "Random sequences with initial law P1P_1P1​ and conditional laws pnp_npn​" is encoded through their joint law, which these data determine: the law of (Xn)(X_n)(Xn​) is Kernel.trajMeasure P₁ p. Corollary 1 is stated about these laws directly. The goal therefore cannot be met by coupling only one-dimensional marginals or finite prefixes; it asserts the joint laws of the full sequences and one almost-sure event for all indices.
  • Hypothesis (4) is required pointwise for all ordered pairs of histories; no stochastic monotonicity of the kernels is assumed.
  • Corollary 1 (iii) takes expectations in the extended sense, Ef=Ef+−Ef−E f = E f^+ - E f^-Ef=Ef+−Ef− in EReal, and "the expectation exists" means the two parts are not both infinite; expectations equal to ±∞\pm\infty±∞ are covered.
  • Theorem 1 (iii) and Proposition 4 (iii) leave the law of the real variable ZZZ free, as the paper does.
  • Theorem 3 (p. 904) prints the conditional increment law of TTT as qn+1(T1,…,Tn;A+Sn)q_{n+1}(T_1, \dots, T_n; A + S_n)qn+1​(T1​,…,Tn​;A+Sn​); this is a misprint for A+TnA + T_nA+Tn​, and the Lean states the corrected form. Its compatible vector structure is [AddCommGroup E] [Module ℝ E] [ContinuousAdd E] [ContinuousSMul ℝ E] plus the hypothesis that translates of increasing sets are increasing.
  • Not included: Sections 4–5 (continuous time, which needs the Skorohod space).

Welcome contributions: proofs of the milestones, in particular Strassen's coupling theorem on partially ordered Polish spaces, a measurable-selection lemma for monotone couplings, and general lemmas relating StochLE to increasing sets.

The source is the published version (The Annals of Probability), and printed page = PDF page + 898 throughout.

Selected references

  • T. Kamae, U. Krengel, G. L. O'Brien, Stochastic Inequalities on Partially Ordered Spaces, The Annals of Probability 5(6), 1977, 899–912. https://doi.org/10.1214/aop/1176995659
  • V. Strassen, The existence of probability measures with given marginals, The Annals of Mathematical Statistics 36(2), 1965, 423–439. https://doi.org/10.1214/aoms/1177700153
  • G. L. O'Brien, The comparison method for stochastic processes, The Annals of Probability 3, 1975, 80–88. https://projecteuclid.org/journals/annals-of-probability/volume-3/issue-1
  • D. J. Daley, Stochastically monotone Markov chains, Zeitschrift für Wahrscheinlichkeitstheorie und verwandte Gebiete 10, 1968, 307–317 (as cited in the paper).
  • G. I. Kalmykov, On the partial ordering of one-dimensional Markov processes, Theory of Probability and its Applications 7, 1962, 456–459 (as cited in the paper).
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CombinatoricsProbabilityTheoretical Computer Science·Captain: mikedeng1

Packet Routing and Job-Shop Scheduling in O(Congestion + Dilation) Steps: Edge-Simple Paths with Congestion c and Dilation d Admit an O(c + d)-Step Schedule with Constant-Size QueuesResearch Paper

Motivation

In a store-and-forward network, messages are cut into packets that travel from node to node along wires, one wire per time step, and wait in buffers between moves. Routing such traffic splits into two problems: choosing a path for each packet, and scheduling the packets along their paths, deciding at every step which packets move and which wait. Leighton, Maggs and Rao showed that the second problem always has an essentially optimal solution: once paths are fixed, two simple parameters of the paths determine the routing time up to a constant factor, on every network.

The result separates path selection from timing in the design of routing algorithms for parallel machines, and it reaches beyond networks: a job-shop problem in which every operation takes one unit of time and no job visits a machine twice is the same problem with jobs as packets and machines as edges.

Timeline.

  • 1988: Leighton, Maggs and Rao, extended abstract at FOCS, Universal packet routing algorithms.
  • 1994: the full paper in Combinatorica 14, with the existence theorem of an O(c+d)O(c+d)O(c+d) schedule with constant queues (Theorem 3.4) and a randomized on-line algorithm.
  • 1999: Leighton, Maggs and Richa, Fast algorithms for finding O(congestion + dilation) packet routing schedules, make the construction algorithmic, using the algorithmic Local Lemma of Beck.

Setting

A network is a directed multigraph: a type VVV of nodes, a type EEE of edges, and maps src,tgt:E→V\mathrm{src},\mathrm{tgt}:E\to Vsrc,tgt:E→V. A path is a list of edges e0,…,eℓ−1e_0,\dots,e_{\ell-1}e0​,…,eℓ−1​ with tgt(ek)=src(ek+1)\mathrm{tgt}(e_k)=\mathrm{src}(e_{k+1})tgt(ek​)=src(ek+1​); it is edge-simple if no edge occurs twice. A finite set PPP of packets is given, each with its path.

The dilation ddd is the largest number of edges on a path; the congestion ccc is the largest number of paths through one edge. Since a packet crosses at most one edge per step and an edge carries at most one packet per step, every schedule needs at least max⁡(c,d)\max(c,d)max(c,d) steps.

A schedule assigns to every packet ppp and every index kkk of its path the step τ(p,k)≥1\tau(p,k)\ge1τ(p,k)≥1 at which ppp crosses its kkk-th edge, strictly increasing in kkk. Its length is the last crossing step. A packet waits in its initial queue before its first crossing, in the edge queue at the head of the edge it last crossed between two crossings, and in its final queue after its last crossing. Only edge queues count for queue size: the other two are fixed by the instance. A schedule is valid if at most one packet crosses each edge at each step.

For the proof, the paper measures schedules that are not yet valid through frames: a TTT-frame is a run of TTT consecutive steps, and the relative congestion of a frame is the largest number of packets crossing one edge in it, divided by TTT.

Formalization targets

Goal: Theorem 3.4 (p. 11)

There are absolute constants KKK and QQQ such that every finite set of packets with edge-simple paths of congestion at most ccc and dilation at most ddd, on any network, has a valid schedule with

length≤K (c+d),every edge queue≤Q at every step.\text{length}\le K\,(c+d),\qquad \text{every edge queue}\le Q \text{ at every step}.length≤K(c+d),every edge queue≤Q at every step.

The constants are not fixed: the paper proves existence, and any constants are a valid answer.

Milestones, in the order the proof uses them

  1. Lemma 3.1 (p. 8), the Lovász Local Lemma: events of probability at most ppp with dependence at most b≥1b\ge1b≥1 and 4pb<14pb<14pb<1 all fail together with positive probability.
  2. Lemma 3.2 (p. 8): with congestion and dilation at most ddd, a schedule of length O(d)O(d)O(d) with no waiting in edge queues and at most TTT packets per edge in every frame of size T≥log⁡2dT\ge\log_2 dT≥log2​d.
  3. The recurrences (pp. 12–13): I(1)=log⁡dI^{(1)}=\log dI(1)=logd, I(i+1)=log⁡5I(i)I^{(i+1)}=\log^5 I^{(i)}I(i+1)=log5I(i), r(1)=1r^{(1)}=1r(1)=1, r(i+1)=r(i)(1+κ/log⁡I(i))r^{(i+1)}=r^{(i)}(1+\kappa/\sqrt{\log I^{(i)}})r(i+1)=r(i)(1+κ/logI(i)​) stop at some j=O(log⁡∗d)j=O(\log^* d)j=O(log∗d) with r(j)=O(1)r^{(j)}=O(1)r(j)=O(1).
  4. Lemma 3.5 (p. 14): frame bounds for all sizes from TTT to 2T−12T-12T−1 imply them for all sizes ≥T\ge T≥T.
  5. The final simulation (pp. 12–13): a schedule with relative congestion O(1)O(1)O(1) in frames of constant size, in which every packet waits at most once every k1≥2k_1\ge2k1​≥2 steps, becomes a valid schedule, a constant factor longer, with constant edge queues.

Significance

The theorem shows that congestion and dilation, two quantities read off the paths alone, determine the optimal schedule length up to a constant factor on every network, with buffers of constant size. Path-selection algorithms that minimize c+dc+dc+d therefore yield near-optimal routing, and the same bound holds for unit-time job shops without repeated machines. It is also an often-cited application of the Local Lemma beyond a single round of random choices: the proof applies it O(log⁡∗d)O(\log^* d)O(log∗d) times in succession.

The result has been proved since 1994, and the algorithmic version since 1999. As far as is known, no part of it is machine-checked. The mission produces a formal model of store-and-forward schedules with edge queues and frames, a formal statement of the theorem that excludes the degenerate readings, and formal versions of the steps of its proof.

Difficulty

The naive approach gives each packet a random initial delay and then lets it move without waiting. That gives O(log⁡(Nd))O(\log(Nd))O(log(Nd)) packets per edge per step, and O(c+dlog⁡(Nd))O(c+d\log(Nd))O(c+dlog(Nd)) steps after slowing down. Lemma 3.2 does better only in frames of size log⁡d\log dlogd, not in single steps. Recursing on frames, as in Theorem 3.3, loses a constant factor per level and gives (c+d)2O(log⁡∗(c+d))(c+d)2^{O(\log^* (c+d))}(c+d)2O(log∗(c+d)).

Removing that factor is the central difficulty. Each refinement must keep the relative congestion nearly unchanged, r(i+1)=r(i)(1+O(1)/log⁡I(i))r^{(i+1)}=r^{(i)}(1+O(1)/\sqrt{\log I^{(i)}})r(i+1)=r(i)(1+O(1)/logI(i)​), which requires second-order terms in the tail estimates, delays spread over the block rather than inserted at its start, and careful handling of block boundaries. The constant queue bound needs an invariant: every packet waits at most once every I(i)I^{(i)}I(i) steps. The Local Lemma gives existence only; the construction is non-constructive.

Formalization scope

Conventions committed to in Lean:

  • The network is arbitrary: V E : Type with src tgt : E → V, no finiteness or degree bound. Packets form a Fintype P; paths are List E with matching endpoints (List.IsChain) and List.Nodup.
  • Congestion and dilation are bounds (CongestionLE path c, DilationLE path d), equivalent to exact values because every bound is monotone.
  • A schedule is a Timetable: time p k : ℕ, the step at which packet p crosses its k-th edge, at least 1 and strictly increasing in k. Length at most L means every crossing step is ≤ L. Valid means two different crossings of one edge happen at different steps.
  • The edge-queue size of edge g at the end of step t counts crossings of g at a step ≤ t whose packet crosses its next edge at a step > t. Initial and final queues are not counted.
  • Frames: frameCount τ g t T counts the packets that cross g at a step in [t, t+T). Relative congestion at most r in frames of size ≥ T₀ is frameCount ≤ r·T for all T ≥ 1 with T₀ ≤ T.
  • log is Real.logb 2. Every O(1) and "sufficiently large" is a constant quantified before the instance, and in the goal ∃ K Q comes before the network.
  • Lemma 3.1 is stated on an arbitrary probability space with measurable events. Dependence uses independence from the generated σ-algebra, and the statement adds 1 ≤ b, without which the page's statement is false.

These choices rule out the trivial formalizations: constants chosen after the instance would make K=cdK=cdK=cd suffice; a schedule without edge exclusivity would make the greedy schedule of length ddd a solution; a missing queue bound drops half of the theorem; and a statement without its length bound is solved by sending one packet at a time.

Not stated: Lemmas 3.6–3.10 and the summary of the refinement step (p. 20). They concern the block decomposition and delay-insertion rules of pp. 13–14 and 17–18, which the paper defines only in prose. Contributions are welcome on the Local Lemma (finite or general), the probabilistic estimates of Lemma 3.2, Lemma 3.5, the elementary simulation step, and formal definitions of the block operations from which Lemmas 3.6–3.10 can be stated. The schedule model is reusable for other routing results, such as the on-line algorithm of §2 or the O(c+d)O(c+d)O(c+d) results for leveled networks.

Selected references

  • F. T. Leighton, B. M. Maggs, S. B. Rao, Packet routing and job-shop scheduling in O(congestion + dilation) steps, Combinatorica 14 (1994) 167–186. https://doi.org/10.1007/BF01215349 (this mission cites the authors' manuscript).
  • F. T. Leighton, B. M. Maggs, S. B. Rao, Universal packet routing algorithms, Proc. 29th IEEE FOCS (1988) 256–269.
  • F. T. Leighton, B. M. Maggs, A. W. Richa, Fast algorithms for finding O(congestion + dilation) packet routing schedules, Combinatorica 19 (1999) 375–401. https://doi.org/10.1007/s004930050061
  • P. Erdős, L. Lovász, Problems and results on 3-chromatic hypergraphs and some related questions, in Infinite and Finite Sets, Colloq. Math. Soc. János Bolyai 10 (1975) 609–627.
  • J. Spencer, Ten Lectures on the Probabilistic Method, SIAM (1987), pp. 57–58. https://doi.org/10.1137/1.9780898719918
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Combinatorics·Captain: mikedeng1

Sequencing Jobs to Minimize Total Weighted Completion Time Subject to Precedence Constraints 1: The Series-Parallel Algorithm Returns an Optimal SequenceResearch Paper

Motivation

Sequencing jobs on one machine to minimize the total weighted completion time ∑jwjCj\sum_j w_jC_j∑j​wj​Cj​ is one of the basic problems of scheduling theory. Without precedence constraints it is solved by Smith's ratio rule (1956): sequence the jobs in nonincreasing order of wj/pjw_j/p_jwj​/pj​. With arbitrary precedence constraints the problem is NP-hard, as Lawler shows in the same report. Between these two extremes lies the class of series parallel precedence constraints, which covers chains, rooted trees and their nested combinations, and which is exactly the class that can be assembled from single jobs by "all of this before all of that" and "this independently of that".

Lawler's report (IRIA-LABORIA No. 205, 1976; Annals of Discrete Mathematics 2, 1978) gives an O(nlog⁡n)O(n\log n)O(nlogn) algorithm for this class, once a decomposition of the constraints is known. It is the standard reference for the polynomial case of 1 ∣ prec ∣ ∑wjCj1\,|\,\mathrm{prec}\,|\,\sum w_jC_j1∣prec∣∑wj​Cj​.

Timeline.

  • 1956, Smith: the ratio rule, no precedence constraints.
  • 1964, Conway, Maxwell and Miller: parallel chains with unit weights.
  • 1971–1973, Baker, Horn, Adolphson and Hu: rooted trees, in O(nlog⁡n)O(n\log n)O(nlogn).
  • 1975, Sidney: job modules and ρ\rhoρ-maximal initial sets, for arbitrary precedence constraints, without an efficient method to find them.
  • 1976/1978, Lawler: series parallel constraints in O(nlog⁡n)O(n\log n)O(nlogn); NP-completeness for arbitrary constraints.

Setting

There are nnn jobs, forming a finite set NNN, to be processed by a single machine. Precedence constraints are given by an acyclic digraph G=(N,A)G=(N,A)G=(N,A): job iii must precede job jjj if there is a directed path from iii to jjj. Job jjj has a processing time pj>0p_j>0pj​>0 and a real weight wjw_jwj​ (of either sign). A feasible sequence lists every job once and never puts a job before one that must precede it. The machine starts at time 000 without idle time, so the completion time CjC_jCj​ of job jjj is the sum of the processing times of jjj and the jobs before it. The cost of a sequence is ∑j∈NwjCj\sum_{j\in N} w_jC_j∑j∈N​wj​Cj​, and an optimal sequence is a feasible sequence of minimum cost.

A transitive series parallel digraph is built from single nodes by series composition (N1∪N2, A1∪A2∪N1×N2)(N_1\cup N_2,\ A_1\cup A_2\cup N_1\times N_2)(N1​∪N2​, A1​∪A2​∪N1​×N2​) and parallel composition (N1∪N2, A1∪A2)(N_1\cup N_2,\ A_1\cup A_2)(N1​∪N2​, A1​∪A2​) of digraphs on disjoint node sets. GGG is series parallel if its transitive closure is transitive series parallel. A decomposition tree records such a construction: its leaves are the jobs, and its internal nodes are marked SSS (left son before right son) or PPP.

Sidney's theory works with three notions. A nonempty M⊆NM\subseteq NM⊆N is a module if every job outside MMM must precede all of MMM, must follow all of MMM, or is unconstrained with respect to all of MMM. A subset I⊆MI\subseteq MI⊆M is an initial set of MMM if it contains, with each of its jobs, every job of MMM that must precede it. With

ρ(I)=∑j∈Iwj∑j∈Ipj,\rho(I)=\frac{\sum_{j\in I}w_j}{\sum_{j\in I}p_j},ρ(I)=∑j∈I​pj​∑j∈I​wj​​,

an initial set I∗I^*I∗ is ρ\rhoρ-maximal if ρ(I∗)≥ρ(I)\rho(I^*)\ge\rho(I)ρ(I∗)≥ρ(I) for every initial set III of MMM.

A composite job is a sequence of jobs treated as one job whose weight and processing time are the sums over its jobs. Lawler's algorithm works bottom-up on the decomposition tree and represents an optimal sequence for each node's module as a set of composite jobs: a leaf gives its job, a PPP-node the union of its sons' sets, and an SSS-node the output of the series procedure (Steps 1–3, p. 11), which absorbs elements of M1M_1M1​ and M2M_2M2​ into a composite job k=(i,k)k=(i,k)k=(i,k) or k=(k,j)k=(k,j)k=(k,j) until the ratios separate.

Formalization targets

Goal: the algorithm is correct

For a decomposition tree TTT of GGG and any run of the algorithm on TTT producing the set FFF of composite jobs, every arrangement LLL of FFF in nonincreasing ratio order, with each composite job expanded into its sequence, is an optimal sequence:

flatten⁡(L) is an optimal sequence for N.\operatorname{flatten}(L)\ \text{is an optimal sequence for } N .flatten(L) is an optimal sequence for N.

Ties, both in the algorithm's choice of minimal and maximal elements and in the final sort, may be broken arbitrarily.

Milestones

  1. Subtrees are modules (p. 7): the leaves of any subtree of a decomposition tree form a module.
  2. Theorem 1 (p. 7): an optimal sequence for a module extends to an optimal sequence for NNN in which the module's jobs keep their order.
  3. Existence (Definition 3, p. 8): every module has a ρ\rhoρ-maximal initial set.
  4. Theorem 2 (p. 8): for a ρ\rhoρ-maximal initial set III of a module MMM, some optimal sequence for NNN runs III consecutively, before all other jobs of MMM.
  5. Ratio order (p. 10): composite jobs in nonincreasing ratio order cost no more than in any other order.
  6. Parallel composition (p. 10): the union of two representing sets represents the parallel composition.
  7. Series composition, separated ratios (p. 10): if every ratio in M1M_1M1​ exceeds every ratio in M2M_2M2​, the union represents the series composition.
  8. Series composition, Steps 1–3 (pp. 10–11): the output of the series procedure represents the series composition.

A companion theorem states that the algorithm always has a run.

Significance

The result places 1 ∣ prec ∣ ∑wjCj1\,|\,\mathrm{prec}\,|\,\sum w_jC_j1∣prec∣∑wj​Cj​ with series parallel constraints in polynomial time. Sidney's module and ρ\rhoρ-maximal initial set machinery, which it uses, is the decomposition later used in approximation algorithms for general precedence constraints; Theorems 1 and 2 hold for arbitrary acyclic constraints.

The algorithm and Sidney's theorems are proved in the literature; none of them has a machine-checked proof. The report proves Theorems 1 and 2 by citation to Sidney's 25 lemmas and argues the algorithm's correctness informally in a few paragraphs. A formal development supplies the missing invariants (what exactly is preserved at each node of the tree) and checks that arbitrary tie-breaking is harmless.

Difficulty

The algorithm's correctness rests on an invariant the page states only in passing: that composite jobs behave as single jobs. Read literally, "nonincreasing ratio order is optimal for M1M_1M1​ and for M2M_2M2​" does not imply the same for their parallel composition. A composite job of two unconstrained jobs can be optimal on its own module yet blocks a better interleaving once another module is added. The invariant that makes the induction go through is that every composite job is a ρ\rhoρ-maximal initial set of itself, and the series procedure must be shown to preserve it while merging. The second difficulty is Theorems 1 and 2 themselves: exchange arguments over sequences of a whole module, with real weights of either sign, under arbitrary acyclic constraints.

Formalization scope

Jobs have a type ι\iotaι with decidable equality; NNN is a Finset ι; the arcs are a relation G : ι → ι → Prop with endpoints in NNN, acyclic in the sense that no job reaches itself by a path, and "must precede" is Relation.TransGen G. Processing times are positive on NNN, weights are arbitrary reals. The objective and optimality are the published SingleMachinePrec.Biclique.WeightedCompletion; feasibility is checked against the arcs of GGG, which for a sequence listing each job once is the same as against its paths. "Optimal for MMM" uses the job set MMM.

The decomposition tree is an inductive SPTree with leaves, SSS-nodes and PPP-nodes. It is tied to GGG by three hypotheses: its leaves are distinct, they are exactly NNN, and on NNN the transitive closure of GGG equals the arc relation of the transitive series parallel digraph the tree builds. Composite jobs are lists of jobs, sets of composite jobs are lists of lists, and their ratio is ∑w/∑p\sum w/\sum p∑w/∑p over the list. The algorithm is an inductive relation (Run, SeriesMerge, SeriesLoop), so every tie-breaking is a run. The dummies of ratio ±∞\pm\infty±∞ in Step 1 become vacuous conditions on empty sets.

Committed conventions and disclosed deviations:

  • Lean evaluates 0/0=00/0=00/0=0, so ρ(∅)=0\rho(\emptyset)=0ρ(∅)=0; ρ\rhoρ-maximality compares nonempty initial sets only and requires the maximal set to be nonempty, and composite jobs are nonempty.
  • Modules use the disjunction of (4.1)–(4.3); for acyclic GGG and nonempty MMM this is the page's "exactly one".
  • "Represents an optimal sequence" includes the composite invariant (every nonempty proper prefix has ratio at most the whole). Without it the parallel and series milestones are false.
  • In the parallel and series milestones, M1M_1M1​, M2M_2M2​ and M1∪M2M_1\cup M_2M1​∪M2​ are assumed to be modules, as on the page (the sons of a node of the decomposition tree and the node itself). Without this, a path through a job outside M1∪M2M_1\cup M_2M1​∪M2​ could constrain two of its jobs while feasibility on M1∪M2M_1\cup M_2M1​∪M2​ does not see it, and the series milestone would be false.
  • The introduction's "nondecreasing order of the ratios" (p. 1) is a slip for the nonincreasing order used in §5; the formalization uses nonincreasing.

The goal assumes nothing about modules, Theorem 2 or subtrees, only the tree, the data and a run, and the companion theorem shows that runs exist, so the goal is neither trivial nor vacuous.

Needed infrastructure: exchange arguments for ∑wjCj\sum w_jC_j∑wj​Cj​ on lists, restriction of sequences to modules, induction over decomposition trees. Theorems 1 and 2 and the ratio-order lemma are reusable in any 1 ∣ prec ∣ ∑wjCj1\,|\,\mathrm{prec}\,|\,\sum w_jC_j1∣prec∣∑wj​Cj​ development. Proofs of any milestone, and of the goal from the milestones, are welcome.

Selected references

  • E. L. Lawler, Sequencing jobs to minimize total weighted completion time subject to precedence constraints, IRIA-LABORIA Rapport de Recherche No. 205, December 1976. https://hal.science/hal-04716371v1. Journal version: Annals of Discrete Mathematics 2 (1978), 75–90. https://doi.org/10.1016/S0167-5060(08)70323-6
  • J. B. Sidney, Decomposition algorithms for single-machine sequencing with precedence relations and deferral costs, Operations Research 23 (1975), 283–298. https://doi.org/10.1287/opre.23.2.283
  • W. E. Smith, Various optimizers for single-stage production, Naval Research Logistics Quarterly 3 (1956), 59–66. https://doi.org/10.1002/nav.3800030106
  • J. Valdes, R. E. Tarjan, E. L. Lawler, The recognition of series parallel digraphs, SIAM Journal on Computing 11 (1982), 298–313. https://doi.org/10.1137/0211023
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Theoretical Computer Science·Captain: mikedeng1

Flowshop and Jobshop Schedules: Complexity and Approximation 6: Concatenating Optimal Schedules of Processor Pairs Gives Finish Time within ⌈m/2⌉ Times the OptimumResearch Paper

Motivation

A flow shop is the simplest model of a production line: every job passes through the same sequence of machines in the same order. Minimizing the time at which the last job leaves the last machine (the finish time, or makespan) is easy for two machines, where Johnson's rule gives an optimal schedule in time O(nlog⁡n)O(n\log n)O(nlogn) (Johnson 1954), and NP-hard from three machines on (Garey, Johnson, Sethi 1976; Gonzalez, Sahni 1978, §1, for preemptive schedules). For three or more machines the practical question is therefore how far a fast heuristic can be from the optimum in the worst case.

Gonzalez and Sahni (1978, §2) first show that any busy schedule, one that never leaves a processor idle while a task is available for it, has finish time at most mmm times the optimum (their Lemma 10), and that the natural "longest job first" order does no better. They then give a heuristic, H, that uses Johnson's two-machine algorithm as a subroutine and has the better worst-case ratio ⌈m/2⌉\lceil m/2\rceil⌈m/2⌉ (Lemma 11). This mission formalizes that bound.

Setting

A flow shop has mmm processors P1,…,PmP_1,\dots,P_mP1​,…,Pm​ and nnn jobs. Job iii consists of mmm tasks; task jjj runs on processor PjP_jPj​ and needs processing time tj,i≥0t_{j,i}\ge 0tj,i​≥0 (zero is allowed). A non-preemptive schedule gives every task a start time sj,is_{j,i}sj,i​, and the task occupies PjP_jPj​ during [sj,i,sj,i+tj,i)[s_{j,i},s_{j,i}+t_{j,i})[sj,i​,sj,i​+tj,i​). It is feasible when all start times are nonnegative, when task j+1j+1j+1 of a job starts only after task jjj of that job has completed (sj,i+tj,i≤sj+1,is_{j,i}+t_{j,i}\le s_{j+1,i}sj,i​+tj,i​≤sj+1,i​), and when no processor runs two tasks of positive length at the same time. Its finish time FT(s)\mathrm{FT}(s)FT(s) is the latest completion time sj,i+tj,is_{j,i}+t_{j,i}sj,i​+tj,i​ (and 000 if there is no task). An optimal finish time (OFT) schedule S∗S^*S∗ is a feasible schedule of least finish time.

Heuristic H (p. 48) divides the processors into ⌈m/2⌉\lceil m/2\rceil⌈m/2⌉ groups: group ggg consists of P2g−1P_{2g-1}P2g−1​ and P2gP_{2g}P2g​, and when mmm is odd the last group is PmP_mPm​ alone. The flow shop FgF_gFg​ on group ggg has the same jobs, with task times t2g−1,i,t2g,it_{2g-1,i},t_{2g,i}t2g−1,i​,t2g,i​ only. For each group an OFT schedule R(g)R(g)R(g) of FgF_gFg​ is computed (by Johnson's algorithm), with finish time f(R(g))f(R(g))f(R(g)). The schedule SSS generated by H runs these schedules one after another: every task of group ggg starts at its time in R(g)R(g)R(g) plus the offset ∑h<gf(R(h))\sum_{h<g}f(R(h))∑h<g​f(R(h)).

Formalization targets

Goal: Lemma 11 (p. 49)

For every flow shop with mmm processors and nnn jobs, every choice of OFT schedules R(g)R(g)R(g) of the group flow shops, the schedule SSS generated by H, and every feasible non-preemptive schedule τ\tauτ of the flow shop (in particular an OFT schedule S∗S^*S∗),

FT(S)≤⌈m2⌉ FT(τ).\mathrm{FT}(S)\le\Big\lceil\frac m2\Big\rceil\,\mathrm{FT}(\tau).FT(S)≤⌈2m​⌉FT(τ).

The paper writes this as FT(S)/FT(S∗)≤⌈m/2⌉\mathrm{FT}(S)/\mathrm{FT}(S^*)\le\lceil m/2\rceilFT(S)/FT(S∗)≤⌈m/2⌉.

Milestones

  1. Every group flow shop FgF_gFg​ has an OFT schedule (p. 48; the paper obtains it with Johnson's algorithm).
  2. The concatenation of feasible group schedules is a feasible schedule of the mmm-processor flow shop (p. 48).
  3. FT(S∗)≥max⁡gf(R(g))\mathrm{FT}(S^*)\ge\max_g f(R(g))FT(S∗)≥maxg​f(R(g)) (proof of Lemma 11, p. 49).
  4. FT(S)≤∑gf(R(g))≤⌈m/2⌉max⁡gf(R(g))\mathrm{FT}(S)\le\sum_g f(R(g))\le\lceil m/2\rceil\max_g f(R(g))FT(S)≤∑g​f(R(g))≤⌈m/2⌉maxg​f(R(g)) (proof of Lemma 11, p. 49).

Significance

The result gives a polynomial-time heuristic, running in time O(mnlog⁡n)O(mn\log n)O(mnlogn), whose finish time is within a factor ⌈m/2⌉\lceil m/2\rceil⌈m/2⌉ of the optimum on every instance, half the factor mmm that every busy schedule already achieves. The paper's Example 4 (p. 49) shows that for m=3m=3m=3 the factor 2=⌈3/2⌉2=\lceil 3/2\rceil2=⌈3/2⌉ is approached, so the analysis cannot be improved in general.

Formalizing it produces a reusable, machine-checked model of the mmm-processor flow shop with non-preemptive schedules, its restriction to blocks of consecutive processors, and the concatenation of schedules of disjoint blocks. Lemma 11 has a short paper proof; to our knowledge it has no machine-checked proof. The model's careful treatment of zero task times, of odd mmm, and of the restriction of a schedule to a processor group is itself part of the value: these are the points at which an informal argument is silent.

Difficulty

The arithmetic of the bound is short; the content lies in two structural facts about a precise model: that the concatenated object is a feasible schedule of the whole flow shop, including across the boundary between consecutive groups and for a final group of a single processor, and that the optimum of the whole flow shop is at least the optimum of each group flow shop, whose first processor has no predecessor. Informal arguments are silent on exactly these points. A naive formalization that leaves the offsets free, or that does not require the R(g)R(g)R(g) to be optimal, makes the statement false; one that lets SSS be any schedule with the right restrictions makes it witnessed by S∗S^*S∗ itself.

The existence of OFT schedules for the group flow shops (milestone 1) is a separate matter: the paper takes it from Johnson's algorithm, and proving it requires an optimality argument for two-machine flow shops (or a compactness argument over finitely many processing orders).

Formalization scope

  • Representation. Processors are Fin m and jobs Fin n, both 0-based; times are real numbers. A flow shop is a matrix t : Fin m → Fin n → ℝ with t ≥ 0. A schedule is a matrix of start times. Group ggg (0-based, g : Fin ((m + 1) / 2)) contains the processors 2g2g2g and 2g+12g+12g+1 that exist, min 2 (m - 2g) of them.
  • Finish time. A fold of max with baseline 000 over all tasks; it is 000 for m=0m=0m=0 or n=0n=0n=0. The maximum of the group finish times is also a fold of max with baseline 000. No supremum over an unbounded or empty set is taken.
  • Zero task times. A zero-length task occupies no processor time (it is exempt from the non-overlap condition) but still respects the job's task order and counts in the finish time when it is placed late; optimal schedules are unaffected.
  • The optimum. "OFT schedule S∗S^*S∗" is replaced by a universally quantified feasible schedule τ\tauτ, which is stronger and needs no existence assumption for S∗S^*S∗. Optimality of R(g)R(g)R(g) is the predicate "feasible, and finish time at most that of every feasible schedule of the same group flow shop".
  • Constants. ⌈m/2⌉\lceil m/2\rceil⌈m/2⌉ is (m + 1) / 2 in N\mathbb NN, cast to R\mathbb RR; the two agree for every m≥0m\ge 0m≥0. The ratio is multiplied out, so a zero optimum does not divide by zero.
  • Not formalized. Johnson's algorithm and the running times O(nlog⁡n)O(n\log n)O(nlogn) and O(mnlog⁡n)O(mn\log n)O(mnlogn). The goal quantifies over every family of optimal group schedules, which includes the one Johnson's algorithm computes; the paper's proof uses only optimality. The preemptive comparison on p. 50 and the tightness examples are outside the mission.
  • Trivialization ruled out. The schedule SSS is a function of the R(g)R(g)R(g) with the fixed offsets ∑h<gf(R(h))\sum_{h<g}f(R(h))∑h<g​f(R(h)), and each R(g)R(g)R(g) must be optimal for its group flow shop; neither the offsets nor SSS are free.
  • Infrastructure. Elementary facts about Finset.fold max and sums over Fin; nothing beyond Mathlib. Contributions proving milestone 1 by formalizing Johnson's rule for this model are welcome and reusable.

Selected references

  • T. Gonzalez, S. Sahni, Flowshop and Jobshop Schedules: Complexity and Approximation, Operations Research 26(1), 36–52, 1978. https://doi.org/10.1287/opre.26.1.36
  • S. M. Johnson, Optimal two- and three-stage production schedules with setup times included, Naval Research Logistics Quarterly 1(1), 61–68, 1954. https://doi.org/10.1002/nav.3800010110
  • M. R. Garey, D. S. Johnson, R. Sethi, The Complexity of Flowshop and Jobshop Scheduling, Mathematics of Operations Research 1(2), 117–129, 1976. https://doi.org/10.1287/moor.1.2.117
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CombinatoricsProbability·Captain: mikedeng1

Subadditive Ergodic Theory 6: The Ulam–Hammersley Constant of the Longest Increasing Subsequence Satisfies (8/π)^{1/2} ≤ c ≤ δ^{1/2} + δ^{−1/2}Research Paper

Motivation

Ulam asked, in the early 1960s, how long the longest increasing subsequence of a random permutation typically is. The question is a test case for a whole class of problems in combinatorial probability: a quantity defined by an optimisation over exponentially many configurations (here, all subsequences) whose typical size has to be found without enumerating them. The same structure appears in last-passage percolation, in the patience-sorting card game, in sequence comparison, and in scheduling problems where the longest chain of a random partial order determines a makespan.

J. F. C. Kingman's IMS Special Invited Paper Subadditive ergodic theory (Ann. Probab. 1 (1973), 883–899) presents Ulam's problem in §2.4 as an application of the subadditive ergodic theorem, following J. M. Hammersley, and then proves explicit bounds on the limiting constant. This mission formalizes those bounds.

Timeline.

  • Ulam (1961) raised the question and conjectured, from simulations, that the typical length grows like a constant times n\sqrt nn​.
  • Erdős and Szekeres (1935) had already shown deterministically that every permutation of nnn elements has an increasing or a decreasing subsequence of length at least n\sqrt nn​.
  • Hammersley (A few seedlings of research, Proc. Sixth Berkeley Symp. 1 (1972), 345–394) proved, by embedding the permutation in a planar Poisson process and applying a subadditivity argument, that n−1/2n^{-1/2}n−1/2 times the length converges in probability to a constant ccc, and showed 12π≤c≤e\tfrac12\pi\le c\le e21​π≤c≤e.
  • Kingman (1973), the paper of this mission, refined the bounds to (8/π)1/2≤c≤δ1/2+δ−1/2(8/\pi)^{1/2}\le c\le\delta^{1/2}+\delta^{-1/2}(8/π)1/2≤c≤δ1/2+δ−1/2, that is 1.59<c<2.491.59<c<2.491.59<c<2.49.
  • Logan and Shepp (1977) and Vershik and Kerov (1977) later identified the constant exactly. That later work is not part of this mission.

Setting

Let Sn\mathcal S_nSn​ be the group of permutations of {1,2,…,n}\{1,2,\dots,n\}{1,2,…,n}. For σ∈Sn\sigma\in\mathcal S_nσ∈Sn​, a sequence of positions 1≤i1<i2<⋯<ik≤n1\le i_1<i_2<\dots<i_k\le n1≤i1​<i2​<⋯<ik​≤n is ascending if σ(i1)<σ(i2)<⋯<σ(ik)\sigma(i_1)<\sigma(i_2)<\dots<\sigma(i_k)σ(i1​)<σ(i2​)<⋯<σ(ik​). The length of the longest ascending sequence, l(σ)l(\sigma)l(σ), is the largest such kkk.

The uniform distribution on Sn\mathcal S_nSn​ gives each of the n!n!n! permutations probability 1/n!1/n!1/n!, so an event A⊆SnA\subseteq\mathcal S_nA⊆Sn​ has probability P{A}=∣A∣/n!P\{A\}=|A|/n!P{A}=∣A∣/n!. Let πn\pi_nπn​ denote a permutation drawn from this distribution. The sequence n−1/2 l(πn)n^{-1/2}\,l(\pi_n)n−1/2l(πn​) converges in probability to a real number ccc if, for every ε>0\varepsilon>0ε>0,

P{∣n−1/2 l(πn)−c∣>ε}→0(n→∞).P\{|n^{-1/2}\,l(\pi_n)-c|>\varepsilon\}\to0\qquad(n\to\infty).P{∣n−1/2l(πn​)−c∣>ε}→0(n→∞).

For integers k≥1k\ge1k≥1, νk(σ)\nu_k(\sigma)νk​(σ) denotes the number of ascending sequences of length kkk in σ\sigmaσ. For 0<α<b0<\alpha<b0<α<b write

E(α,b)=2α+(b−α)log⁡(b−α)−αlog⁡α−blog⁡b,E(\alpha,b)=2\alpha+(b-\alpha)\log(b-\alpha)-\alpha\log\alpha-b\log b ,E(α,b)=2α+(b−α)log(b−α)−αlogα−blogb,

the exponent appearing in Kingman's condition (2.4.8). (The paper calls the second variable β\betaβ; in Lean it is bbb, because β\betaβ also names the constant below.)

Formalization targets

Goal: Theorem 8

Let ccc be the limit in probability of n−1/2l(πn)n^{-1/2}l(\pi_n)n−1/2l(πn​). Let δ\deltaδ be the unique positive root of log⁡(1+δ)=2δ/(1+δ)\log(1+\delta)=2\delta/(1+\delta)log(1+δ)=2δ/(1+δ), and β=δ1/2+δ−1/2\beta=\delta^{1/2}+\delta^{-1/2}β=δ1/2+δ−1/2. Then

(8π)1/2≤c≤β(2.4.4)\Big(\frac8\pi\Big)^{1/2}\le c\le\beta \tag{2.4.4}(π8​)1/2≤c≤β(2.4.4)

and hence

1.59<c<2.49.(2.4.5)1.59<c<2.49. \tag{2.4.5}1.59<c<2.49.(2.4.5)

Numerically δ≈3.9216\delta\approx3.9216δ≈3.9216 and β≈2.4853\beta\approx2.4853β≈2.4853. The goal states (2.4.4) and (2.4.5) as printed, together with the existence and uniqueness of δ\deltaδ.

Milestones

  1. Theorem 7 (Hammersley). There is an absolute constant ccc such that n−1/2l(πn)→cn^{-1/2}l(\pi_n)\to cn−1/2l(πn​)→c in probability.
  2. The double integral of the lower-bound argument: ∫0∞ ⁣∫0∞x e−12(x+y)2 dx dy=(π/8)1/2\int_0^\infty\!\int_0^\infty x\,e^{-\frac12(x+y)^2}\,dx\,dy=(\pi/8)^{1/2}∫0∞​∫0∞​xe−21​(x+y)2dxdy=(π/8)1/2.
  3. First moment of ν\nuν: for k≥1k\ge1k≥1, E(νk)=(nk)(k!)−1E(\nu_k)=\binom nk(k!)^{-1}E(νk​)=(kn​)(k!)−1.
  4. Subsequence count: if l(σ)≥kl(\sigma)\ge kl(σ)≥k then νk(σ)≥(l(σ)k)\nu_k(\sigma)\ge\binom{l(\sigma)}kνk​(σ)≥(kl(σ)​).
  5. (2.4.7): for 1≤k≤r1\le k\le r1≤k≤r, P{l(π)≥r}≤(nk)[k!(rk)]−1P\{l(\pi)\ge r\}\le\binom nk\big[k!\binom rk\big]^{-1}P{l(π)≥r}≤(kn​)[k!(kr​)]−1.
  6. (2.4.8): if 0<α<b0<\alpha<b0<α<b and E(α,b)<0E(\alpha,b)<0E(α,b)<0, then P{l(π)≥b n1/2}→0P\{l(\pi)\ge b\,n^{1/2}\}\to0P{l(π)≥bn1/2}→0.
  7. The optimisation: inf⁡{b>0:∃α∈(0,b), E(α,b)<0}=δ1/2+δ−1/2\inf\{b>0:\exists\alpha\in(0,b),\ E(\alpha,b)<0\}=\delta^{1/2}+\delta^{-1/2}inf{b>0:∃α∈(0,b), E(α,b)<0}=δ1/2+δ−1/2.

Significance

The result. Theorem 8 was, at the time, the sharpest rigorous information about Ulam's constant. Its upper bound comes from a first-moment computation that applies to any random structure in which long increasing chains can be counted; the same calculation bounds the longest chain in random partial orders and the height of random kkk-dimensional orders. The lower bound shows that a greedy path through a Poisson process already achieves a positive fraction of the optimum, an argument reused for last-passage percolation models.

Formalizing it. All results here are proved in the literature; none is open. To the best of our knowledge none of them has been machine-checked. A complete development produces: an exact non-asymptotic tail bound on the longest increasing subsequence ((2.4.7)), useful on its own; a Stirling-type asymptotic for products of binomial coefficients at scale n\sqrt nn​; and, through Theorem 7, a formal existence proof of the Ulam constant. Theorem 7 is the substantial piece: the paper's proof is a sketch through a continuous-parameter subadditive ergodic theorem and a planar Poisson process.

Difficulty

The upper bound needs care with asymptotics: (2.4.7) holds for each nnn, but the passage to (2.4.8) uses Stirling's formula with kkk and rrr integers tending to infinity at rate n\sqrt nn​, and the infimum in milestone 7 is not attained (the admissible set is open), so the argument must produce, for every b>βb>\betab>β, a suitable α\alphaα.

The lower bound is where the obvious approach fails. A first-moment or counting argument gives only upper bounds on lll; a lower bound requires exhibiting long ascending sequences. The paper's argument does this with a greedy path in a planar Poisson process, which is not a statement about uniform permutations and needs the Poissonization of Theorem 7 to transfer back. Proving c≥(8/π)1/2c\ge(8/\pi)^{1/2}c≥(8/π)1/2 directly from the counting definition of lll is not available. Both bounds also need the existence of ccc, which is Theorem 7.

Formalization scope

  • Sn\mathcal S_nSn​ is Equiv.Perm (Fin n); the permutation is named σ\sigmaσ in Lean because π\piπ is Real.pi.
  • l(σ)l(\sigma)l(σ) is the largest cardinality of a finite set of positions on which σ\sigmaσ is strictly increasing; νk(σ)\nu_k(\sigma)νk​(σ) is the number of such sets with exactly kkk elements.
  • Probabilities under the uniform law are proportions ∣A∣/n!|A|/n!∣A∣/n!. Convergence in probability depends only on the law of each πn\pi_nπn​ (the paper notes this on p. 895), so no common probability space is introduced. Time is n∈Nn\in\mathbb Nn∈N, and at n=0n=0n=0 Lean reads l/0l/\sqrt0l/0​ as 000, which does not affect a limit.
  • The constant of Theorem 8 enters as a hypothesis: the goal is stated for every real ccc to which n−1/2l(πn)n^{-1/2}l(\pi_n)n−1/2l(πn​) converges in probability. Such a limit is unique, and Theorem 7 shows it exists. No sign or bound on ccc is assumed. The goal and milestone 7 assert that the positive root δ\deltaδ exists and is unique, so the upper bound cannot hold vacuously for want of a root.
  • The double integral is stated as a value only; the claim that it is the mean of i.i.d. increments of a Poisson greedy path is not formalized.

Needed infrastructure: counting arguments on Finset.powersetCard; Stirling's formula (Mathlib has Stirling.tendsto_stirlingSeq_sqrt_pi); elementary calculus for the root δ\deltaδ; Gaussian-type integrals on the quadrant. Theorem 7 additionally needs Poisson point processes in the plane, which Mathlib lacks, or another route to the existence of the limit. Contributions of that infrastructure are welcome and reusable well beyond this mission. The counting bounds (milestones 3–5) are independent of Theorem 7 and are a natural first target.

Selected references

  • J. F. C. Kingman, Subadditive ergodic theory, Annals of Probability 1(6) (1973), 883–899. https://doi.org/10.1214/aop/1176996798
  • J. M. Hammersley, A few seedlings of research, Proc. Sixth Berkeley Symposium on Mathematical Statistics and Probability, Vol. 1 (1972), 345–394. https://projecteuclid.org/euclid.bsmsp/1200514101
  • S. M. Ulam, Monte Carlo calculations in problems of mathematical physics, in Modern Mathematics for the Engineer, Second Series, McGraw-Hill (1961), 261–281.
  • P. Erdős and G. Szekeres, A combinatorial problem in geometry, Compositio Mathematica 2 (1935), 463–470. http://www.numdam.org/item/CM_1935__2__463_0/
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Dynamical SystemsProbabilityStochastic Systems·Captain: mikedeng1

Subadditive Ergodic Theory 3: A Separable Continuous-Parameter Subadditive Process with E(Ω_I) < ∞ Has x₀ₜ/t → ξ with Probability One and in Mean, E(ξ) = γResearch Paper

Motivation

Many random growth models attach a cost, distance, or accumulated quantity to each time interval. If the quantity over a long interval is at most the sum over two adjacent intervals, it is subadditive. A stable rate over long intervals is useful even when the underlying random system does not have a simple independent-increment description. Kingman's 1973 account places percolation and products of random matrices among the applications of this framework. Its basic theorem treats integer time; the present mission concerns the change to every nonnegative real time. Kingman, §1.1 and §1.4.

The change matters because observing a process at integer times can miss arbitrarily large behavior between observations. Kingman gives a smooth stationary-increment example in Theorem 3 where no prescribed increasing bound eventually controls the values at all real times. Theorem 4 gives a condition that does control this local behavior. These two results together explain why the continuous-time statement needs more than a formal change of index type. Kingman, Theorems 3–4.

Setting

Work on a probability space (S,F,P)(S,\mathcal F,P)(S,F,P). A continuous-parameter subadditive process is a family of real random variables xstx_{st}xst​ indexed by 0≤s<t0\le s<t0≤s<t. It satisfies xsu≤xst+xtux_{su}\le x_{st}+x_{tu}xsu​≤xst​+xtu​ whenever 0≤s<t<u0\le s<t<u0≤s<t<u. Its stationarity says that, for every real shift τ≥0\tau\ge0τ≥0, the complete shifted family (xs+τ,t+τ)s<t(x_{s+\tau,t+\tau})_{s<t}(xs+τ,t+τ​)s<t​ has the same joint distribution as (xst)s<t(x_{st})_{s<t}(xst​)s<t​. Finally, the mean gt=E(x0t)g_t=E(x_{0t})gt​=E(x0t​) exists as a finite real number for each t>0t>0t>0 and obeys gt≥−Atg_t\ge-Atgt​≥−At for one constant AAA. These are the continuous-time versions of Kingman's conditions S₁, S₂, and S₃. Equality of the law of each individual coordinate would be too weak for S₂. Kingman, pp. 883–884, 887.

For an interval I⊆[0,∞)I\subseteq[0,\infty)I⊆[0,∞), the oscillation is ΩI=sup⁡s<t, s,t∈I∣xst∣\Omega_I=\sup_{s<t,\,s,t\in I}|x_{st}|ΩI​=sups<t,s,t∈I​∣xst​∣. A process has the local oscillation condition when E(ΩI)<∞E(\Omega_I)<\inftyE(ΩI​)<∞ for one nondegenerate interval III. In §1.4 the paper states that this condition then holds on every bounded interval. The quantity γ=inf⁡t>0gt/t\gamma=\inf_{t>0}g_t/tγ=inft>0​gt​/t is the mean growth rate. Kingman, (1.4.1), (1.4.6)–(1.4.7).

The paper also requires separability. Here this means that, outside one probability-zero set, the value at every admissible pair of times can be approached along a sequence from a fixed countable dense set of pairs. It is a condition on how the uncountable family is observed, and it permits discontinuous sample paths. Kingman cites the standard separability convention without defining it in §1.4; the mission makes the convention explicit. Kingman, Theorem 4 and following discussion.

Formalization targets

Theorem 4: convergence at all real times

For a separable process satisfying S₁–S₃ and the local oscillation condition, the target is an integrable real random variable ξ\xiξ such that

x0tt⟶ξalmost surely and in L1(P)(t→∞, t∈R),E(ξ)=γ=inf⁡u>0guu.\frac{x_{0t}}{t}\longrightarrow\xi\quad\text{almost surely and in }L^1(P) \quad(t\to\infty,\ t\in\mathbb R), \qquad E(\xi)=\gamma=\inf_{u>0}\frac{g_u}{u}.tx0t​​⟶ξalmost surely and in L1(P)(t→∞, t∈R),E(ξ)=γ=u>0inf​ugu​​.

The time variable ranges over all real values tending to infinity. The integer-time assertion (1.4.8) is a separate milestone, not the goal. Other milestones identify the deterministic mean rate, extend the oscillation condition to bounded intervals, record the bounds at times between consecutive integers, and state the vanishing of normalized unit-interval oscillation. Each comes from a sentence or display in §1.4. Kingman, Theorem 4 and proof.

Significance

Theorem 4 gives a long-run rate for a broad stationary subadditive family without assuming that its trajectories are continuous or differentiable. Its mean statement identifies the expectation of the random limit with a deterministic infimum of normalized expectations. The paper notes that the usual continuous-time ergodic theorem for additive integral processes is a corollary, because such processes satisfy its oscillation condition. The condition is therefore compatible with familiar additive models while also covering discontinuous subadditive ones. Kingman, p. 890.

The theorem and its argument are known from the paper; this mission asks for a machine-checked development of the stated result. A completed development would provide reusable formal interfaces for joint-law stationarity of an uncountable family, separability, extended-valued oscillations, and almost-sure plus L1L^1L1 limits. The proposed theorem statements compile as open goals; compiling them does not prove them. Kingman, §1.4.

Difficulty

An integer-time subadditive theorem controls x0n/nx_{0n}/nx0n​/n only on its discrete skeleton. It cannot by itself rule out large values of x0t/tx_{0t}/tx0t​/t for times between consecutive integers; Theorem 3 shows that such behavior can occur even with smooth stationary sample paths and finite expectations. A continuous-time result must control local variation as the observation window moves to large times. The uncountable supremum in ΩI\Omega_IΩI​ also needs a genuine extended value and a measurability convention. Kingman, Theorems 3–4.

The paper's displayed assertion (1.4.1), that gt/tg_t/tgt​/t converges as real t→∞t\to\inftyt→∞ under S₁–S₃ alone, needs qualification: a discontinuous additive function can make a deterministic subadditive mean with no such limit. The corresponding milestone includes the finite-oscillation condition assumed by Theorem 4. This correction is disclosed in the theorem's formalization note. Kingman, (1.4.1), (1.4.7).

Formalization scope

Lean indexes the family by R→R→S→R\mathbb R\to\mathbb R\to S\to\mathbb RR→R→S→R and uses only 0≤s<t0\le s<t0≤s<t; values at other pairs carry no meaning. Each valid coordinate is measurable. S₁ is pointwise in the sample point, while S₂ is equality of the probability laws of the complete path before and after every nonnegative real shift. S₃ states integrability of x0tx_{0t}x0t​ for every t>0t>0t>0 and one linear lower bound on its mean. The sample-space type is called SSS so that ΩI\Omega_IΩI​ denotes oscillation without ambiguity.

Oscillation is valued in the extended nonnegative reals. Thus an unbounded supremum is infinite, and the expected-oscillation condition cannot hold because of a default real-supremum value. Separability is the countable graph-approximation condition outside one measurable null set. The premise is finiteness on one closed interval [a,b][a,b][a,b] with 0≤a<b0\le a<b0≤a<b; the propagation milestone covers bounded intervals with either choice of endpoints. The limit in Theorem 4 uses the filter at infinity on real time. Its almost-sure convergence and convergence in mean are separate clauses, and ξ\xiξ is integrable. No continuity of paths, ergodicity, or independence is assumed.

A complete proof may build on Mathlib's measure theory, filters, Borel–Cantelli lemmas, and real analysis. The process and oscillation definitions, along with the finite-oscillation and skeleton statements, are reusable beyond this particular theorem. Contributions that establish these interfaces and the named intermediate statements are within scope. A statement limited to integer times, a marginal-only version of stationarity, or a real supremum that returns a default value on an unbounded set would not express this target.

Selected references

  • J. F. C. Kingman, Subadditive ergodic theory, The Annals of Probability 1(6):883–899, 1973. DOI: 10.1214/aop/1176996798.
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Dynamical SystemsProbabilityStochastic Systems·Captain: mikedeng1

Subadditive Ergodic Theory 1: Kingman's Theorem — x₀ₜ/t Converges with Probability One and in Mean to a Finite ξ with E(ξ) = γResearch Paper

Motivation

Many stochastic systems have an accumulated cost or passage time that grows with the length of an observation window. Splitting a window at an intermediate time gives an upper bound on the total, even when it does not give an equality. First-passage percolation and products of random matrices are examples discussed by Kingman (1973). A central question is whether the cost per unit time settles to a long-run rate on almost every sample path.

Hammersley and Welsh introduced subadditive stochastic processes in 1965. Kingman's 1968 theorem established the almost-sure limit for the stationary process considered here. His 1973 paper restated the result, identified the limit's expectation, and set out the process model and related formulations used in later applications. The 1973 article explicitly attributes the proof of Theorem 1, its maximal inequality, and its decomposition to the 1968 work; this mission formalizes the statements as presented in the 1973 paper.

Setting

Fix a probability space (Ω,F,P)(\Omega,\mathcal F,P)(Ω,F,P). For integers 0≤s<t0\le s<t0≤s<t, let xst:Ω→Rx_{st}:\Omega\to\mathbb Rxst​:Ω→R be a measurable random variable. Think of xstx_{st}xst​ as the accumulated quantity over the interval from sss to ttt. Condition S₁ is subadditivity: xsu≤xst+xtux_{su}\le x_{st}+x_{tu}xsu​≤xst​+xtu​ whenever s<t<us<t<us<t<u. The stronger equality defines an additive process.

Condition S₂ says that shifting every interval one time unit forward leaves the joint distribution of the whole family unchanged. It concerns the law of (xs+1,t+1)s<t(x_{s+1,t+1})_{s<t}(xs+1,t+1​)s<t​ as a single countable path. The one-dimensional statement that each xstx_{st}xst​ has a distribution depending only on t−st-st−s is called S₂′ in the paper; it does not replace S₂. Condition S₃ says that every positive-time variable x0tx_{0t}x0t​ is integrable and that, for some real AAA, its mean gt=EP(x0t)g_t=E_P(x_{0t})gt​=EP​(x0t​) satisfies gt≥−Atg_t\ge-Atgt​≥−At for all t≥1t\ge1t≥1. A subadditive process satisfies S₁, S₂, and S₃.

The associated mean growth rate is γ(x)=inf⁡t≥1gt/t\gamma(x)=\inf_{t\ge1}g_t/tγ(x)=inft≥1​gt​/t. The time index starts at one in this infimum because division by zero would introduce an unrelated value. S₃ gives a finite lower bound; the subadditive mean relation makes this infimum the limit of gt/tg_t/tgt​/t. Kingman also considers the invariant σ-field I\mathcal II: events determined by the entire path and unchanged by the shift (xst)↦(xs+1,t+1)(x_{st})\mapsto(x_{s+1,t+1})(xst​)↦(xs+1,t+1​). It records the part of the process that can remain random in the limiting rate. Source: §§1.1–1.2, pp. 883–885.

Formalization targets

Theorem 1: finite long-run rate

For every subadditive process there is an integrable real random variable ξ\xiξ such that

x0t(ω)t⟶ξ(ω)for P-almost every ω,∫Ω∣x0tt−ξ∣ dP⟶0,EP(ξ)=γ(x).\frac{x_{0t}(\omega)}{t}\longrightarrow\xi(\omega) \quad\text{for }P\text{-almost every }\omega, \qquad \int_\Omega\left|\frac{x_{0t}}{t}-\xi\right|\,dP\longrightarrow0, \qquad E_P(\xi)=\gamma(x).tx0t​(ω)​⟶ξ(ω)for P-almost every ω,∫Ω​​tx0t​​−ξ​dP⟶0,EP​(ξ)=γ(x).

These are three distinct claims: almost-sure convergence, convergence in mean, and an identity for the limit's expectation. The limit need not be a constant. The milestone list also covers the convergence of the means, the maximal inequality (1.2.5), the decomposition (1.2.6), the measure-preserving formulation (1.3.4)–(1.3.5), and the representation through I\mathcal II in (1.2.4). Each statement comes from the 1973 paper, §§1.1–1.3.

Measure-preserving formulation

For a measure-preserving map θ:Ω→Ω\theta:\Omega\to\Omegaθ:Ω→Ω, condition Sᴱ concerns integrable functions fnf_nfn​ with fm+n(ω)≤fm(ω)+fn(θmω)f_{m+n}(\omega)\le f_m(\omega)+f_n(\theta^m\omega)fm+n​(ω)≤fm​(ω)+fn​(θmω) for m,n≥1m,n\ge1m,n≥1 and a linear lower bound on their expectations. It yields an integrable ξ\xiξ with fn/n→ξf_n/n\to\xifn​/n→ξ almost surely and in mean. Neither invertibility nor ergodicity of θ\thetaθ is assumed. Source: §1.3, p. 886.

Significance

The theorem assigns a finite sample-path growth rate to an entire class of stationary subadditive systems. It also ties the average of that rate to the infimum of the normalized expected values. The equality is useful when a system's finite-window means are easier to estimate than its individual long-run paths. The invariant σ-field statement explains why a stationary system can have different rates on different paths without contradicting the mean identity. Kingman's examples show how the same framework reaches percolation and matrix products.

The result is mathematically proved, as the 1973 article states. The remaining work here is a machine-checked development of its process model and the displayed results. A published Prove2Me item, PalmQueueing.Ergodic.kingman_subadditive, is open and concerns an ergodic bijective flow with a constant possibly infinite limit; it does not supply this nonergodic finite-limit theorem or convergence in mean. The draft statements in this mission compile with proof placeholders. No machine-checked proof of these new statements is claimed.

Difficulty

Subadditivity gives bounds when intervals are joined, but it does not express x0tx_{0t}x0t​ as a sum of identically distributed increments. The ordinary additive ergodic theorem therefore does not immediately give convergence for x0t/tx_{0t}/tx0t​/t. Nor does convergence of the expectations gt/tg_t/tgt​/t alone force convergence on individual paths or in mean. The process may also have a nonconstant limit: stationarity is a joint-law condition, not an ergodicity assumption that erases invariant information. These are the gaps the formalization must close without strengthening the hypotheses. Source: §§1.1–1.3.

Formalization scope

Lean represents xxx as a function of two natural-number indices and a sample point, but every condition and conclusion uses only s<ts<ts<t. S₁ and the Sᴱ cocycle inequality follow the paper's pointwise displays; random-variable equality in the decomposition is stated almost surely at each valid index pair. Measurability of each valid coordinate is explicit. S₂ compares two probability measures on the full countable path space, so equality of individual marginals cannot satisfy it. The probability measure is normalized, and S₃ requires Bochner integrability to keep every mean finite. The constant γ\gammaγ is a real infimum over positive times, used only under S₃. The goal's limit is real and integrable; Kingman's separate Theorem 2 treats a possible value of −∞-\infty−∞ and belongs to another mission.

The invariant σ-field is the pullback of the path-space shift-invariant σ-field. Its measurability claim is made through an almost-everywhere equal version, since almost-sure limits are insensitive to changes on null sets. Conditional expectation uses this explicit sub-σ-field. All limits in this mission take positive integer time to infinity. The Sᴱ definition requires measure preservation but no inverse and no ergodicity. Thus neither a marginal-law substitution for S₂ nor an ergodic constant-limit specialization can trivialize the target.

A complete proof can reuse Mathlib's measure theory, conditional expectation, product measurable spaces, measure-preserving maps, and subadditive-sequence results. The two process definitions, the invariant σ-field interface, and auxiliary statements are useful to later formalizations of stationary growth models. Contributions that prove the stated results or develop general-purpose lemmas for these objects fit the mission.

Selected references

  • J. F. C. Kingman, Subadditive ergodic theory, The Annals of Probability 1(6), 883–899, 1973. DOI: 10.1214/aop/1176996798.
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Theoretical Computer Science·Captain: mikedeng1

Flowshop and Jobshop Schedules: Complexity and Approximation 5: Sequencing Jobs by Shortest Total Processing Time Gives Mean Flow Time within m Times the OptimumResearch Paper

Motivation

Flow shops and job shops are the basic models of multi-stage production: every job passes through several machines, and each machine handles one task at a time. Gonzalez and Sahni (Operations Research 26(1), 1978) showed that minimizing the finish time of such shops is NP-complete even with preemption, and then asked how well simple heuristics do. One of the two objectives they study is the mean flow time, the average completion time of the jobs, which measures how long a job spends in the shop on average.

Their Lemma 9 analyses the most natural rule for that objective, shortest processing time first (SPT): process the jobs in order of nondecreasing total work. For a single machine SPT is optimal (Smith 1956; Conway, Maxwell and Miller, Theory of Scheduling, 1967, p. 76, which the paper cites). Lemma 9 shows that with mmm processors the same rule loses at most a factor mmm, and the paper's Example 2 shows that the factor mmm is attained in the limit.

Setting

A job shop has m≥1m\ge1m≥1 processors P1,…,PmP_1,\dots,P_mP1​,…,Pm​ and nnn jobs. Job iii is a finite sequence of tasks; each task names the processor that must process it and a processing time p≥0p\ge0p≥0. The tasks of a job are processed in their order: a task may start only after the job's previous task has completed. A flow shop is the special case in which every job has mmm tasks and its kkk-th task runs on PkP_kPk​.

A non-preemptive schedule gives every task a start time s≥0s\ge0s≥0; the task then occupies its processor during [s,s+p)[s,s+p)[s,s+p), and two distinct tasks on one processor never overlap. The finish time fi(S)f_i(S)fi​(S) of job iii in schedule SSS is the time at which all tasks of job iii have been completed, and the mean flow time is

MFT(S)=1n∑i=1nfi(S).\mathrm{MFT}(S)=\frac1n\sum_{i=1}^n f_i(S).MFT(S)=n1​i=1∑n​fi​(S).

An OMFT schedule S∗S^*S∗ is a feasible schedule of least mean flow time.

Let LiL_iLi​ be the sum of the task times of job iii. An SPT order is a listing σ(0),…,σ(n−1)\sigma(0),\dots,\sigma(n-1)σ(0),…,σ(n−1) of the jobs with Lσ(0)≤⋯≤Lσ(n−1)L_{\sigma(0)}\le\dots\le L_{\sigma(n-1)}Lσ(0)​≤⋯≤Lσ(n−1)​, ties broken arbitrarily. The SPT schedule processes the jobs in that order: each job in turn, each positive-time task starting at the later of the completion of the job's previous task and the latest completion of a task already placed on its processor. A zero-time task completes when its preceding task completes and leaves the processor available.

Formalization targets

Goal: Lemma 9

For every job shop, every SPT order σ\sigmaσ with SPT schedule SSS, and every feasible non-preemptive schedule τ\tauτ,

MFT(S)≤m⋅MFT(τ),\mathrm{MFT}(S)\le m\cdot\mathrm{MFT}(\tau),MFT(S)≤m⋅MFT(τ),

so that MFT(S)/MFT(S∗)≤m\mathrm{MFT}(S)/\mathrm{MFT}(S^*)\le mMFT(S)/MFT(S∗)≤m for an OMFT schedule S∗S^*S∗. A separate item states the same bound for flow shops.

Milestones

The milestones follow the paper's proof of Lemma 9 on p. 47:

  1. the SPT schedule is a feasible schedule;
  2. the job in position kkk of the SPT schedule finishes by ∑j≤kLσ(j)\sum_{j\le k}L_{\sigma(j)}∑j≤k​Lσ(j)​;
  3. in any feasible schedule, if i1,…,ini_1,\dots,i_ni1​,…,in​ is the order in which the jobs finish, then fik≥∑j≤kLij/mf_{i_k}\ge\sum_{j\le k}L_{i_j}/mfik​​≥∑j≤k​Lij​​/m;
  4. a prefix sum of any listing of the LiL_iLi​ is at least the corresponding prefix sum of the sorted listing;
  5. MFT(S∗)≥1n∑k=1n∑j=1kLj/m\mathrm{MFT}(S^*)\ge\frac1n\sum_{k=1}^n\sum_{j=1}^k L_j/mMFT(S∗)≥n1​∑k=1n​∑j=1k​Lj​/m.

Significance

Lemma 9 is one of the earliest performance guarantees for a shop-scheduling heuristic under a sum objective. It shows that a rule computable by one sort, O(nlog⁡n)O(n\log n)O(nlogn) plus a linear pass, is within a factor equal to the number of machines, independently of the number of jobs. This is better than the factor nnn that Lemma 8 of the same paper gives for an arbitrary busy schedule, whenever m<nm<nm<n. Example 2 of the paper shows that the factor cannot be improved for SPT. Lower bounds of the work-per-machine type used here recur in later approximation results for total completion time in shops and on parallel machines.

The result has been proved since 1978. As far as a search of the Prove2Me catalogue shows, it has not been machine-checked. This mission adds a machine-checked statement of the SPT schedule as an algorithm, not as an arbitrary schedule with a property. It also adds the guarantee for job shops with arbitrary task sequences, of which flow shops are a special case.

Difficulty

The arithmetic of the proof is short. The work is in two places. The SPT schedule is a concrete construction, a fold over jobs and tasks that updates processor availability, so the bound on its finish times has to be carried through an invariant of that construction. The lower bound on every feasible schedule needs a packing argument: tasks on one processor that complete by a given time have total length at most that time. That argument is not available in Mathlib in a form that applies here. The tempting shortcut of comparing the two schedules job by job fails: the optimal schedule need not finish jobs in SPT order, so the comparison must go through sums of sorted prefixes.

Formalization scope

  • Instance and feasibility. The job data come from the published JobShopLTAS.Core.Instance (jobs Fin n, processors Fin m, tasks Fin (μ j), real processing times ≥0\ge0≥0). The local IsPaperFeasibleSchedule requires nonnegative starts, in-job order, and no overlap of positive-time tasks on the same processor. Zero-time tasks occupy no processor interval. The paper assumes m≥1m\ge1m≥1, and the scheduling theorems carry that hypothesis.
  • Finish time and MFT. fif_ifi​ is the maximum completion time over the tasks of job iii, with baseline 000; MFT divides by nnn, and for n=0n=0n=0 both sides are 000.
  • The SPT schedule is the list schedule constructed by the heuristic, for every SPT order, so ties are broken in every possible way. Zero-time tasks do not advance processor availability. It is not "any feasible schedule whose jobs finish in SPT order", a reading under which an optimal schedule would qualify and the goal would be trivial.
  • The optimum. The paper's "Let S∗S^*S∗ be an OMFT schedule" is stated as "for every feasible non-preemptive schedule τ\tauτ". This is equivalent and does not assume that an optimum exists.
  • Ratios multiplied out. MFT(S)/MFT(S∗)≤m\mathrm{MFT}(S)/\mathrm{MFT}(S^*)\le mMFT(S)/MFT(S∗)≤m is stated as MFT(S)≤m⋅MFT(τ)\mathrm{MFT}(S)\le m\cdot\mathrm{MFT}(\tau)MFT(S)≤m⋅MFT(τ), and the proof's bounds divided by mmm are multiplied by mmm. No division by a possibly zero quantity occurs.
  • Constants. The only constant is the factor mmm, the number of processors.
  • Not formalized. The running time of SPT, the comparison with Lemma 8 and its remark "it is assumed that m<nm<nm<n" (not a hypothesis of Lemma 9), and the tightness Example 2.
  • Milestones 2 and 4 are stated for every listing σ\sigmaσ and every pair of listings; they are combinatorial facts about list schedules and sorted sums that are reusable beyond this mission. Proofs of any milestone, and a reusable packing lemma for non-overlapping intervals, are welcome.

Selected references

  • T. Gonzalez, S. Sahni, Flowshop and Jobshop Schedules: Complexity and Approximation, Operations Research 26(1), 36–52, 1978. https://doi.org/10.1287/opre.26.1.36
  • R. W. Conway, W. L. Maxwell, L. W. Miller, Theory of Scheduling, Addison-Wesley, 1967.
  • W. E. Smith, Various optimizers for single-stage production, Naval Research Logistics Quarterly 3, 59–66, 1956. https://doi.org/10.1002/nav.3800030106
  • K. Jansen, R. Solis-Oba, M. Sviridenko, Makespan Minimization in Job Shops: A Linear Time Approximation Scheme, SIAM J. Discrete Math. 16(2), 288–300, 2003 (source of the job-shop definition reused here). https://doi.org/10.1137/S0895480199363908
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Flowshop and Jobshop Schedules: Complexity and Approximation 4: The 3-Partition Job Shop on Two Machines Has a Schedule of Finish Time 2tB, Preemptive or Not, Iff a 3-Partition ExistsResearch Paper

Motivation

In a job shop, each job has an ordered list of tasks, and each task must use a specified machine. A production planner may be able to interrupt a task and resume it later, yet still be unable to find a short schedule efficiently. Gonzalez and Sahni studied how the ability to preempt changes the difficulty of minimizing the time at which all jobs are complete. Their 1978 paper establishes complexity results for flow shops and job shops and then examines approximation schedules when exact optimization is difficult (Gonzalez and Sahni, 1978).

This mission concerns the paper's two-machine construction from 3-PARTITION. It is the reduction used for their claim that finding an optimal preemptive or non-preemptive finish-time schedule remains NP-complete when input size is measured by the sum of task lengths. The formal target isolates the decision equivalence proved for the constructed instance. It gives solvers a precise schedule statement to establish without conflating the scheduling theorem with the separate complexity-class argument.

Setting

A 3-PARTITION input CCC comprises an integer t≥0t\ge0t≥0, a positive integer target BBB, and s=3ts=3ts=3t positive integer sizes a1,…,asa_1,\ldots,a_sa1​,…,as​. A valid input satisfies

∑i=1sai=tB,B/4<ai<B/2(1≤i≤s).\sum_{i=1}^{s}a_i=tB,\qquad B/4<a_i<B/2\quad(1\le i\le s).i=1∑s​ai​=tB,B/4<ai​<B/2(1≤i≤s).

It has a 3-partition when the sss indices can be assigned to ttt disjoint groups, each containing exactly three indices and having total size BBB. The strict bounds rule out groups of other cardinalities at that sum and are part of the source problem, not a convenience added for formalization. When t=0t=0t=0, the input has no sizes or groups and the empty grouping is a solution.

The paper builds a job shop JS(C)JS(C)JS(C) with machines P1,P2P_1,P_2P1​,P2​ and s+1s+1s+1 jobs. For 1≤i≤s1\le i\le s1≤i≤s, job iii has a task of length aia_iai​ on P1P_1P1​, followed by a task of length aia_iai​ on P2P_2P2​. The last job, s+1s+1s+1, has 2t2t2t tasks, each of length BBB. Its tasks alternate between the machines, starting on P2P_2P2​. This last order matters: reversing it would produce a different construction. The source gives the decision threshold τ=2tB\tau=2tBτ=2tB (Lemma 7, p. 44).

A non-preemptive schedule gives each task one continuous processing interval. A preemptive schedule may divide an operation into finitely many positive-length intervals. In either case, jobs obey their task order, distinct pieces on the same machine do not overlap, and processing starts at or after time zero. Write FT(S)FT(S)FT(S) for the time at which all tasks of schedule SSS have finished. The paper allows the same job to revisit a machine, as the final job does here.

Formalization targets

Lemma 7: the constructed instance

For every valid CCC, the main goal states both the paper's preemptive equivalence and the non-preemptive form established by its two directions:

∃S preemptive for JS(C):FT(S)≤2tB  ⟺  C has a 3-partition,∃S non-preemptive for JS(C):FT(S)≤2tB  ⟺  C has a 3-partition.\begin{aligned} \exists S\text{ preemptive for }JS(C): FT(S)\le2tB &\iff C\text{ has a 3-partition},\\ \exists S\text{ non-preemptive for }JS(C): FT(S)\le2tB &\iff C\text{ has a 3-partition}. \end{aligned}∃S preemptive for JS(C):FT(S)≤2tB∃S non-preemptive for JS(C):FT(S)≤2tB​⟺C has a 3-partition,⟺C has a 3-partition.​

The milestones follow the source's proof text: the schedule supplied by a solution in Lemma 7(a), the timing forced on the final job, the exclusion of short preemptive schedules in Lemma 7(b), and the preemptive equivalence stated in the proof. The final-job milestone describes the occupied time slots of the paper's Figure 4, independently of how many adjacent pieces a formal schedule uses.

Significance

The equivalence transfers the distinction between solvable and unsolvable 3-PARTITION inputs to an exact finish-time threshold in a job shop with only two machines. It also establishes the non-preemptive decision equivalence for this construction, because a solution yields a non-preemptive schedule while an unsolvable input rules out even the more permissive preemptive schedules. These facts supply the scheduling part of the paper's strong NP-completeness claim; membership in NP is handled separately by Lemma 6 (pp. 44–45).

The result is proved in the source paper. The work proposed here is a machine-checked development of its explicit reduction equivalence and the finite-piece scheduling model needed to state it. The milestone theorems expose statements that can be reused when formalizing other preemptive scheduling reductions. The shared job-shop instance and 3-PARTITION input are already published definitions; the particular reduction instance and preemptive schedule predicate are new local definitions.

Difficulty

Preemption makes a task's processing time distributable across intervals. Merely checking that each machine has enough total capacity does not characterize feasible schedules: every job's operations must still occur in order, and pieces of different jobs must avoid one another on each machine. The converse direction must constrain every preemptive schedule that meets the threshold, not only a schedule having the visual pattern in Figure 4. A schedule can also represent one uninterrupted interval as several adjacent pieces, so piece count alone cannot express the source's timing claim.

Formalization scope

Machines are Fin 2, with P1P_1P1​ numbered 000 and P2P_2P2​ numbered 111. Jobs are Fin (3t+1); the final job is index 3t3t3t. Its zero-based operation iii uses machine 111 for even iii and machine 000 for odd iii. The definition uses the published JobShopLTAS.Core.Instance for ordered operations, machines, real processing times, and non-preemptive feasibility. It uses the published ResourceScheduling.Chain.ThreePartition for t,B,ait,B,a_it,B,ai​, validity, and solutions. Those objects have the same conventions as the paper here. The local preemptive model uses finitely many positive-length pieces per task, whose durations sum exactly to that task's processing time. The threshold is the explicit real cast of 2tB2tB2tB; no real infimum or division by a possibly zero optimum appears.

The goal is specific to the instance the paper constructs from each valid input. A free job-shop instance, a freely chosen schedule, or an extra hypothesis that the input has a solution would erase the reduction's content. Validity carries B>0B>0B>0, total tBtBtB, and both strict bounds on each aia_iai​. The case t=0t=0t=0 remains in scope: the constructed job has no tasks and its finish-time threshold is zero. All existing tasks have positive length whenever the input is valid, so the published non-preemptive predicate's treatment of zero-length tasks does not affect this mission.

The paper states “3-Partition α\alphaα preemptive JOFT with m=2m=2m=2” and connects it to NP-completeness for preemptive and non-preemptive scheduling when complexity is measured by the sum of task lengths. The Lean theorem states the proof's two decision equivalences with constant 2tB2tB2tB and machine count 222. It does not formalize α\alphaα as a polynomial-time many-one reduction, the unary task-length measure, NP-completeness, or Lemma 6's NP-membership argument. Contributions that establish the reduction statements, the schedule construction, or general finite-piece scheduling facts are within scope.

Selected references

  • Teofilo Gonzalez and Sartaj Sahni, Flowshop and Jobshop Schedules: Complexity and Approximation, Operations Research 26(1), 1978, pp. 36–52. DOI: 10.1287/opre.26.1.36.
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Flowshop and Jobshop Schedules: Complexity and Approximation 2: The 3-Partition Flow Shop on Three Machines Has a Preemptive Schedule of Finish Time 2tB Iff a 3-Partition ExistsResearch Paper

Motivation

A flow shop is the simplest multi-stage production model: every job passes through the same machines in the same order, and the question is how to sequence the work so that everything is finished as early as possible. Gonzalez and Sahni's 1978 paper in Operations Research settled the computational complexity of the preemptive version of this problem, in which a task may be interrupted and resumed later on the same machine. For two machines the optimal finish time is computed by Johnson's rule, and preemption does not help. The paper shows that with three machines the preemptive problem is already NP-complete, and its second result, Theorem 2, strengthens this to NP-completeness in the strong sense: the problem stays hard even when the task times are written in unary.

Strong NP-completeness matters because it rules out a pseudo-polynomial algorithm, one whose running time is polynomial in the sum of the task lengths. The preemptive three-machine flow shop (F3∣pmtn∣Cmax⁡F3\mid pmtn\mid C_{\max}F3∣pmtn∣Cmax​ in later notation) is a standard entry in complexity classifications of scheduling problems, and Theorem 2 is the reference for its status. The paper notes that Garey and Johnson obtained a proof independently.

Timeline. Johnson (1954) solved the two-machine flow shop. Garey, Johnson and Sethi (1976) proved the non-preemptive three-machine flow shop NP-complete in the strong sense. Gonzalez and Sahni (1978) treated the preemptive case: Theorem 1 (ordinary NP-completeness via Partition) and Theorem 2 (strong NP-completeness via 3-Partition), the subject of this mission.

Setting

A flow shop has m≥1m\ge 1m≥1 processors P1,…,PmP_1,\dots,P_mP1​,…,Pm​ and a finite set of jobs. Task jjj of job iii runs on PjP_jPj​ and takes time tj,i≥0t_{j,i}\ge 0tj,i​≥0; zero times are allowed. For each job, task j≥2j\ge 2j≥2 can begin only after task j−1j-1j−1 has been completed, and schedules start at time 000.

A preemptive schedule is given, as in the paper's footnote on p. 40, by finitely many pieces (s,f)(s,f)(s,f) per task: the task is processed on its processor during [s,f)[s,f)[s,f). Pieces have 0≤s<f0\le s<f0≤s<f, pieces on the same processor do not overlap, the pieces of a task have total length tj,it_{j,i}tj,i​, and a piece of task j′j'j′ starts only after every piece of every earlier task j<j′j<j'j<j′ of the same job has ended. A task has no preemption when it is processed in at most one piece; a non-preemptive schedule has no preemption anywhere. The completion time of job iii is fi(S)f_i(S)fi​(S), the end of its last piece, and the finish time is FT(S)=max⁡ifi(S)FT(S)=\max_i f_i(S)FT(S)=maxi​fi​(S).

3-Partition. An instance C=(a1,…,as,B)C=(a_1,\dots,a_s,B)C=(a1​,…,as​,B), s=3ts=3ts=3t, consists of a positive integer BBB and nonnegative integers aia_iai​ with ∑iai=tB\sum_i a_i=tB∑i​ai​=tB and B/4<ai<B/2B/4<a_i<B/2B/4<ai​<B/2. It has a 3-partition if {1,…,s}\{1,\dots,s\}{1,…,s} splits into ttt disjoint sets L1,…,LtL_1,\dots,L_tL1​,…,Lt​, each of three elements with ∑j∈Lkaj=B\sum_{j\in L_k}a_j=B∑j∈Lk​​aj​=B.

The instance FS. From CCC the proof of Lemma 4 builds a three-processor flow shop with s+t+2s+t+2s+t+2 jobs:

  • jobs 1,…,s1,\dots,s1,…,s: times (ai, 0, ai)(a_i,\,0,\,a_i)(ai​,0,ai​) on (P1,P2,P3)(P_1,P_2,P_3)(P1​,P2​,P3​);
  • job s+1s+1s+1: (0, 2B, B)(0,\,2B,\,B)(0,2B,B);
  • jobs s+i+1s+i+1s+i+1, 1≤i≤t−21\le i\le t-21≤i≤t−2: (B, 2B, B)(B,\,2B,\,B)(B,2B,B);
  • job s+ts+ts+t: (B, 2B, 0)(B,\,2B,\,0)(B,2B,0);
  • job s+t+1s+t+1s+t+1: (0, 0, B)(0,\,0,\,B)(0,0,B);
  • job s+t+2s+t+2s+t+2: (B, 0, 0)(B,\,0,\,0)(B,0,0).

Each processor carries total work exactly 2tB2tB2tB, and the threshold is τ=2tB\tau=2tBτ=2tB.

Formalization targets

Goal: Theorem 2, as Lemma 4's equivalence

For every 3-Partition instance CCC with t≥2t\ge 2t≥2,

∃ S preemptive schedule of FS: FT(S)≤2tB⟺C has a 3-partition,\exists\,S \text{ preemptive schedule of } FS:\ FT(S)\le 2tB \quad\Longleftrightarrow\quad C \text{ has a 3-partition},∃S preemptive schedule of FS: FT(S)≤2tB⟺C has a 3-partition,

and every task time of FSFSFS satisfies tj,i≤2Bt_{j,i}\le 2Btj,i​≤2B. The second clause records that the instance's numbers are bounded by those of CCC, the reduction's contribution to "the problem size being measured as the sum of the length of the tasks".

Milestones

  1. Lemma 3 (p. 40). For a flow shop with m≥1m\ge 1m≥1 processors, every preemptive schedule SSS can be replaced by a schedule S′S'S′ with no preemptions on P1P_1P1​ and on PmP_mPm​ and FT(S′)=FT(S)FT(S')=FT(S)FT(S′)=FT(S).
  2. Lemma 4(a) (p. 41). If CCC has a 3-partition, FSFSFS has a non-preemptive schedule with FT≤2tBFT\le 2tBFT≤2tB.
  3. First step of Lemma 4(b) (pp. 41–42). In every preemptive schedule of FSFSFS with FT≤2tBFT\le 2tBFT≤2tB, the jobs among 1,…,s1,\dots,s1,…,s whose P1P_1P1​ task is completed by time 2B2B2B have ∑t1,i=B\sum t_{1,i}=B∑t1,i​=B.
  4. Lemma 4(b) (pp. 41–42). If CCC has no 3-partition, every preemptive schedule of FSFSFS has FT>2tBFT>2tBFT>2tB.

Significance

The result. Theorem 2 places the preemptive three-machine flow shop among the strongly NP-hard scheduling problems, so no algorithm polynomial in the number of jobs and the total processing time exists unless P = NP. It also shows that allowing preemption, which makes several single-stage problems (such as P∣pmtn∣Cmax⁡P\mid pmtn\mid C_{\max}P∣pmtn∣Cmax​) easy, does not help for three-stage flow shops. Lemma 3, that preemptions on the first and last machines can be removed without changing the finish time, holds for any number of machines and is a reusable structural fact about flow-shop schedules.

Formalizing it. The result is proved on paper, and its proof of the "only if" direction is an informal busy-processor argument repeated window by window. As far as is known, no part of it has a machine-checked proof. A formal development has to make the interval accounting rigorous: why each processor is busy throughout [0,2tB][0,2tB][0,2tB], why exactly BBB units of element jobs end on P1P_1P1​ by 2B2B2B, and how the argument restarts on [2B,2tB][2B,2tB][2B,2tB]. Lemma 3 needs an exchange argument on piece representations. Both are new for the platform.

Difficulty

The direction "3-partition ⇒\Rightarrow⇒ schedule" is a direct construction. The hard direction is the converse. Preemption lets any task be split across many intervals, so the usual non-preemptive arguments, which reason about the order in which whole tasks run, do not apply directly. The paper's argument that the window [0,2B][0,2B][0,2B] must contain exactly three element jobs of total BBB on P1P_1P1​ depends on every processor being continuously busy, and turning "otherwise there is idle time" into a precise statement requires a careful account of which jobs are available to each processor at each moment. The induction on windows is then not a literal restriction of the schedule to a smaller instance, because pieces may cross the boundary at 2B2B2B.

Formalization scope

  • Times are real numbers; the 3-Partition data are natural numbers cast to R\mathbb RR. Processors are Fin m with P1=P_1=P1​= 0; in FSFSFS, P1,P2,P3P_1,P_2,P_3P1​,P2​,P3​ are 0, 1, 2.
  • The 3-Partition instance is the published ResourceScheduling.Chain.ThreePartition (0-based indices, t parts, b =B=B=B, Valid, HasSolution), whose definition is the paper's p. 37 problem.
  • Preemptive schedules are finite sets of pieces per task; precedence is required between every pair of tasks j<j′j<j'j<j′ of a job, so a zero task occupies no time and does not break the job's order. Finish time is a maximum over jobs with baseline 000, never an unattained supremum.
  • "No preemptions on PjP_jPj​" means at most one piece per task on PjP_jPj​.
  • Not formalized: "NP-complete", the reduction symbol ∝\propto∝, "the problem size being measured as the sum of the length of the tasks", and Lemma 2 (membership in NP). The goal states the equivalence Lemma 4 proves for the constructed instance, plus the bound tj,i≤2Bt_{j,i}\le 2Btj,i​≤2B on its task times.
  • Added hypothesis: t≥2t\ge 2t≥2. For t<2t<2t<2 the paper's job indices s+1s+1s+1 and s+ts+ts+t coincide with conflicting times, so the construction is undefined there.
  • The first step of Lemma 4(b) is read as: the sum of t1,it_{1,i}t1,i​ over element jobs whose P1P_1P1​ task has completed by time 2B2B2B equals BBB.
  • A trivializing formalization is ruled out: the schedules quantified over are all preemptive schedules of the specific instance FS(C)FS(C)FS(C), the threshold 2tB2tB2tB is explicit, and the validity conditions B/4<ai<B/2B/4<a_i<B/2B/4<ai​<B/2 are kept, so the schedule side cannot match the weaker "partition into ttt groups of sum BBB".

Contributions welcome: the interval-accounting lemmas (busy processors, work available before a time), Lemma 3's exchange argument, and the schedule of Figure 2.

Selected references

  • T. Gonzalez, S. Sahni, Flowshop and Jobshop Schedules: Complexity and Approximation, Operations Research 26(1):36–52, 1978. https://doi.org/10.1287/opre.26.1.36
  • M. R. Garey, D. S. Johnson, R. Sethi, The Complexity of Flowshop and Jobshop Scheduling, Mathematics of Operations Research 1(2):117–129, 1976. https://doi.org/10.1287/moor.1.2.117
  • S. M. Johnson, Optimal two- and three-stage production schedules with setup times included, Naval Research Logistics Quarterly 1(1):61–68, 1954. https://doi.org/10.1002/nav.3800010110
  • M. R. Garey, D. S. Johnson, "Strong" NP-Completeness Results: Motivation, Examples, and Implications, Journal of the ACM 25(3):499–508, 1978. https://doi.org/10.1145/322077.322090
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Flowshop and Jobshop Schedules: Complexity and Approximation 1: The Partition Flow Shop on Three Machines Has a Schedule of Finish Time 2T, Preemptive or Not, Iff a Partition ExistsResearch Paper

Motivation

Shop scheduling asks how to process a set of jobs, each a sequence of tasks on prescribed machines, so that all work is completed as early as possible. The flow shop, in which every job visits the machines in the same order, is the most studied of these models in operations research. For two machines, Johnson's rule (1954) computes a schedule of minimum finish time in O(nlog⁡n)O(n\log n)O(nlogn) time, and for two machines allowing interruptions (preemption) does not help. Garey, Johnson and Sethi (Math. Oper. Res. 1976) and Brucker, Lenstra and Rinnooy Kan (Mathematisch Centrum report BW 43/75, 1975; Ann. Discrete Math. 1977) showed that the non-preemptive problem becomes NP-complete with three machines.

Gonzalez and Sahni (Operations Research, 1978) extended this picture to preemptive schedules. Their Theorem 1 shows that minimizing the finish time of a three-machine flow shop is NP-complete with or without preemption, even when every job has at most two tasks of nonzero length. When every job has only one task the problem is trivial, so this is the simplest NP-complete case of the flow-shop finish-time problem. The result explains why no exact polynomial algorithm is expected and motivates the approximation results in the second half of the same paper.

Timeline:

  • 1954: Johnson, two-machine flow shop, optimal non-preemptive rule.
  • 1975: Brucker, Lenstra, Rinnooy Kan: non-preemptive three-machine flow shop NP-complete, including the case of two nonzero tasks per job (the paper's Corollary 1, which it says "is also obtained in [1]").
  • 1976: Garey, Johnson, Sethi: non-preemptive three-machine flow shop NP-complete even when the problem size is the sum of the task lengths.
  • 1978: Gonzalez, Sahni: the preemptive three-machine flow shop is NP-complete, by a reduction from PARTITION that also covers the non-preemptive case with two nonzero tasks per job.

Setting

A flow shop has m≥1m\ge1m≥1 processors P1,…,PmP_1,\dots,P_mP1​,…,Pm​ and a finite set of jobs. Task jjj of job iii must be processed on PjP_jPj​ for tj,i≥0t_{j,i}\ge 0tj,i​≥0 time units, and it can start only after task j−1j-1j−1 of the same job is completed. Task times may be zero. The schedule starts at time 000.

A preemptive schedule gives every task a finite set of pieces (s,f)(s,f)(s,f) with 0≤s<f0\le s<f0≤s<f: the task is processed on its processor during [s,f)[s,f)[s,f). Pieces on the same processor do not overlap, the pieces of a task have total length tj,it_{j,i}tj,i​, and every piece of a task starts after every piece of the earlier tasks of the same job has ended. A non-preemptive schedule processes every task in at most one piece. For a schedule SSS, fi(S)f_i(S)fi​(S) is the time at which job iii is completed and the finish time is FT(S)=max⁡ifi(S)\mathrm{FT}(S)=\max_i f_i(S)FT(S)=maxi​fi​(S).

PARTITION. A multiset S={a1,…,an}S=\{a_1,\dots,a_n\}S={a1​,…,an​} of nonnegative integers has a partition if some set uuu of indices satisfies ∑i∈uai=12∑i=1nai\sum_{i\in u}a_i=\tfrac12\sum_{i=1}^n a_i∑i∈u​ai​=21​∑i=1n​ai​.

The instance FS. With T=∑iaiT=\sum_i a_iT=∑i​ai​, the flow shop FS\mathrm{FS}FS has m=3m=3m=3 processors and n+2n+2n+2 jobs:

t⋅,i=(ai,0,ai) (1≤i≤n),t⋅,n+1=(T/2, T, 0),t⋅,n+2=(0, T, T/2),t_{\cdot,i}=(a_i,0,a_i)\ (1\le i\le n),\qquad t_{\cdot,n+1}=(T/2,\,T,\,0),\qquad t_{\cdot,n+2}=(0,\,T,\,T/2),t⋅,i​=(ai​,0,ai​) (1≤i≤n),t⋅,n+1​=(T/2,T,0),t⋅,n+2​=(0,T,T/2),

and the threshold is τ=2T\tau=2Tτ=2T. Every job has at most two nonzero tasks.

Formalization targets

Goal: Theorem 1 (p. 38)

For every nnn and every a1,…,an∈Na_1,\dots,a_n\in\mathbb Na1​,…,an​∈N:

FS has at most two nonzero tasks per job,\mathrm{FS}\ \text{has at most two nonzero tasks per job},FS has at most two nonzero tasks per job, (∃ S′ preemptive, FT(S′)≤2T)  ⟺  S has a partition,\big(\exists\,S'\ \text{preemptive},\ \mathrm{FT}(S')\le 2T\big)\iff S\ \text{has a partition},(∃S′ preemptive, FT(S′)≤2T)⟺S has a partition, (∃ S′ non-preemptive, FT(S′)≤2T)  ⟺  S has a partition.\big(\exists\,S'\ \text{non-preemptive},\ \mathrm{FT}(S')\le 2T\big)\iff S\ \text{has a partition}.(∃S′ non-preemptive, FT(S′)≤2T)⟺S has a partition.

Milestones

  • Lemma 1(a) (p. 39): a partition of SSS yields a non-preemptive schedule of FS with FT=2T\mathrm{FT}=2TFT=2T (Figure 1).
  • Observations (i)–(ii) (p. 39): in any preemptive schedule with FT≤2T\mathrm{FT}\le 2TFT≤2T, task t1,n+1t_{1,n+1}t1,n+1​ ends by TTT and task t3,n+2t_{3,n+2}t3,n+2​ starts no earlier than TTT.
  • Lemma 1(b) (p. 39): if SSS has no partition, every preemptive schedule of FS has FT>2T\mathrm{FT}>2TFT>2T.
  • Lemma 1 (p. 38): the preemptive equivalence.
  • Corollary 1 (p. 39): the non-preemptive equivalence.

Significance

The result. The equivalence shows that a polynomial-time algorithm deciding whether a three-machine flow shop (with two nonzero tasks per job) has a schedule of finish time at most τ\tauτ, preemptive or not, would decide PARTITION in polynomial time. With membership in NP (the paper's Lemma 2) it gives NP-completeness of both versions. It also marks the boundary with the polynomial cases: m=2m=2m=2 (Johnson's rule), and jobs with a single task.

Formalizing it. The result is proved in the paper; nothing here is open. To our knowledge no machine-checked proof of this reduction exists, and the platform has no formal model of preemptive shop schedules. The mission produces one: a piece-based preemptive schedule with zero task times, which subsumes the non-preemptive case, together with the first formal NP-hardness reduction for preemptive flow shops on the platform.

Difficulty

Direction (a) asks for one explicit schedule. Direction (b) must rule out every preemptive schedule of FS, and preemption is exactly what makes this delicate: a task may be split into any finite number of pieces placed anywhere on its processor, so the naive intuition that a failed partition leaves "unusable gaps" is no longer evident, since pieces of different jobs can be interleaved to fill gaps. A correct argument must therefore be an accounting of processor time over the pieces, valid for any number and position of pieces, and it must treat zero-length tasks, which occupy no processor time but still delay the next task of their job. Finally, the non-preemptive equivalence is not a formal consequence of the preemptive one alone: it uses that (a) produces a non-preemptive schedule and that (b) holds for all preemptive schedules, which include the non-preemptive ones.

Formalization scope

  • Processors are Fin 3 (P1,P2,P3P_1,P_2,P_3P1​,P2​,P3​ = 0, 1, 2), jobs of FS are Fin n ⊕ Fin 2 (Sum.inr 0 is job n+1n+1n+1, Sum.inr 1 is job n+2n+2n+2). Task times are real numbers; the PARTITION data are a : Fin n → ℕ, and "has a partition" is 2∑i∈uai=∑iai2\sum_{i\in u}a_i=\sum_i a_i2∑i∈u​ai​=∑i​ai​ in N\mathbb NN.
  • A schedule is a Finset of pieces (s,f)(s,f)(s,f) per task with the four conditions of the Setting. Non-preemptive means at most one piece per task. Zero tasks have no pieces and occupy no processor time. fi(S)f_i(S)fi​(S) and FT(S)\mathrm{FT}(S)FT(S) are maxima with baseline 000.
  • Explicit constants: T=∑iaiT=\sum_i a_iT=∑i​ai​, the times T/2T/2T/2 and TTT of jobs n+1,n+2n+1,n+2n+1,n+2, and the threshold τ=2T\tau=2Tτ=2T, as on p. 39. Lemma 1(a) states finish time exactly 2T2T2T, as the paper does.
  • Not formalized: "NP-complete", the reducibility relation "α\alphaα", membership in NP (Lemma 2, p. 40) and the polynomial computability of FS from SSS. The paper states Theorem 1, Lemma 1 and Corollary 1 as complexity claims. Their proofs establish the equivalences above for the constructed instance, and those equivalences are what is formalized.
  • No trivializing reading is admitted: the statements concern the specific instance FS built from aaa, not an arbitrary flow shop, and the schedule model rules out overlapping pieces and out-of-order tasks while still allowing zero tasks. A sorry-free check confirms that Figure 1 is a valid non-preemptive schedule of FS for S={1,1}S=\{1,1\}S={1,1}.
  • Reusable beyond this mission: the flow-shop model with preemptive pieces, and PARTITION. Proofs of the milestones and a proof of the goal from them are welcome.

Selected references

  • T. Gonzalez, S. Sahni, Flowshop and Jobshop Schedules: Complexity and Approximation, Operations Research 26(1), 36–52, 1978. https://doi.org/10.1287/opre.26.1.36
  • M. R. Garey, D. S. Johnson, R. Sethi, The Complexity of Flowshop and Jobshop Scheduling, Mathematics of Operations Research 1(2), 117–129, 1976. https://doi.org/10.1287/moor.1.2.117
  • S. M. Johnson, Optimal two- and three-stage production schedules with setup times included, Naval Research Logistics Quarterly 1(1), 61–68, 1954. https://doi.org/10.1002/nav.3800010110
  • P. Brucker, J. K. Lenstra, A. H. G. Rinnooy Kan, Complexity of Machine Scheduling Problems, Mathematisch Centrum report BW 43/75, 1975; Annals of Discrete Mathematics 1, 343–362, 1977. https://doi.org/10.1016/S0167-5060(08)70743-X
  • R. M. Karp, Reducibility among Combinatorial Problems, in Complexity of Computer Computations, Plenum, 85–103, 1972. https://doi.org/10.1007/978-1-4684-2001-2_9
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Dynamic ProgrammingLinear OptimizationMarkov Chain+1·Captain: mikedeng1

Linear Programming and Finite Markovian Control Problems VII: Bias Optimality — Four Equivalent Characterizations of a Bias Optimal Pure Stationary PolicyTextbook

Why bias optimality matters

A finite Markov decision problem can have several policies with the same long-run average reward even though they behave differently before reaching a recurrent class. Average reward discards a finite initial advantage: dividing total reward by the horizon sends any fixed transient contribution to zero. In applications where the process runs for a long but unspecified time, this leaves a genuine choice among average-optimal policies. Bias optimality uses discounted values as the discount factor approaches one to resolve that choice. Kallenberg develops the criterion in Chapter 5 of Linear Programming and Finite Markovian Control Problems, following Blackwell's earlier notion of a nearly optimal policy.

The chapter asks for a way to recognize a bias optimal policy without comparing it directly with every possible sequence of history-dependent, randomized decisions. Its central characterization, Theorem 5.2.2, gives three tests involving only pure stationary policies. That change in the comparison class matters: there are finitely many pure stationary rules, while a general policy may depend on an unbounded history. The theorem also explains why average optimality alone leaves policies untied only at the gain level, while their bias vectors can differ. The two-state example in Figure 5.2.1 exhibits an average-optimal stationary rule that is not bias optimal (Kallenberg, pp. 163–164).

The finite decision model

The state space is E={1,…,N}E=\{1,\ldots,N\}E={1,…,N} with N>0N>0N>0. Each state iii has its own finite, nonempty set A(i)A(i)A(i) of admissible actions. Choosing a∈A(i)a\in A(i)a∈A(i) earns the real reward riar_{ia}ria​ immediately and moves to state jjj with probability piajp_{iaj}piaj​. The rows are stochastic: piaj≥0p_{iaj}\ge0piaj​≥0 and ∑jpiaj=1\sum_jp_{iaj}=1∑j​piaj​=1. This is the chapter's standing assumption, stated at the opening of §5.2 (Kallenberg, p. 162).

A policy R∈CR\in CR∈C assigns a probability distribution on the actions available at the current state after every finite history. Thus CCC includes randomized and history-dependent decisions. A pure stationary policy f∞∈CDf^\infty\in C_Df∞∈CD​ instead fixes one action f(i)∈A(i)f(i)\in A(i)f(i)∈A(i) for each state and uses it in every period. Time starts at t=1t=1t=1. From initial state iii, let vit(R)v_i^t(R)vit​(R) be expected reward in period ttt, and let viα(R)=∑t≥1αt−1vit(R)v_i^\alpha(R)=\sum_{t\ge1}\alpha^{t-1}v_i^t(R)viα​(R)=∑t≥1​αt−1vit​(R) be expected discounted reward for 0≤α<10\le\alpha<10≤α<1. The optimal discounted value viαv_i^\alphaviα​ is the supremum of viα(R)v_i^\alpha(R)viα​(R) over all R∈CR\in CR∈C, taken separately for each state.

The average reward is ϕi(R)=lim inf⁡T→∞T−1∑t=1Tvit(R)\phi_i(R)=\liminf_{T\to\infty}T^{-1}\sum_{t=1}^T v_i^t(R)ϕi​(R)=liminfT→∞​T−1∑t=1T​vit​(R); its optimal value ϕi\phi_iϕi​ is again the statewise supremum over CCC. The corresponding limsup is ϕ^i(R)\hat\phi_i(R)ϕ^​i​(R). A policy is average optimal when ϕi(R)=ϕi\phi_i(R)=\phi_iϕi​(R)=ϕi​ at every state. It is bias optimal when viα(R)−viα→0v_i^\alpha(R)-v_i^\alpha\to0viα​(R)−viα​→0 as α↑1\alpha\uparrow1α↑1 at every state (Kallenberg, pp. 21–23 and 161).

For a pure stationary fff, write P(f)ij=pi,f(i),jP(f)_{ij}=p_{i,f(i),j}P(f)ij​=pi,f(i),j​ and r(f)i=ri,f(i)r(f)_i=r_{i,f(i)}r(f)i​=ri,f(i)​. The stationary matrix P∗(f)P^*(f)P∗(f) is the Cesàro limit of the powers of P(f)P(f)P(f). The deviation matrix and bias vector are

D(f)=(I−P(f)+P∗(f))−1−P∗(f),u(f∞)=D(f)r(f).D(f)=\bigl(I-P(f)+P^*(f)\bigr)^{-1}-P^*(f), \qquad u(f^\infty)=D(f)r(f).D(f)=(I−P(f)+P∗(f))−1−P∗(f),u(f∞)=D(f)r(f).

The gain of f∞f^\inftyf∞ is ϕ(f∞)=P∗(f)r(f)\phi(f^\infty)=P^*(f)r(f)ϕ(f∞)=P∗(f)r(f). These are the book's Definitions 2.4.2 and notation (2.5.5), and their finite stochastic-matrix properties are recorded in Theorem 2.4.1 (Kallenberg, pp. 29 and 33–34).

Formalization targets

The goal is Theorem 5.2.2, for each pure stationary f∗∞f_*^\inftyf∗∞​. It asserts that the following are equivalent: (i) f∗∞f_*^\inftyf∗∞​ is bias optimal against the value over all CCC; (ii) for every f∞∈CDf^\infty\in C_Df∞∈CD​, the vector limit lim⁡α↑1(vα(f∗∞)−vα(f∞))\lim_{\alpha\uparrow1}\bigl(v^\alpha(f_*^\infty)-v^\alpha(f^\infty)\bigr)limα↑1​(vα(f∗∞​)−vα(f∞)) exists and is nonnegative; (iii) f∗∞f_*^\inftyf∗∞​ is average optimal and, at every state iii, its bias is maximal among pure stationary policies optimal at that state; and (iv) for every f∞∈CDf^\infty\in C_Df∞∈CD​, the vector limit

lim⁡T→∞1T∑t=1T∑s=1t(vs(f∗∞)−vs(f∞))≥0\lim_{T\to\infty}\frac1T\sum_{t=1}^T\sum_{s=1}^t \bigl(v^s(f_*^\infty)-v^s(f^\infty)\bigr)\ge0T→∞lim​T1​t=1∑T​s=1∑t​(vs(f∗∞​)−vs(f∞))≥0

exists. These limits are read in the extended reals: a strictly positive gain gap makes a difference tend to +∞+\infty+∞, which satisfies the displayed inequality. The period rewards vsv^svs in the last formula are not discounted values. The printed clause (iii) has a typographical error in the set following max; its proof on p. 164 supplies the intended restriction ϕi(f∞)=ϕi\phi_i(f^\infty)=\phi_iϕi​(f∞)=ϕi​. The formal statement uses that reading and records it explicitly.

Two earlier chapter results are milestones. Lemma 5.2.1 bounds the Abel limsup (1−α)viα(R)(1-\alpha)v_i^\alpha(R)(1−α)viα​(R) by the upper average reward ϕ^i(R)\hat\phi_i(R)ϕ^​i​(R) for any policy. Theorem 5.2.1 says that bias optimality implies average optimality, with a two-state example showing that average optimality does not imply bias optimality. The later §5.3 linear-programming construction is a separate route to finding a bias optimal policy and is outside this mission's statement layer.

What the characterization gives

The equivalence allows bias optimality to be checked through discounted differences, bias vectors, or accumulated period rewards. Its bias-vector clause identifies the relevant comparison set: a policy can be optimal in average reward at one state without being optimal at every state, and that statewise distinction cannot be discarded. It also makes the bias criterion finer than average optimality, as the two-state counterexample shows (Kallenberg, pp. 163–165).

The source proves these results; the mission asks for machine-checked proofs of the stated equivalence and its two supporting results. The finite history and reward definitions can support later formalizations of Kallenberg's average-reward and constrained models. The matrix definitions can also be reused for results about finite Markov chains, including the Cesàro limit and deviation matrix. No machine-checked proof of this specific four-way characterization is claimed here.

Where the difficulty lies

Comparing average rewards alone loses the transient reward information that bias is designed to retain. The discounted comparison has an apparent divergence of order (1−α)−1(1-\alpha)^{-1}(1−α)−1 whenever two policies have different gains; the finite residual becomes meaningful only after those gain terms are separated. The accumulated-reward test also takes a second Cesàro average, so a direct identification with ordinary average reward would miss its bias contribution. The theorem has to preserve both the gain ordering and the finite bias remainder across these limiting regimes (Kallenberg, pp. 164–165).

For a general policy, the discounted value is defined by infinitely many history-dependent choices, whereas P∗P^*P∗ and DDD apply to one fixed stationary matrix. This is the obstacle in passing from clause (i), which compares with all CCC, to the tests against CDC_DCD​. Treating the optimal discounted value as a supremum over stationary policies by definition would erase that obstacle and change the theorem.

Formalization scope

Lean uses Fin N with N>0N>0N>0 for states and a finite action type with nonempty state-dependent admissible subsets. Rewards are real, transition rows are stochastic, and policies are distributions on actions after complete finite histories. The finite sum of products defining each history's probability supplies vit(R)v_i^t(R)vit​(R); no measure-theoretic path space is needed. Discounted rewards use a real infinite sum for α\alphaα near one from below. Average rewards are real liminf and limsup values of bounded sequences. Optimal values are statewise suprema over all admissible policies, not over just CDC_DCD​.

The matrix P∗P^*P∗ is a Cesàro limit and DDD uses the book's inverse formula. Their convergence and nonsingularity for finite stochastic matrices are mathematical obligations, not added assumptions of the goal. The T=0T=0T=0 value of a totalized average and the value of a discounted sum outside its summable range do not enter the limits. Contributions to the finite-chain limit facts, boundedness of reward criteria, the Abel/Cesàro comparison, and the four equivalences are all within scope. A formalization that restricts the optimum in bias optimality to pure stationary policies would not be this theorem.

Selected references

  • L. C. M. Kallenberg, Linear Programming and Finite Markovian Control Problems, Mathematical Centre Tracts 148, Mathematisch Centrum, Amsterdam, 1983. Publisher repository.
  • David Blackwell, “Discrete Dynamic Programming,” Annals of Mathematical Statistics 33 (1962), 719–726. DOI.
8 thms1 active userReviewed
Dynamic ProgrammingLinear OptimizationOptimization·Captain: mikedeng1

Linear Programming and Finite Markovian Control Problems IV: Positive Dynamic Programming — an Extreme Optimal Dual Solution Yields a Pure Stationary Optimal PolicyTextbook

Motivation

A finite Markov decision system asks how to choose actions while a changing state controls which actions are available. The choice can affect both the reward now and the distribution of future states. In a total-reward problem, the process can continue indefinitely; if transitions are substochastic, it can also terminate. Kallenberg's treatment of positive dynamic programming connects this control problem to a finite linear program and uses an optimal flow to identify a policy that is optimal from every initial state Kallenberg, Chapter 3.

The connection matters when an optimizer is easier to compute as state-action frequencies than as an infinite sequence of decisions. A policy may depend on the entire observed history and may randomize at every decision. The theorem under study says that, when the dual linear program has an extreme optimal solution, its positive coordinates identify a pure stationary policy as good as every such general policy. This assertion includes models in which the selected policy is not transient Kallenberg, Theorem 3.5.2 and Example 3.5.1.

Setting

Let E={1,…,N}E=\{1,\ldots,N\}E={1,…,N} be a nonempty finite state space. At state iii, one chooses an action from the nonempty finite set A(i)A(i)A(i). Action a∈A(i)a\in A(i)a∈A(i) earns a real reward riar_{ia}ria​ and moves to state jjj with probability piaj≥0p_{iaj}\ge0piaj​≥0. The sum ∑jpiaj\sum_jp_{iaj}∑j​piaj​ is at most one; any missing probability represents termination. Positive dynamic programming imposes ria≥0r_{ia}\ge0ria​≥0 for every admissible state-action pair Kallenberg, Section 2.2 and Assumption 3.5.1.

A policy R=(π1,π2,…)R=(\pi^1,\pi^2,\ldots)R=(π1,π2,…) assigns an action distribution after each finite history. The history records the initial state and all states and actions observed so far. Write CCC for this full policy class. A pure stationary policy f∞f^\inftyf∞ always chooses one fixed admissible action f(i)f(i)f(i) whenever it visits iii.

Starting from state iii, let vi(R)v_i(R)vi​(R) be the expected total reward of RRR, and let vi=sup⁡R∈Cvi(R)v_i=\sup_{R\in C}v_i(R)vi​=supR∈C​vi​(R). Because the rewards are nonnegative, the finite-horizon expected totals increase to vi(R)v_i(R)vi​(R); both policy and optimal values may be +∞+\infty+∞. A real vector w=(wi)w=(w_i)w=(wi​) is TMD superharmonic if wi≥ria+∑jpiajwjw_i\ge r_{ia}+\sum_jp_{iaj}w_jwi​≥ria​+∑j​piaj​wj​ for each iii and a∈A(i)a\in A(i)a∈A(i) Kallenberg, Definition 3.3.1.

For given weights βj>0\beta_j>0βj​>0, equation (3.5.2) is a linear program over nonnegative state-action flows xiax_{ia}xia​. It maximizes ∑i,ariaxia\sum_{i,a}r_{ia}x_{ia}∑i,a​ria​xia​ subject to

∑a∈A(j)xja−∑i∈E∑a∈A(i)piajxia≤βj(j∈E).\sum_{a\in A(j)}x_{ja}-\sum_{i\in E}\sum_{a\in A(i)}p_{iaj}x_{ia}\le\beta_j\qquad(j\in E).a∈A(j)∑​xja​−i∈E∑​a∈A(i)∑​piaj​xia​≤βj​(j∈E).

The set Ex={i:∑a∈A(i)xia>0}E_x=\{i:\sum_{a\in A(i)}x_{ia}>0\}Ex​={i:∑a∈A(i)​xia​>0} records the states with positive flow Kallenberg, Notation 3.1.1 and equation (3.5.2).

Formalization targets

Superharmonic value

Theorem 3.5.1 identifies vvv as the smallest nonnegative TMD-superharmonic vector. The original superharmonic definition ranges over real vectors; the value must also be allowed to take +∞+\infty+∞. The formal statement gives both the extended superharmonic inequalities for vvv and its componentwise bound by every nonnegative real superharmonic vector:

0≤vi,vi≥ria+∑jpiajvj,vi≤wifor every nonnegative real superharmonic w.0\le v_i,\qquad v_i\ge r_{ia}+\sum_jp_{iaj}v_j,\qquad v_i\le w_i\quad\text{for every nonnegative real superharmonic }w.0≤vi​,vi​≥ria​+j∑​piaj​vj​,vi​≤wi​for every nonnegative real superharmonic w.

This is the numbered milestone of the mission Kallenberg, Theorem 3.5.1.

Policy from an extreme optimal dual flow

The main goal is Theorem 3.5.2. Suppose x∗x^*x∗ is an extreme point of the feasible region of (3.5.2) and has objective at least as large as every feasible flow. Then every pure stationary rule fff choosing xi,f(i)∗>0x^*_{i,f(i)}>0xi,f(i)∗​>0 on Ex∗E_{x^*}Ex∗​, with arbitrary admissible actions outside Ex∗E_{x^*}Ex∗​, is total optimal:

vi(f∞)=vi=sup⁡R∈Cvi(R)(i∈E).v_i(f^\infty)=v_i=\sup_{R\in C}v_i(R)\qquad(i\in E).vi​(f∞)=vi​=R∈Csup​vi​(R)(i∈E).

The goal records the finite-optimum case by requiring an actual optimal solution x∗x^*x∗; it makes no extra assumption that every policy is transient Kallenberg, Theorem 3.5.2.

Significance

The theorem converts an extreme optimal state-action flow into one decision rule that is optimal simultaneously for all starting states. Since CCC includes policies that change with time, use full histories or randomize, it establishes that those freedoms cannot improve the total-reward value when the stated dual optimizer exists. The positive model also reveals when extended values matter: outside the finite-optimum case, the dual objective may be unbounded and an optimal total reward may be infinite Kallenberg, pp. 78–83.

A formal development of this result must keep the policy class and the linear program in the same model. It provides reusable definitions of a finite substochastic control system, finite-history probabilities, extended total rewards, state-action flows and superharmonic vectors. The result is proved in the book; the mission asks for a machine-checked proof of its statement, not for a new optimization theorem. The draft statements are open goals until such proofs are supplied.

Difficulty

A finite linear program has finitely many variables, while the value compares a flow-derived rule with every infinite, history dependent policy. A comparison restricted to stationary or transient policies would miss the main assertion. The proof also has to treat zero-flow states: the dual solution does not prescribe an action there, yet the theorem permits every admissible choice. Finally, the superharmonic milestone permits an infinite value, while the linear program itself uses real coordinates. A development must connect those representations without turning an infinite reward into an arbitrary default real value.

Formalization scope

Lean represents states as Fin N with N>0N>0N>0, and each A(i)A(i)A(i) as a nonempty finite subset of a finite action type. The transition rows sum to at most one. Histories carry past states and actions, and the first action has internal index zero, corresponding to the book's time one. Policy probabilities are nonnegative, vanish outside A(i)A(i)A(i) and sum to one. The expected reward of a finite horizon is a finite sum over histories; nonnegative total reward is its supremum in EReal. The value takes the supremum over the full policy type, not over a restricted policy subclass.

Dual flow coordinates exist only for admissible state-action pairs. Extreme points are taken in the feasible set of (3.5.2) in those coordinates, and its balance constraints retain the source's ≤βj\le\beta_j≤βj​ direction. The goal quantifies over every admissible pure rule satisfying the positive-flow condition. These choices exclude the vacuous shortcut of defining the value using only pure stationary policies or assuming optimality as a hypothesis.

The mathematical infrastructure includes finite policy histories, state-action occupancy probabilities, extended nonnegative reward limits, the real flow polyhedron and its extreme points. The history and flow definitions can support later work on other finite decision models. Contributions to those foundational results and to the full comparison with history dependent policies are within the mission's scope. The negative-reward regime of Section 3.6, whose claims involve average optimality and recurrent classes in an extended stochastic model, requires its own representation and is outside this positive-reward target.

Selected references

  • L. C. M. Kallenberg, Linear Programming and Finite Markovian Control Problems, Mathematical Centre Tracts 148, Mathematisch Centrum, Amsterdam, 1983. Publisher repository.
6 thms1 active userReviewed
Discrete GeometryGroup Theory·Captain: mikedeng1

Some Polyhedra Related to Combinatorial Problems III: Faces of P(H, ψg0) Lift Through a Homomorphism ψ of G onto H to Faces of P(G, g0)Research Paper

Motivation

Every pure integer program, once its linear programming relaxation has been solved, can be relaxed further to a problem over a finite Abelian group: the nonbasic variables must satisfy a single congruence in the group of the optimal basis. R. E. Gomory introduced this relaxation in 1965 and studied the convex hull of its solutions in Some polyhedra related to combinatorial problems (Linear Algebra Appl. 2 (1969), 451–558, doi:10.1016/0024-3795(69)90017-2). The facets of these group polyhedra are valid inequalities for every integer program whose basis produces the same group, and they became the source of cutting planes for integer programming: Gomory's mixed-integer cut and the later theory of corner polyhedra and subadditive valid inequalities descend from them.

The number of facets grows fast with the order of the group, so a list of facets for each group is not a usable description. Section 3D of the paper, "Lifting up Faces", gives a structural tool instead: facets of the polyhedron of a small group produce facets of the polyhedron of every larger group that maps onto it. This mission formalizes that construction (THEOREM 19) and its converse (THEOREM 20).

Setting

Let G\mathcal GG be a finite Abelian group, written additively, with zero 0ˉ\bar 00ˉ and order D=∣G∣D=|\mathcal G|D=∣G∣, and let G+=G−{0ˉ}\mathcal G^+=\mathcal G-\{\bar 0\}G+=G−{0ˉ}. For a right-hand side g0∈Gg_0\in\mathcal Gg0​∈G, the group equation asks for nonnegative integers t(g)t(g)t(g), g∈G+g\in\mathcal G^+g∈G+, with

∑g∈G+t(g)⋅g=g0.\sum_{g\in\mathcal G^+}t(g)\cdot g=g_0 .g∈G+∑​t(g)⋅g=g0​.

Its solution set is T(G,g0)T(\mathcal G,g_0)T(G,g0​); when g0=0ˉg_0=\bar 0g0​=0ˉ, the zero solution is excluded, as stipulated on p. 474. Each solution is a point of RG+\mathbb R^{\mathcal G^+}RG+, a space of dimension n′=D−1n'=D-1n′=D−1, and the master polyhedron P(G,g0)P(\mathcal G,g_0)P(G,g0​) is the convex hull of T(G,g0)T(\mathcal G,g_0)T(G,g0​). Gomory reads a solution ttt as a path from 0ˉ\bar 00ˉ to g0g_0g0​ that uses the element ggg exactly t(g)t(g)t(g) times; with arc lengths π(g)\pi(g)π(g) its length is π⋅t=∑gπ(g)t(g)\pi\cdot t=\sum_g\pi(g)t(g)π⋅t=∑g​π(g)t(g).

An inequality π⋅t≥π0\pi\cdot t\ge\pi_0π⋅t≥π0​, written (π,π0)(\pi,\pi_0)(π,π0​), is a face of P(G,g0)P(\mathcal G,g_0)P(G,g0​) when π≠0\pi\ne0π=0, every t∈T(G,g0)t\in T(\mathcal G,g_0)t∈T(G,g0​) satisfies it, and the solutions on the hyperplane π⋅t=π0\pi\cdot t=\pi_0π⋅t=π0​ generate that hyperplane: every point of it is a weighted sum, with total weight 1, of such solutions. A face is therefore a facet; Gomory says "face" throughout, and so does this mission. The solutions with π⋅t=π0\pi\cdot t=\pi_0π⋅t=π0​ are the minimal paths.

Let ψ\psiψ be a homomorphism of G\mathcal GG onto a second finite Abelian group H\mathcal HH, with kernel K=ψ−1(0ˉ)\mathcal K=\psi^{-1}(\bar 0)K=ψ−1(0ˉ). A coefficient vector π′\pi'π′ on H+\mathcal H^+H+ is lifted to G+\mathcal G^+G+ by

π(g)=π′(ψg),π′(0ˉ)=0,\pi(g)=\pi'(\psi g),\qquad \pi'(\bar 0)=0,π(g)=π′(ψg),π′(0ˉ)=0,

so that every element of the kernel gets coefficient 000.

Formalization targets

Goal: THEOREM 19 (p. 486)

If ψg0≠0ˉ\psi g_0\ne\bar 0ψg0​=0ˉ and (π′,π0)(\pi',\pi_0)(π′,π0​) is a face of P(H,ψg0)P(\mathcal H,\psi g_0)P(H,ψg0​) with π0>0\pi_0>0π0​>0, then

(π,π0),π(g)=π′(ψg),is a face of P(G,g0).\Big(\pi,\pi_0\Big),\quad \pi(g)=\pi'(\psi g),\quad\text{is a face of } P(\mathcal G,g_0).(π,π0​),π(g)=π′(ψg),is a face of P(G,g0​).

For instance, the face t1≥1t_1\ge1t1​≥1 of P(G2,(1))P(\mathcal G_2,(1))P(G2​,(1)) lifts through reduction mod 2 to the face t1+0t2+t3+0t4+t5≥1t_1+0t_2+t_3+0t_4+t_5\ge1t1​+0t2​+t3​+0t4​+t5​≥1 of P(G6,(5))P(\mathcal G_6,(5))P(G6​,(5)) (p. 488).

Milestones

  1. Face criterion (p. 469, proof of THEOREM 7). For π0>0\pi_0>0π0​>0, (π,π0)(\pi,\pi_0)(π,π0​) is a face iff it is valid on TTT and n′n'n′ linearly independent solutions satisfy π⋅t=π0\pi\cdot t=\pi_0π⋅t=π0​.
  2. Pushing paths forward (p. 486). A solution ttt for G\mathcal GG maps to the solution τ(h)=∑g∈ψ−1ht(g)\tau(h)=\sum_{g\in\psi^{-1}h}t(g)τ(h)=∑g∈ψ−1h​t(g) for H\mathcal HH of the same length; hence the lifted inequality is valid.
  3. Lifted minimal paths (p. 487). From a minimal path τ\tauτ for H\mathcal HH and a kernel element kkk, the path Tk(τ)T_k(\tau)Tk​(τ) is a minimal path for G\mathcal GG.
  4. Rank D−1D-1D−1 (pp. 487–488). The lifted inequality has D−1D-1D−1 linearly independent minimal paths.

Further result: THEOREM 20 (p. 489)

Conversely, a face (π,π0)(\pi,\pi_0)(π,π0​) of P(G,g0)P(\mathcal G,g_0)P(G,g0​) with π0>0\pi_0>0π0​>0 and π(g)=0\pi(g)=0π(g)=0 for some g≠0ˉg\ne\bar0g=0ˉ is the lift of a face of P(H,ψg0)P(\mathcal H,\psi g_0)P(H,ψg0​) for some ψ\psiψ of G\mathcal GG onto a group H\mathcal HH with ψg=0ˉ\psi g=\bar 0ψg=0ˉ.

Significance

THEOREMS 19 and 20 together say that the faces of P(G,g0)P(\mathcal G,g_0)P(G,g0​) with π0>0\pi_0>0π0​>0 and a zero coefficient are exactly the faces lifted from proper quotients of G\mathcal GG. A catalogue of these positive-right-hand-side faces therefore only needs the faces with all coefficients positive; the rest come from smaller groups. Gomory uses this to explain the tables of Appendix 5: for G2,2\mathcal G_{2,2}G2,2​ and G2,2,2\mathcal G_{2,2,2}G2,2,2​ every such face is a lift of the single nontrivial face of P(G2,(1))P(\mathcal G_2,(1))P(G2​,(1)), and for a direct sum G=K1⊕K2\mathcal G=\mathcal K_1\oplus\mathcal K_2G=K1​⊕K2​ every face of the polyhedron of a summand not containing g0g_0g0​ extends to a face for G\mathcal GG (p. 489). The same idea, sending facets of one group problem to facets of another by a homomorphism, recurs in the later theory of group relaxations and infinite group problems.

Both theorems are proved in the paper. To the knowledge of this mission, neither the group polyhedra, their faces nor the lifting theorem has a machine-checked formalization; the platform's other "lifting" results (lifted cover facets of the knapsack polytope, cone lifts of convex sets) are different notions. The mission produces a reusable definition layer for master group polyhedra and their faces, stated once for an arbitrary finite Abelian group, and a formal version of the passage between faces of a group and of its quotients.

Difficulty

Validity of the lifted inequality is a one-line computation (milestone 2). The content is that the lift is a facet: one must exhibit D−1D-1D−1 affinely independent solutions on the hyperplane, while the quotient face only supplies ∣H∣−1|\mathcal H|-1∣H∣−1 of them. A solution of the quotient problem does not determine a solution for G\mathcal GG; coset representatives must be chosen and the path closed by a kernel element, and the kernel coordinates, which all have coefficient zero, must still be filled to full rank.

The obvious reduction, "a facet pulls back to a facet under a surjective map", fails here because the map from G\mathcal GG-paths to H\mathcal HH-paths is a linear projection from dimension D−1D-1D−1 to ∣H∣−1|\mathcal H|-1∣H∣−1 and does not lift independence by itself. It also fails at π0=0\pi_0=0π0​=0: with G=Z6\mathcal G=\mathbb Z_6G=Z6​, H=Z3\mathcal H=\mathbb Z_3H=Z3​, g0=1g_0=1g0​=1, the face t(1)≥0t(1)\ge0t(1)≥0 of P(Z3,(1))P(\mathbb Z_3,(1))P(Z3​,(1)) lifts to t(1)+t(4)≥0t(1)+t(4)\ge0t(1)+t(4)≥0, which is valid but not a face of P(Z6,(1))P(\mathbb Z_6,(1))P(Z6​,(1)).

Formalization scope

The formalization uses one definition file for an arbitrary finite Abelian group ([AddCommGroup G] [Fintype G] [DecidableEq G]), instantiated at both G\mathcal GG and H\mathcal HH. It commits to the following readings:

  • Vectors are indexed by the subtype {g∣g≠0ˉ}\{g\mid g\ne\bar0\}{g∣g=0ˉ}, so 0ˉ\bar00ˉ has no coordinate and the dimension of T-space is ∣G∣−1|\mathcal G|-1∣G∣−1 for each group separately; solutions are N\mathbb NN-valued, nonzero, and cast to R\mathbb RR. When g0≠0ˉg_0\ne\bar0g0​=0ˉ, the group equation already excludes the zero vector.
  • "Generated by the points on the hyperplane" means that the affine span of the tight solutions equals the hyperplane {x:π⋅x=π0}\{x:\pi\cdot x=\pi_0\}{x:π⋅x=π0​}; π≠0\pi\ne0π=0 is part of the definition and π0≥0\pi_0\ge0π0​≥0 is not (THEOREM 6 proves it).
  • Paths are vectors and "minimal path" means a solution with π⋅t=π0\pi\cdot t=\pi_0π⋅t=π0​.
  • π′(0ˉ)=0\pi'(\bar0)=0π′(0ˉ)=0 is built into the lift; "onto" is Function.Surjective ψ, and g0∉Kg_0\notin\mathcal Kg0​∈/K is ψg0≠0ˉ\psi g_0\ne\bar0ψg0​=0ˉ.
  • Added hypothesis π0>0\pi_0>0π0​>0 in THEOREM 19 and the rank milestone: the page omits it, and the Z6→Z3\mathbb Z_6\to\mathbb Z_3Z6​→Z3​ example above shows the statement is false without it.
  • The face criterion is THEOREM 7's "basic feasible solution" unfolded into n′n'n′ linearly independent tight solutions, stated for the master polyhedron only and for nontrivial groups, where n′>0n'>0n′>0.
  • Pushing a path forward uses ψg0≠0ˉ\psi g_0\ne\bar0ψg0​=0ˉ, the standing hypothesis of THEOREM 19, so the pushed path cannot become the excluded zero solution.
  • The lifted path Tk(τ)T_k(\tau)Tk​(τ) puts its closing 111 on the kernel element g0−∑g∉Ktk(g)⋅gg_0-\sum_{g\notin\mathcal K}t_k(g)\cdot gg0​−∑g∈/K​tk​(g)⋅g itself, and adds nothing when that element is 0ˉ\bar00ˉ.
  • In THEOREM 20, "nontrivial" is π0>0\pi_0>0π0​>0, the printed (π1,π0)(\pi_1,\pi_0)(π1​,π0​) is read as (π,π0)(\pi,\pi_0)(π,π0​), and the conclusion requires ψg=0ˉ\psi g=\bar0ψg=0ˉ for the given zero coefficient: without it the identity map would satisfy the statement.

A formalization that proves only validity of the lift, or that defines faces without the generation condition (so that every valid inequality is a "face"), would be trivial and is ruled out by the definition of IsFace.

The development needs finite sums over fibres of ψ\psiψ, a choice of coset representatives, and a rank argument for a block matrix with full-rank diagonal blocks. The definitions of TTT, P(G,g0)P(\mathcal G,g_0)P(G,g0​) and face are reusable for any further result on master group polyhedra. Proofs of the milestones in any order, alternative proofs of the face criterion, and a proof of THEOREM 20 are welcome.

Selected references

  • R. E. Gomory, Some polyhedra related to combinatorial problems, Linear Algebra and Its Applications 2 (1969), 451–558. doi:10.1016/0024-3795(69)90017-2
  • R. E. Gomory, On the relation between integer and noninteger solutions to linear programs, Proceedings of the National Academy of Sciences 53 (1965), 260–265. doi:10.1073/pnas.53.2.260
  • R. E. Gomory and E. L. Johnson, Some continuous functions related to corner polyhedra, Mathematical Programming 3 (1972), 23–85. doi:10.1007/BF01584976
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Dynamic ProgrammingLinear OptimizationOptimization+2·Captain: mikedeng1

Linear Programming and Finite Markovian Control Problems IX: Discounted Semi-Markov Decision Processes — the Value Vector Is the Smallest Superharmonic VectorTextbook

Why semi-Markov control matters

A decision maker may choose an action whenever a system changes state, even though the time until the next change is random. Maintenance and replacement decisions are examples: a repair choice changes both the next operating state and the length of time before another choice is available. A discrete-time Markov decision model gives every decision epoch the same duration. A semi-Markov decision process allows the holding time to depend on the current state, action, and next state. Kallenberg's Chapter 7 develops discounted and average-reward versions of this finite-state model, including a characterization of optimal discounted value through inequalities and a linear program (Kallenberg 1983, Chapter 7).

The mission focuses on the discounted part of that chapter. The result is known: Kallenberg proves it as Theorem 7.2.1. The formalization task is to state and eventually prove the result for the same policy class and the same general holding-time distributions. Those details matter because a stationary-policy-only version would omit the comparison that the theorem makes with all admissible policies.

Model and notation

Let EEE be a nonempty finite state set. Each i∈Ei\in Ei∈E has a finite nonempty set A(i)A(i)A(i) of available actions. If action a∈A(i)a\in A(i)a∈A(i) is chosen, the next state is jjj with probability piaj≥0p_{iaj}\geq0piaj​≥0, with ∑jpiaj=1\sum_jp_{iaj}=1∑j​piaj​=1. Conditional on jjj, the nonnegative time until that transition has distribution FiajF_{iaj}Fiaj​. The action earns a lump reward riar_{ia}ria​ immediately and a reward rate sias_{ia}sia​ during the ensuing sojourn. Neither the holding-time laws nor the reward rates are required to be identical across transitions (Kallenberg 1983, pp. 210–212).

Fix a positive continuous discount rate λ\lambdaλ. The factor applied after elapsed time ttt is e−λte^{-\lambda t}e−λt. Assumption 7.2.1 requires the Laplace–Stieltjes transform of each conditional holding-time law to satisfy

Liaj(λ):=∫0∞e−λt dFiaj(t)<1(i,j∈E, a∈A(i)).L_{iaj}(\lambda):=\int_0^\infty e^{-\lambda t}\,dF_{iaj}(t)<1 \qquad(i,j\in E,\ a\in A(i)).Liaj​(λ):=∫0∞​e−λtdFiaj​(t)<1(i,j∈E, a∈A(i)).

The assumption is strict for every triple; it does not require a particular parametric family. Define the one-epoch expected discounted reward and discounted transition entry by

ria∗=∑jpiaj∫0∞(ria+sia∫0te−λu du)dFiaj(t),piaj∗=piajLiaj(λ).r^*_{ia}=\sum_jp_{iaj}\int_0^\infty \left(r_{ia}+s_{ia}\int_0^t e^{-\lambda u}\,du\right)dF_{iaj}(t), \qquad p^*_{iaj}=p_{iaj}L_{iaj}(\lambda).ria∗​=j∑​piaj​∫0∞​(ria​+sia​∫0t​e−λudu)dFiaj​(t),piaj∗​=piaj​Liaj​(λ).

A policy RRR is a sequence of randomized decisions. At each epoch its choice may depend on every previously observed state and chosen action and the current state. It does not observe previous sojourn times when choosing. Let viλ(R)v_i^\lambda(R)viλ​(R) be the expected sum of discounted lump and rate rewards from initial state iii, as in equation (7.2.1). The DRD value vector is viλ=sup⁡Rviλ(R)v_i^\lambda=\sup_R v_i^\lambda(R)viλ​=supR​viλ​(R), with the supremum over this full policy class. A real vector www is DRD-superharmonic when

wi≥ria∗+∑jpiaj∗wj(i∈E, a∈A(i)).w_i\geq r^*_{ia}+\sum_jp^*_{iaj}w_j \qquad(i\in E,\ a\in A(i)).wi​≥ria∗​+j∑​piaj∗​wj​(i∈E, a∈A(i)).

These are Kallenberg's Definitions 7.2.1 and 7.2.2 (p. 214).

Formalization targets

Goal: the smallest superharmonic vector

Theorem 7.2.1 states that the value vector itself satisfies the superharmonic inequalities and lies below every other vector satisfying them:

vλ is DRD-superharmonic,w is DRD-superharmonic⟹viλ≤wi(i∈E).v^\lambda\text{ is DRD-superharmonic},\qquad w\text{ is DRD-superharmonic}\Longrightarrow v_i^\lambda\leq w_i\quad(i\in E).vλ is DRD-superharmonic,w is DRD-superharmonic⟹viλ​≤wi​(i∈E).

This is the mission goal. Its content includes both clauses: merely showing that every superharmonic vector bounds policy rewards would leave out the assertion that the value is superharmonic.

Supporting results

Lemma 7.2.1 expresses viλ(R)v_i^\lambda(R)viλ​(R) as a sum of expected discounted state-action occupancies. Theorem 7.2.2 says that a feasible action choice attaining equality in the value equation at every state yields an optimal pure stationary policy. Theorem 7.2.3 says that positive support in an optimal solution of the dual linear program (7.2.11) yields such a policy. Their statements and exact conditions are the mission's milestones (Kallenberg 1983, pp. 212, 216–217).

What the results provide

The goal turns a supremum over possibly history-dependent randomized policies into a finite system of inequalities indexed by available state-action pairs. The next two results explain how equality in those inequalities and an optimal linear-programming solution identify a pure stationary policy with the same value. Thus the finite programs are statements about the original semi-Markov process rather than a separate discounted matrix model (Kallenberg 1983, pp. 214–217).

Kallenberg supplies paper proofs. This mission supplies Lean statements and a shared model interface; the proposed theorem items still require machine-checked proofs. The model's arbitrary conditional holding-time measures, its distinction between the current reward and future discounted value, and its history-dependent policy class are reusable for other finite semi-Markov arguments. The average-reward section of Chapter 7 is outside this mission.

Main difficulty

The straightforward finite-state discounted equation concerns expected one-step rewards. The source value, however, is defined from rewards earned at random physical times, and a policy can react to its entire discrete history. Identifying these two descriptions requires accounting for the joint state, action, and elapsed-time law without giving the policy access to past holding times. Holding times may equal zero with positive probability; the strict transform assumption excludes a degenerate law concentrated entirely at zero, but it does not make each holding time strictly positive. A second issue is real-valuedness: all-policy suprema and integrals must represent the finite quantities in the source rather than Lean's default values on ill-posed inputs.

Formalization scope

Lean uses finite types for states and actions, with a nonempty available-action finset for each state. Conditional sojourn laws are probability measures on R\mathbb RR with no mass on negative times. This includes arbitrary distributions on [0,∞)[0,\infty)[0,∞) and permits an atom at zero. The model stores stochastic transition rows, lump rewards, and rate rewards. The discounted layer stores λ>0\lambda>0λ>0 and the source's strict per-triple transform inequality. The expected one-epoch reward uses a Lebesgue integral against each sojourn law.

Policy decision rules read a list of past state-action pairs and the current state. Finite-horizon value recursively integrates the original epoch reward and its holding-time discount; policy value is its limit. The coordinatewise supremum ranges over all such policies. The occupation sum of Lemma 7.2.1 is a separate theorem, not the definition of policy value. A proof must establish convergence and boundedness from Assumption 7.2.1 so that Lean's total limit, integral, and real-supremum operations have their intended meanings.

The dual program sums only over actions available in each state and uses strictly positive weights βj\beta_jβj​, equality flow constraints, and nonnegative variables as printed. Its pure stationary selector must choose an available action with positive dual mass in every state. Contributions establishing the occupation identity, value bounds, Bellman inequalities, or dual decoding all fit the mission.

Selected references

  • L. C. M. Kallenberg, Linear Programming and Finite Markovian Control Problems, Mathematical Centre Tracts 148, Mathematisch Centrum, Amsterdam, 1983, Chapter 7, pp. 210–227. Publisher repository.
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Graph TheoryProbabilityStochastic Systems·Captain: mikedeng1

Loss Networks 3: If E e^{λX} < ∞ and 2E(X − C)⁺ < E(C − X)⁺, i.i.d. Loads on the Complete Graph Fit on Direct and Two-Edge Routes with Probability → 1Research Paper

Motivation

Telephone and data networks are often fully connected at the core: every pair of switching centres has a direct trunk group, and a call that finds its direct trunk full may be alternatively routed over a two-link path through a third, tandem centre. Whether such a network can absorb fluctuations of demand without losing traffic depends on how the spare capacity of lightly loaded links can be borrowed by heavily loaded ones. F. P. Kelly's survey Loss networks (Ann. Appl. Probab., 1991, doi:10.1214/aoap/1177005872) studies this question in §4.6, "Results respecting graph structure", by looking at a static snapshot of the network: random loads on the edges of a complete graph, and the question whether they can all be carried at once.

Timeline:

  • Hajek (1986, personal communication cited as [21] in the survey) considered edges coloured independently red (a pair that needs twice an edge's capacity) or white (an idle edge), and proved that for red probability p<1/3p<1/3p<1/3 all red pairs can be served over white two-edge paths with probability tending to one (Theorem 4.40, p. 357).
  • Hajek (1987, [22]) showed the same threshold for routing each red pair on a single two-edge path (Theorem 4.43, quoted, p. 358).
  • Kelly (1991, Theorem 4.45, p. 358) extended Theorem 4.40 from two-valued loads to general i.i.d. loads with an exponential moment, under the condition 2 E(X−C)+<E(C−X)+2\,\mathbb E(X-C)^+<\mathbb E(C-X)^+2E(X−C)+<E(C−X)+. The survey gives the proof in full on p. 359.

Setting

Let K≥1K\ge 1K≥1 and consider the complete graph on the nodes {0,…,K−1}\{0,\dots,K-1\}{0,…,K−1}: every unordered pair e={a,b}e=\{a,b\}e={a,b} of distinct nodes is an edge, and there are 12K(K−1)\tfrac12K(K-1)21​K(K−1) edges. Every edge has capacity CCC.

An offered load xe≥0x_e\ge 0xe​≥0 is attached to every edge eee. The load of e={a,b}e=\{a,b\}e={a,b} may be carried on its direct edge, or on a two-edge route a−k−ba-k-ba−k−b through a tandem node k∉ek\notin ek∈/e; such a route uses the two edges {a,k}\{a,k\}{a,k} and {b,k}\{b,k\}{b,k}. Loads are divisible. The loads are routable with capacity CCC if there are flows fe,k≥0f_{e,k}\ge 0fe,k​≥0 (zero when k∈ek\in ek∈e) with ∑kfe,k≤xe\sum_k f_{e,k}\le x_e∑k​fe,k​≤xe​ such that on every edge ggg

(xg−∑kfg,k)+∑(e,k): g on the route of e via kfe,k≤C,\Big(x_g-\sum_k f_{g,k}\Big)+\sum_{(e,k):\ g \text{ on the route of } e \text{ via } k} f_{e,k}\le C,(xg​−k∑​fg,k​)+(e,k): g on the route of e via k∑​fe,k​≤C,

i.e. the directly carried part of ggg's own load plus all two-edge traffic passing through ggg is at most CCC.

The loads are random: XXX is a nonnegative real random variable with law μ\muμ, and the loads (xe)(x_e)(xe​) are independent, each distributed as XXX. Let

P(K)=P{the loads on the complete graph on K nodes are routable}.P(K)=\mathbb P\{\text{the loads on the complete graph on } K \text{ nodes are routable}\}.P(K)=P{the loads on the complete graph on K nodes are routable}.

Write E(X−C)+\mathbb E(X-C)^+E(X−C)+ for the mean excess of a load over the capacity and E(C−X)+\mathbb E(C-X)^+E(C−X)+ for the mean spare capacity.

Formalization targets

Goal: Theorem 4.45

If E eλX<∞\mathbb E\,e^{\lambda X}<\inftyEeλX<∞ for some λ>0\lambda>0λ>0 and

2 E(X−C)+<E(C−X)+(4.46)2\,\mathbb E(X-C)^+<\mathbb E(C-X)^+ \tag{4.46}2E(X−C)+<E(C−X)+(4.46)

then

P(K)→1(K→∞).P(K)\to 1 \qquad (K\to\infty).P(K)→1(K→∞).

Milestones (in the order of the proof on pp. 358–359)

With m=E(C−X)+m=\mathbb E(C-X)^+m=E(C−X)+ and 0<ε<m20<\varepsilon<m^20<ε<m2:

  1. In each triangle the cyclic reservations (xe−C)+(C−xak)+(C−xbk)+/((K−2)(m2−ε))(x_e-C)^+(C-x_{ak})^+(C-x_{bk})^+/((K-2)(m^2-\varepsilon))(xe​−C)+(C−xak​)+(C−xbk​)+/((K−2)(m2−ε)) are positive at most once, and only for an overloaded edge routed through two underloaded ones.
  2. If every overloaded edge's reservations cover its excess and no underloaded edge has more reserved through it than it has spare, the loads are routable.
  3. E Y=m2/(m2−ε)>1\mathbb E\,Y=m^2/(m^2-\varepsilon)>1EY=m2/(m2−ε)>1 for Y=(C−X2)+(C−X3)+/(m2−ε)Y=(C-X_2)^+(C-X_3)^+/(m^2-\varepsilon)Y=(C−X2​)+(C−X3​)+/(m2−ε).
  4. E Z=2 E(X−C)+ m/(m2−ε)<1\mathbb E\,Z=2\,\mathbb E(X-C)^+\,m/(m^2-\varepsilon)<1EZ=2E(X−C)+m/(m2−ε)<1 for small ε\varepsilonε, for Z=((X1−C)+(C−X2)++(X2−C)+(C−X1)+)/(m2−ε)Z=\big((X_1-C)^+(C-X_2)^++(X_2-C)^+(C-X_1)^+\big)/(m^2-\varepsilon)Z=((X1​−C)+(C−X2​)++(X2​−C)+(C−X1​)+)/(m2−ε).
  5. (4.48): P{∑i≤nn−1Yi<1}≤e−nI1\mathbb P\{\sum_{i\le n} n^{-1}Y_i<1\}\le e^{-nI_1}P{∑i≤n​n−1Yi​<1}≤e−nI1​ for some I1>0I_1>0I1​>0 and all n≥1n\ge 1n≥1.
  6. (4.49): P{∑i≤nn−1Zi>1}≤e−nI2\mathbb P\{\sum_{i\le n} n^{-1}Z_i>1\}\le e^{-nI_2}P{∑i≤n​n−1Zi​>1}≤e−nI2​ for some I2>0I_2>0I2​>0 and all n≥1n\ge 1n≥1, when XXX has an exponential moment.
  7. 1−P(K)≤12K(K−1) (P1(K−2)+P2(K−2))1-P(K)\le\tfrac12K(K-1)\,(P_1(K-2)+P_2(K-2))1−P(K)≤21​K(K−1)(P1​(K−2)+P2​(K−2)).

Significance

The result. Theorem 4.45 says that in a large fully connected network, a load distribution whose mean overload is less than half its mean spare capacity can be served, with probability tending to one, using only direct and two-link routes. The factor 222 reflects that each unit of overflow occupies two edges. The condition is explicit and depends on the load law only through two expectations, which makes it a design rule: capacity CCC suffices for large networks as soon as (4.46) holds. Comment 4.47 (p. 358) observes that the two-valued law P{X=2}=p\mathbb P\{X=2\}=pP{X=2}=p, P{X=0}=1−p\mathbb P\{X=0\}=1-pP{X=0}=1−p with C=1C=1C=1 turns (4.46) into Hajek's threshold p<1/3p<1/3p<1/3.

Formalizing it. The theorem is proved, in full, on one page; nothing here is open. The mission produces a machine-checked version of that proof: a precise routing event on the complete graph, the deterministic reservation argument, two Cramér–Chernoff tail bounds with rates uniform in the network size, and the union bound over edges. To our knowledge none of these steps is formalized anywhere, and no routing-feasibility event on a complete graph exists in Mathlib or on the platform.

Difficulty

The obvious approach is to fix an overloaded edge and argue that, among its K−2K-2K−2 two-edge routes, enough pass through two underloaded edges. That count is binomial and concentrates, but it does not prove the theorem: the same underloaded edge sits on two-edge routes of many overloaded pairs, so the reservations of different pairs compete for its spare capacity, and the routing decisions of different pairs are dependent. Any argument has to control, simultaneously at every edge, both the flow an overloaded edge can shed and the flow an underloaded edge is asked to absorb, with exponential rates uniform in KKK so that a union bound over 12K(K−1)\tfrac12K(K-1)21​K(K−1) edges survives; this is the central difficulty. On the upper side the variables are unbounded, so the exponential moment of XXX has to be carried over to the reserved flows.

Formalization scope

The Lean development lives in the namespace KellyLossNetworks.Routing.

  • Edges are {e : Sym2 (Fin K) // ¬ e.IsDiag}; loads are functions Edge K → ℝ.
  • The routing event Feasible C x lets each pair split its load arbitrarily over its direct edge and all its two-edge routes, charges a two-edge flow to both edges of its route, requires the flow sent away from a pair not to exceed its load, and imposes the capacity constraint on every edge, including overloaded ones. Flows are real (divisible loads); integer routing is a different problem.
  • "Independent and identically distributed" is the product measure Measure.pi (fun _ : Edge K => μ), with μ a probability measure on ℝ and X ≥ 0 almost surely. P(K)P(K)P(K) is the measure of the routing event; no measurability of that event is presupposed.
  • Expectations are Bochner integrals. The exponential moment is integrability of t↦eθtt\mapsto e^{\theta t}t↦eθt for some θ>0\theta>0θ>0; it implies that (X−C)+(X-C)^+(X−C)+ is integrable, so (4.46) compares genuine expectations.
  • The limit is along K∈NK\in\mathbb NK∈N with no lower bound on KKK; for K≤2K\le 2K≤2 there are no two-edge routes, which does not affect the limit.
  • The rates I1,I2I_1,I_2I1​,I2​ in (4.48) and (4.49) are quantified before nnn and may not depend on it.

A trivializing formalization is ruled out: an event that allows only direct routing, ignores the capacity of the edges a detour passes through, or lets a pair send away more than its load would make the statement false or empty, and the routing event here does none of these.

A complete development needs Fubini on product measures, the Cramér–Chernoff method for averages of i.i.d. variables (both tails, with the lower tail for a bounded nonnegative variable and the upper tail under an exponential moment), measure-preserving projections of Measure.pi, and combinatorics of the triangles through an edge of the complete graph. The Cramér–Chernoff bounds with rates uniform in nnn are reusable well beyond this mission. Contributions to any milestone are welcome; the deterministic milestones 1–2 and the expectation computations 3–4 are independent of the tail bounds and of each other.

Selected references

  • F. P. Kelly, Loss networks, Annals of Applied Probability 1(3):319–378, 1991. doi:10.1214/aoap/1177005872
  • B. Hajek, Average case analysis of greedy algorithms for Kelly's triangle problem and the independent set problem, 26th IEEE Conference on Decision and Control, 1987 (reference [22] of the survey; the source of Theorem 4.43).
  • H. Chernoff, A measure of asymptotic efficiency for tests of a hypothesis based on the sum of observations, Annals of Mathematical Statistics 23(4):493–507, 1952. doi:10.1214/aoms/1177729330
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CombinatoricsLinear Optimization·Captain: mikedeng1

Some Polyhedra Related to Combinatorial Problems I: A Minimizing Vertex of the Corner Polyhedron Gives an Optimal Integer Solution When the Right-Hand Side Lies Deep in the Basis ConeResearch Paper

Motivation

An integer program in standard form, max⁡c⋅x\max c\cdot xmaxc⋅x subject to Ax=bAx=bAx=b, x≥0x\ge 0x≥0, xxx integer, is easy to state and hard to solve. Its linear programming relaxation is solved by the simplex method, which ends at an optimal basis BBB. The difference between the two problems is the integrality of xxx, and R. E. Gomory's 1969 paper Some polyhedra related to combinatorial problems (Linear Algebra Appl. 2, 451–558, doi:10.1016/0024-3795(69)90017-2) measures that difference with a finite Abelian group. Dropping only the nonnegativity of the basic variables leaves a problem whose feasible set has a simple structure: its convex hull is the corner polyhedron, and its integer points are the solutions of an equation in the group Zm/BZm\mathbb Z^m/B\mathbb Z^mZm/BZm.

The first section of the paper uses this to explain when the integer program reduces to the group problem. Its asymptotic theorem says that if bbb lies far enough inside the cone where BBB is LP-optimal, then an optimal solution of the group problem gives an optimal integer solution. Gomory first proved this in 1965 (On the relation between integer and noninteger solutions to linear programs, PNAS 53, 260–265, doi:10.1073/pnas.53.2.260) by a counting argument on group solutions. The 1969 paper proves it again through the geometry of the corner polyhedron. The vertices of that polyhedron are irreducible, and an irreducible point is small. This is the version formalized here.

Setting

Let A=(B,N)A=(B,N)A=(B,N) be an integer m×(m+n)m\times(m+n)m×(m+n) matrix, with BBB a nonsingular m×mm\times mm×m matrix (the first mmm columns) and NNN the m×nm\times nm×n matrix of the remaining columns N1,…,NnN_1,\dots,N_nN1​,…,Nn​. As on p. 452 of the paper, AAA contains an m×mm\times mm×m unit matrix. Let b∈Zmb\in\mathbb Z^mb∈Zm and c=(cB,cN)∈Rm+nc=(c_B,c_N)\in\mathbb R^{m+n}c=(cB​,cN​)∈Rm+n. The integer program (2) is

max⁡ c⋅xsubject toAx=b, x≥0, x integer.\max\ c\cdot x\quad\text{subject to}\quad Ax=b,\ x\ge 0,\ x\ \text{integer}.max c⋅xsubject toAx=b, x≥0, x integer.

The relative prices are cN∗=−cN+cBB−1Nc_N^*=-c_N+c_BB^{-1}NcN∗​=−cN​+cB​B−1N, with components cm+i∗c^*_{m+i}cm+i∗​. BBB is an optimal LP basis when B−1b≥0B^{-1}b\ge0B−1b≥0 and every cm+i∗≥0c^*_{m+i}\ge0cm+i∗​≥0.

The group is G=M(I)/M(B)\mathcal G=M(I)/M(B)G=M(I)/M(B), where M(I)=ZmM(I)=\mathbb Z^mM(I)=Zm and M(B)M(B)M(B) is the lattice of integer combinations of the columns of BBB. Let fff be the quotient map. G\mathcal GG is finite of order D=∣det⁡B∣D=|\det B|D=∣detB∣. Let N\mathcal NN be the set of nonzero elements among fN1,…,fNnfN_1,\dots,fN_nfN1​,…,fNn​, and put g0=fbg_0=fbg0​=fb. The group equation is

∑g∈Nt(g)⋅g=g0,t(g)∈Z≥0,\sum_{g\in\mathcal N}t(g)\cdot g=g_0,\qquad t(g)\in\mathbb Z_{\ge0},g∈N∑​t(g)⋅g=g0​,t(g)∈Z≥0​,

and P(G,N,g0)P(\mathcal G,\mathcal N,g_0)P(G,N,g0​) is the convex hull of its solutions in RN\mathbb R^{\mathcal N}RN. A nonnegative integer ttt is irreducible if any two integer vectors s,rs,rs,r with 0≤s,r≤t0\le s,r\le t0≤s,r≤t and ∑s(g)⋅g=∑r(g)⋅g\sum s(g)\cdot g=\sum r(g)\cdot g∑s(g)⋅g=∑r(g)⋅g are equal.

The group costs are c∗(g)=min⁡i: fNi=gcm+i∗c^*(g)=\min_{i:\,fN_i=g}c^*_{m+i}c∗(g)=mini:fNi​=g​cm+i∗​. Problem (8) minimizes ∑gc∗(g)t(g)\sum_g c^*(g)t(g)∑g​c∗(g)t(g) over the solutions of the group equation. A corresponding vertex of a point t∗t^*t∗ (Remark 1) is a nonnegative integer xN∗x_N^*xN∗​ with three properties: t∗(g)=∑i: fNi=gxm+i∗t^*(g)=\sum_{i:\,fN_i=g}x^*_{m+i}t∗(g)=∑i:fNi​=g​xm+i∗​; at most one positive xm+i∗x^*_{m+i}xm+i∗​ per group element; and xm+i∗=0x^*_{m+i}=0xm+i∗​=0 when fNi=0ˉfN_i=\bar0fNi​=0ˉ. It uses only least cost columns when xm+i∗>0x^*_{m+i}>0xm+i∗​>0 only if cm+i∗=c∗(fNi)c^*_{m+i}=c^*(fN_i)cm+i∗​=c∗(fNi​).

The cone KB={y∈Rm:B−1y≥0}K_B=\{y\in\mathbb R^m:B^{-1}y\ge0\}KB​={y∈Rm:B−1y≥0} is where BBB is LP-optimal. KB(d)K_B(d)KB​(d) is the set of its points at Euclidean distance at least ddd from its frontier, and lmax⁡l_{\max}lmax​ is the largest Euclidean length of a column NiN_iNi​.

Formalization targets

Goal: THEOREM 4 (p. 462)

If b∈KB(lmax⁡(D−1))b\in K_B\bigl(l_{\max}(D-1)\bigr)b∈KB​(lmax​(D−1)), t∗t^*t∗ is a vertex of P(G,N,fb)P(\mathcal G,\mathcal N,fb)P(G,N,fb) minimizing (8), and xN∗x_N^*xN∗​ is a corresponding vertex using only least cost columns, then

x∗=(B−1(b−NxN∗), xN∗) is an optimal integer solution of (2).x^*=\bigl(B^{-1}(b-Nx_N^*),\ x_N^*\bigr)\ \text{is an optimal integer solution of (2)}.x∗=(B−1(b−NxN∗​), xN∗​) is an optimal integer solution of (2).

Milestones

  1. THEOREM 1 (p. 459): an irreducible ttt satisfies ∏g∈N(1+t(g))≤∣G∣\prod_{g\in\mathcal N}(1+t(g))\le|\mathcal G|∏g∈N​(1+t(g))≤∣G∣.
  2. p. 459: ∣M(I)/M(B)∣=∣det⁡B∣|M(I)/M(B)|=|\det B|∣M(I)/M(B)∣=∣detB∣.
  3. THEOREM 2 (p. 460): every vertex of P(G,N,g0)P(\mathcal G,\mathcal N,g_0)P(G,N,g0​) is an irreducible integer point.
  4. Proof of THEOREM 4 (p. 462): an irreducible ttt has ∑gt(g)≤∣G∣−1\sum_g t(g)\le|\mathcal G|-1∑g​t(g)≤∣G∣−1.
  5. THEOREM 3 (p. 461): a corresponding vertex of a minimizing vertex is optimal for (2) whenever B−1(b−NxN∗)≥0B^{-1}(b-Nx_N^*)\ge0B−1(b−NxN∗​)≥0.
  6. Proof of THEOREM 4: ∥NxN∗∥≤lmax⁡(D−1)\|Nx_N^*\|\le l_{\max}(D-1)∥NxN∗​∥≤lmax​(D−1).
  7. Proof of THEOREM 4: y∈KB(d)y\in K_B(d)y∈KB​(d) and ∥v∥≤d\|v\|\le d∥v∥≤d imply y−v∈KBy-v\in K_By−v∈KB​.
  8. Display (10), p. 464: the optimal x∗x^*x∗ of THEOREM 4 satisfies ∏i=1n(1+xm+i∗)≤D\prod_{i=1}^n(1+x^*_{m+i})\le D∏i=1n​(1+xm+i∗​)≤D.

Significance

THEOREM 4 explains why integer programs with large right-hand sides are often easy. For bbb outside a band of width lmax⁡(D−1)l_{\max}(D-1)lmax​(D−1) along the boundary of the cone KBK_BKB​, the integer optimum is the LP solution B−1bB^{-1}bB−1b plus a correction that depends only on the class of bbb modulo the lattice of BBB. That correction is periodic in bbb and can be tabulated once for the DDD group elements. The vertex form adds a structural statement: the correction can be read off a vertex of the corner polyhedron, so the faces and vertices of corner polyhedra (the subject of the rest of the paper) bear directly on integer programming. Display (10) shows such optimal solutions are small: their nonbasic parts satisfy ∏(1+xm+i)≤D\prod(1+x_{m+i})\le D∏(1+xm+i​)≤D.

All results here are proved in the paper. None of them is formalized yet. The 1965 version of the asymptotic theorem, posed as THEOREM 1 of the mission on Gomory (1965), is in the same family of statements. Its hypothesis is that the group solution is optimal; here the hypothesis is that it is a vertex of P(G,N,g0)P(\mathcal G,\mathcal N,g_0)P(G,N,g0​). Its counting lemma (a short optimal group solution) and its bound ∥Ny∥≤(D−1)l\|Ny\|\le(D-1)l∥Ny∥≤(D−1)l have counterparts in milestones 4 and 6.

Difficulty

An optimal solution of the group problem need not be short. Group solutions can be made longer without changing their cost (add a zero-cost cycle), and a long nonbasic part xNx_NxN​ can push b−NxNb-Nx_Nb−NxN​ out of the cone KBK_BKB​, making xBx_BxB​ negative. The goal must therefore use the vertex hypothesis. Vertices are irreducible (THEOREM 2), and irreducibility bounds their size through a pigeonhole count in the finite group (THEOREM 1). The formal work is spread across three settings: the convex geometry of the hull of an infinite integer set, the arithmetic of the quotient Zm/BZm\mathbb Z^m/B\mathbb Z^mZm/BZm and its order, and the linear algebra connecting B−1B^{-1}B−1, the cone KBK_BKB​ and its frontier.

Formalization scope

Vectors on N\mathcal NN are functions on the subtype of a Finset of the group. Integer solutions are N\mathbb NN-valued, and P(G,N,g0)P(\mathcal G,\mathcal N,g_0)P(G,N,g0​) is convexHull ℝ of their real images. THEOREMS 1 and 2 are stated for an arbitrary finite Abelian group and any finite N\mathcal NN not containing 000, as in the paper. The integer program uses the basis as the first mmm columns, Fin m ⊕ Fin n. B−1B^{-1}B−1 is the real matrix inverse, and every statement assumes det⁡B≠0\det B\ne0detB=0. Optimality of xxx means feasibility plus c⋅y≤c⋅xc\cdot y\le c\cdot xc⋅y≤c⋅x for every feasible integer yyy; no supremum is taken. The loose phrases of the paper are read as follows.

  • "Vertex" means an extreme point.
  • "Every vertex is irreducible" means every vertex is the image of an integer solution, and that solution is irreducible.
  • "Optimal linear programming basis" means B−1b≥0B^{-1}b\ge0B−1b≥0 and c∗≥0c^*\ge0c∗≥0.
  • "Corresponding vertex x∗x^*x∗ of Px(B,N,b)P_x(B,N,b)Px​(B,N,b)" means conditions (i)–(iii) of Remark 1, which the paper calls easily verified, taken as the definition.
  • "Minimizing (8)" means minimizing over the integer solutions of the group equation.
  • "Euclidean distance from the frontier" is explicit: ∑kvk2\sqrt{\sum_k v_k^2}∑k​vk2​​, frontier in the usual topology, so KB(0)=KBK_B(0)=K_BKB​(0)=KB​.

The printed slip on p. 462 (∑(1+t∗(g))\sum(1+t^*(g))∑(1+t∗(g)) for THEOREM 1's product) is corrected in the Lean. Dropping the vertex hypothesis from the goal is not a valid formalization (the statement becomes false in general), and neither is defining irreducibility with real r,sr,sr,s (THEOREM 1 then fails). The existence of a minimizing vertex, which the paper takes for granted, is not posed.

A complete development needs the finiteness and order of Zm/BZm\mathbb Z^m/B\mathbb Z^mZm/BZm (Mathlib's AddSubgroup.index_eq_natAbs_det), extreme points of convex hulls (extremePoints_convexHull_subset), and an argument that a closed cone with yyy deep inside it contains the ball around yyy. The group-polyhedron definitions and THEOREMS 1–2 apply to any finite Abelian group and are reusable beyond this mission. Proofs of any milestone are welcome independently.

Selected references

  • R. E. Gomory, Some polyhedra related to combinatorial problems, Linear Algebra and Its Applications 2 (1969), 451–558. https://doi.org/10.1016/0024-3795(69)90017-2
  • R. E. Gomory, On the relation between integer and noninteger solutions to linear programs, Proc. Nat. Acad. Sci. USA 53 (1965), 260–265. https://doi.org/10.1073/pnas.53.2.260
  • R. E. Gomory, Faces of an integer polyhedron, Proc. Nat. Acad. Sci. USA 57 (1967), 16–18. https://doi.org/10.1073/pnas.57.1.16
  • B. L. van der Waerden, Modern Algebra, English translation of the 2nd German edition, Ungar, New York, 1949–1950 (computation of fff, cited p. 456).
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