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Mathematical Logic

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Formal VerificationTheoretical Computer Science·Captain: Rizwan G Mir

The Cook-Levin Theorem: NP-Completeness of Boolean Satisfiability in Lean 4Research Paper

Introduction

The Cook-Levin theorem states that CNF SAT is NP-complete. This mission formalizes the theorem over a multi-tape Turing machine model in Lean 4.

Main Goal

Prove CookLevin.cook_levin_theorem:

NPCompleteSATNPComplete SATNPCompleteSAT

under the decider and reduction hypotheses.

175 thms9 active usersReviewed
🏆Completed
Number Theory·Captain: marwahaha

Single-fold Diophantine representationsOpen Problem

Motivation

A Diophantine representation translates membership in a computably enumerable set into solvability of one polynomial equation. The source asks whether the entire auxiliary witness can always be made unique. The source is OpenAI's September 2026 manuscript.

Setting

A single-fold representation uses an integer polynomial evaluated at natural-number inputs and witnesses. An input belongs to the represented set exactly when a witness makes the polynomial zero, and any two complete witnesses at the same input must coincide.

Formalization targets

a∈S ⟺ ∃w∈Nm: P(a,w)=0,P(a,w)=P(a,v)=0 ⟹ w=v.a\in S\ \Longleftrightarrow\ \exists w\in\mathbb N^m:\ P(a,w)=0,\qquad P(a,w)=P(a,v)=0\ \Longrightarrow\ w=v.a∈S ⟺ ∃w∈Nm: P(a,w)=0,P(a,w)=P(a,v)=0 ⟹ w=v.

The proposition MainStatement holds, which asserts a single-fold Diophantine representation result with positive lengths. For every n ≥ 1 and every set S of n-tuples of natural numbers that is recursively enumerable (the predicate a ∈ S is REPred, i.e. semi-decidable), there exist m ≥ 1 and an integer polynomial P in the n input variables and m witness variables such that P represents S. Here evaluation takes natural-number inputs a and witnesses w, casts them to integers, and evaluates P over ℤ. Representation means that for every input a, first, a lies in S exactly when some witness w ∈ ℕ^m satisfies P(a,w)=0, and second, any two witnesses w and v with P(a,w)=0 and P(a,v)=0 are equal. So there is exactly one witness for members of S and none for nonmembers, and the uniqueness condition is imposed on every input a, with all witness coordinates included in w.

The goal is OAI.SingleFold.main.

Significance

The target strengthens mere existential representation by enforcing uniqueness across every auxiliary coordinate. It covers all recursively enumerable sets of positive-length natural tuples. The manuscript presents a mathematical argument; the attached goal currently has status Open, so the requested contribution is a Lean proof of the displayed formal target.

Difficulty

Encoding a computation with extra witnesses can introduce many choices that do not affect the represented input. Every such auxiliary choice must be controlled, rather than leaving hidden coordinates outside the uniqueness assertion.

Formalization scope

The input length and witness length are positive naturals. REPred expresses recursive enumerability. Polynomial coefficients and evaluation are integral, while variables range over naturals cast into the integers. The goal is existential and does not require one fixed universal polynomial.

The shared definitions are supplied by SingleFold. The target uses Lean 4.33.1 with the pinned Mathlib environment. Contributions should preserve the existing statement and develop reusable supporting lemmas in its definitions.

Selected references

  • OpenAI, Single-fold Diophantine representations, preprint, September 2026. Manuscript.
  • OpenAI, accompanying Lean statement. Pinned source.
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Model Theory·Captain: wurtle

A CH obstruction to a prescribed categoricity thresholdResearch Paper

Motivation: categoricity transfer beyond first-order logic

A class of structures is categorical in a cardinal λ\lambdaλ if it has exactly one model of size λ\lambdaλ up to isomorphism. Morley's theorem (1965) says that a complete first-order theory in a countable language categorical in one uncountable cardinal is categorical in all of them. Abstract elementary classes (AECs), introduced by Shelah in the 1970s, axiomatize classes of structures with a well-behaved notion of strong substructure (closure under isomorphism, coherence, unions of chains, a Löwenheim–Skolem number) without requiring first-order axiomatizability; they cover classes defined in infinitary logics and many classes of modules. Shelah's categoricity conjecture for AECs, in its eventual form, asks whether categoricity in one sufficiently large cardinal transfers to all sufficiently large cardinals. A sharper, prescribed-threshold form names the threshold explicitly: with H(K)=ℶ(2LS(K))+H(K)=\beth_{(2^{\mathrm{LS}(K)})^+}H(K)=ℶ(2LS(K))+​ (the Hanf number for existence of arbitrarily large models), categoricity in some λ≥H(K)\lambda\ge H(K)λ≥H(K) should imply categoricity in every μ≥H(K)\mu\ge H(K)μ≥H(K). Deciding which form holds is a central question in non-elementary model theory.

Timeline

  • 1965 — Morley proves the categoricity theorem for countable first-order theories (Trans. AMS 1965).
  • 1970s — Shelah introduces abstract elementary classes (see the survey of Boney–Vasey 2017, §2).
  • 1990 — Hart and Shelah show categoricity in Lω1,ωL_{\omega_1,\omega}Lω1​,ω​ can stop at ℵk\aleph_kℵk​ while holding for ℵ0,…,ℵk−1\aleph_0,\dots,\aleph_{k-1}ℵ0​,…,ℵk−1​, using finite-support group constructions (Israel J. Math. 1990).
  • 1999 — Shelah proves categoricity transfer for AECs with amalgamation and records H(K)H(K)H(K) as the existence bound (APAL 1999).
  • 2016 — Kolesnikov and Lambie-Hanson study Hanf numbers for amalgamation of coloring classes (JSL 2016).
  • 2017 — Vasey proves downward transfer from a successor cardinal ≥H(K)\ge H(K)≥H(K) under amalgamation and tameness (APAL 2017).
  • 2021 — Grossberg distinguishes the prescribed-bound and eventual forms of the conjecture (A Course in Model Theory I, draft, Ch. 2 §4).
  • 2022–2023 — Espíndola announces a proof of eventual categoricity and, for arbitrary AECs, a transfer at the prescribed bound via accessible categories (arXiv:1906.09169, arXiv:2301.13167).
  • 2024 — Šaroch and Trlifaj state the prescribed-threshold formulation with both endpoints included (Bull. LMS 2024, §2.3).
  • 2026 — An OpenAI preprint, A CH obstruction to a prescribed categoricity threshold (OpenAI Math Release, September 24, 2026), claims that under CH there is an AEC with LS(K)=ℵ0\mathrm{LS}(K)=\aleph_0LS(K)=ℵ0​, categorical on a tail but with two nonisomorphic models at H(K)=ℶω2H(K)=\beth_{\omega_2}H(K)=ℶω2​​; this conflicts with the transfer claimed in Espíndola 2023 (Theorem 4.1). The preprint has not been peer reviewed and its theorem is not formally verified.

Setting

Work with relational structures in a finitary relational language LLL (a set of relation symbols, each with a finite arity). An abstract elementary class is a class KKK of LLL-structures with a relation M⪯KNM\preceq_K NM⪯K​N ("strong substructure") such that: ⪯K\preceq_K⪯K​ is a partial order on KKK refining substructure; KKK and ⪯K\preceq_K⪯K​ are closed under isomorphism; coherence holds (M0⊆M1M_0\subseteq M_1M0​⊆M1​, M0⪯KM2M_0\preceq_K M_2M0​⪯K​M2​, M1⪯KM2M_1\preceq_K M_2M1​⪯K​M2​ imply M0⪯KM1M_0\preceq_K M_1M0​⪯K​M1​); unions of ⪯K\preceq_K⪯K​-chains are in KKK, are ⪯K\preceq_K⪯K​-extensions of each member, and are ⪯K\preceq_K⪯K​-below any common strong extension (Tarski–Vaught chain axioms); and there is a Löwenheim–Skolem number LS(K)≥ℵ0+∣L∣\mathrm{LS}(K)\ge\aleph_0+|L|LS(K)≥ℵ0​+∣L∣, the least θ\thetaθ such that every subset AAA of a model MMM lies in some N⪯KMN\preceq_K MN⪯K​M with ∣N∣≤∣A∣+θ|N|\le|A|+\theta∣N∣≤∣A∣+θ.

KKK is categorical in μ\muμ if it has a model of size μ\muμ and any two models of size μ\muμ are isomorphic. The beth numbers are ℶ0=ℵ0\beth_0=\aleph_0ℶ0​=ℵ0​, ℶα+1=2ℶα\beth_{\alpha+1}=2^{\beth_\alpha}ℶα+1​=2ℶα​ and suprema at limits. CH is 2ℵ0=ℵ12^{\aleph_0}=\aleph_12ℵ0​=ℵ1​.

Formalization targets

Goal: the CH counterexample (Theorem 1.1)

Assume CH. There is a countable finitary relational language LLL and an AEC KKK in LLL such that

LS(K)=ℵ0,ℶω2=H(K)=ℶ(2ℵ0)+,\mathrm{LS}(K)=\aleph_0,\qquad \beth_{\omega_2}=H(K)=\beth_{(2^{\aleph_0})^+},LS(K)=ℵ0​,ℶω2​​=H(K)=ℶ(2ℵ0​)+​,

KKK has two nonisomorphic models of cardinality ℶω2\beth_{\omega_2}ℶω2​​, and, with Λ=ℶ(2ℵ1)+\Lambda=\beth_{(2^{\aleph_1})^+}Λ=ℶ(2ℵ1​)+​,

K is categorical in every cardinal μ≥Λ.K\ \text{is categorical in every cardinal } \mu\ge\Lambda .K is categorical in every cardinal μ≥Λ.

Hence Λ+>H(K)\Lambda^+>H(K)Λ+>H(K) is a categoricity cardinal from which downward transfer to H(K)H(K)H(K) fails, while eventual categoricity holds for this KKK. The goal statement is published on the platform with status Open: no machine-checked proof exists yet.

Significance

The result itself. Combined with Gödel's relative consistency of CH, the theorem shows that, if ZFC is consistent, the prescribed-threshold form of Shelah's categoricity conjecture is not provable in ZFC (Corollary 6.1 of the source). The example has no amalgamation or joint embedding, which explains why transfer theorems that assume these properties are unaffected, and it is compatible with qualitative eventual categoricity. It also bears directly on a published claim of transfer at the prescribed bound for arbitrary AECs.

Formalizing it. Because the result contradicts a claimed theorem in the literature, an independent machine check is especially valuable. The formal statement verifies every AEC axiom explicitly, not just the two cardinal-arithmetic conclusions. Mathlib has cardinal and ordinal arithmetic, beth numbers and ZFC sets; the AEC framework built here would be reusable for further formal work in non-elementary model theory.

Difficulty

Most known non-transfer examples (Hart–Shelah) fail categoricity at small cardinals; here the failure must happen exactly at the Hanf-type bound ℶω2\beth_{\omega_2}ℶω2​​ while categoricity holds on a tail, with LS(K)=ℵ0\mathrm{LS}(K)=\aleph_0LS(K)=ℵ0​. One must build two nonisomorphic models at ℶω2\beth_{\omega_2}ℶω2​​ and simultaneously rule out any two nonisomorphic models above Λ\LambdaΛ, without appealing to a general eventual-categoricity theorem. All AEC axioms — in particular coherence and smoothness of unions of arbitrary directed chains — must survive the exceptional sets used to separate models, which is where naive constructions break.

Formalization scope

  • Structures are RelModel L with carrier a ZFSet (so cardinalities are cardinals of genuine sets in universe u) and relations indexed by arity; L:N→L:\mathbb N\toL:N→ Type with ΣnLn\Sigma_n L_nΣn​Ln​ countable.
  • ClassData packages a predicate of objects and a strong-substructure relation; IsAEC lists the axioms above, including invariance under isomorphisms compatible with inclusion, coherence, and the union axioms for chains indexed by any nonzero ordinal.
  • HasLSNumber ℵ₀ asserts both that ℵ0\aleph_0ℵ0​ is a Löwenheim–Skolem bound and that it is the least one.
  • hanf κ = beth (succ (2^κ)).ord, endpoint = beth (ω_2), tailThreshold = beth (succ (2^{ℵ₁})).ord. The equality endpoint = hanf ℵ₀ is part of the goal (it follows from CH).
  • Categorical μ includes existence of a model of size μ\muμ; TwoModels asks for two nonisomorphic models of the given size.
  • CH is a hypothesis of the theorem, in the same universe; the consistency corollary is not part of the goal.

Selected references

  • M. Morley, Categoricity in power, Trans. Amer. Math. Soc. (1965). https://doi.org/10.1090/S0002-9947-1965-0175782-0
  • K. Gödel, The consistency of the axiom of choice and of the generalized continuum-hypothesis, Proc. Nat. Acad. Sci. USA (1938). https://doi.org/10.1073/pnas.24.12.556
  • B. Hart and S. Shelah, Categoricity over P for first order T or categoricity for φ∈L_{ω1ω} can stop at ℵ_k while holding for ℵ_0,…,ℵ_{k−1}, Israel J. Math. (1990). https://doi.org/10.1007/BF02807869
  • S. Shelah, Categoricity for abstract classes with amalgamation, Ann. Pure Appl. Logic (1999). https://doi.org/10.1016/S0168-0072(98)00016-5
  • S. Vasey, Downward categoricity from a successor inside a good frame, Ann. Pure Appl. Logic (2017). https://doi.org/10.1016/j.apal.2016.10.003
  • W. Boney and S. Vasey, A survey on tame abstract elementary classes, 2017. https://arxiv.org/abs/1512.00060
  • A. Kolesnikov and C. Lambie-Hanson, The Hanf number for amalgamation of coloring classes, J. Symb. Logic (2016). https://doi.org/10.1017/jsl.2015.48
  • J. Šaroch and J. Trlifaj, Deconstructible abstract elementary classes of modules and categoricity, Bull. London Math. Soc. (2024). https://doi.org/10.1112/blms.13172
  • C. Espíndola, A complete classification of categoricity spectra of accessible categories with directed colimits, preprint, 2023. https://arxiv.org/abs/2301.13167
  • OpenAI, A CH obstruction to a prescribed categoricity threshold, OpenAI Math Release preprint, September 24, 2026 (source of the goal; Theorem 1.1, p. 2). https://github.com/openai/math/blob/adc7f1241b42e322a6451854ab7e4b4c146bf78a/preprints/A-CH-Obstruction-to-a-Prescribed-Categoricity-Threshold-September-24-2026/paper.pdf
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Topology·Captain: Lucas

Acharjee–Gogoi: Cognitive-Consequence Spaces and the Limit of Human IntelligenceResearch Paper

Motivation

In 1998 Smale listed eighteen problems for the twenty-first century; the eighteenth asks: What are the limits of intelligence, both artificial and human? (Smale 1998). The paper of Acharjee and Gogoi (arXiv:2310.10792) proposes a mathematical model of a mind, the cognitive-consequence space, built from a Tarski consequence operator, and claims that two theorems about this model (its Theorems 3.3 and 3.4) show that "human intelligence is limitless" (Discussion, p. 22; Conclusion, p. 23).

This mission records that model and its results in Lean 4 exactly as stated, so that each claim becomes a checkable statement: proved, or refuted, by the verifier rather than by argument.

Setting

A cognitive-consequence space is a set CCC of mental representations (thoughts) with a consequence operator Cn:P(C)→P(C)\mathrm{Cn} : \mathcal P(C) \to \mathcal P(C)Cn:P(C)→P(C) and an implication connective X⇒YX \Rightarrow YX⇒Y satisfying Tarski's axioms as listed in the paper (p. 5):

  1. CCC is countable;
  2. A⊆Cn(A)A \subseteq \mathrm{Cn}(A)A⊆Cn(A);
  3. A⊆B⇒Cn(A)⊆Cn(B)A \subseteq B \Rightarrow \mathrm{Cn}(A) \subseteq \mathrm{Cn}(B)A⊆B⇒Cn(A)⊆Cn(B);
  4. Cn(Cn(A))=Cn(A)\mathrm{Cn}(\mathrm{Cn}(A)) = \mathrm{Cn}(A)Cn(Cn(A))=Cn(A);
  5. if X∈Cn(A)X \in \mathrm{Cn}(A)X∈Cn(A) then X∈Cn(B)X \in \mathrm{Cn}(B)X∈Cn(B) for some finite B⊆AB \subseteq AB⊆A;
  6. if Y∈Cn(A∪{X})Y \in \mathrm{Cn}(A \cup \{X\})Y∈Cn(A∪{X}) then (X⇒Y)∈Cn(A)(X \Rightarrow Y) \in \mathrm{Cn}(A)(X⇒Y)∈Cn(A);

together with the paper's standing assumption Cn(∅)≠∅\mathrm{Cn}(\varnothing) \neq \varnothingCn(∅)=∅ (p. 5, after Definition 3.2).

A set AAA is deductive if Cn(A)=A\mathrm{Cn}(A) = ACn(A)=A. The cognitive-consequence topology is

τ={A⊆C:Cn(C∖A)=C∖A},\tau = \{A \subseteq C : \mathrm{Cn}(C \setminus A) = C \setminus A\},τ={A⊆C:Cn(C∖A)=C∖A},

and members of τ\tauτ are consequence-wise open (CWO). The cognitive closure Cl□(A)\mathrm{Cl}^{\square}(A)Cl□(A) is the intersection of all deductive sets containing AAA.

Separately, a cognitive similarity distance is a function Cog:C×C→[0,1]\mathrm{Cog} : C \times C \to [0,1]Cog:C×C→[0,1] together with a relation x≈yx \approx yx≈y ("xxx and yyy cognitively coincide") such that Cog(x,y)=0  ⟺  x≈y\mathrm{Cog}(x,y) = 0 \iff x \approx yCog(x,y)=0⟺x≈y, Cog\mathrm{Cog}Cog is symmetric, x≈z⇒Cog(x,y)=Cog(z,y)x \approx z \Rightarrow \mathrm{Cog}(x,y) = \mathrm{Cog}(z,y)x≈z⇒Cog(x,y)=Cog(z,y), and the triangle inequality holds. A sequence of thoughts (xn)(x_n)(xn​) converges to xxx if for every ε∈(0,1)\varepsilon \in (0,1)ε∈(0,1) we have Cog(x,xn)<ε\mathrm{Cog}(x, x_n) < \varepsilonCog(x,xn​)<ε for all large nnn.

Formalization targets

Goal: "human intelligence is limitless" (Theorems 3.3 and 3.4)

For every cognitive-consequence space,

(∃f∈C ∀A∈τ, f∉A) ∧ (∃f∈C ∃A∈τ, f∈A).\Bigl(\exists f \in C\ \forall A \in \tau,\ f \notin A\Bigr) \ \wedge\ \Bigl(\exists f \in C\ \exists A \in \tau,\ f \in A\Bigr).(∃f∈C ∀A∈τ, f∈/A) ∧ (∃f∈C ∃A∈τ, f∈A).

Milestones

Theorems 3.1, 3.2, 3.3, 3.5, 3.6, 3.7 and Corollary 3.5 (the topology τ\tauτ and the cognitive closure), Theorems 3.8 and 3.12 (cognitive limits), Theorem 4.3 (the filter fdf_dfd​) and Theorem 5.1 (Gödel's incompleteness black hole).

Significance

The paper presents the goal as its answer to the human-intelligence half of Smale's eighteenth problem. A machine-checked verdict on the goal settles whether that conclusion follows from the paper's own axioms. The milestones are the paper's structural results about τ\tauτ, cognitive closure and cognitive limits.

Status: to the best of current knowledge none of these results has a published machine-checked proof. The first conjunct of the goal (Theorem 3.3) follows from the standing assumption Cn(∅)≠∅\mathrm{Cn}(\varnothing) \neq \varnothingCn(∅)=∅. The second conjunct (Theorem 3.4) is not implied by the listed axioms as formalized: the consequence operator with Cn(A)=C\mathrm{Cn}(A) = CCn(A)=C for every AAA, on a one-point language, satisfies all of them and has τ={∅}\tau = \{\varnothing\}τ={∅}. The goal is therefore expected to be resolved by a disproof.

Difficulty

The content of the goal is the existence of a nonempty CWO set, equivalently of a deductive set different from CCC (a consistent theory). The paper's argument for Theorem 3.4 derives ⋂iAi=∅\bigcap_i A_i = \varnothing⋂i​Ai​=∅ and calls this a contradiction; nothing in the axioms forbids it.

Formalization scope

  • CogCons.CognitiveConsequenceSpace C bundles Cn\mathrm{Cn}Cn, the connective imp, axioms (i)–(vi) and Cn(∅)≠∅\mathrm{Cn}(\varnothing) \neq \varnothingCn(∅)=∅. The syntax σ\sigmaσ and interpretation III of the paper are not modelled beyond imp.
  • The paper also asserts Cn(C)≠C\mathrm{Cn}(C) \neq CCn(C)=C (p. 5). This contradicts axiom (ii), since C⊆Cn(C)⊆CC \subseteq \mathrm{Cn}(C) \subseteq CC⊆Cn(C)⊆C; including it would make every statement vacuously true, so it is omitted. For the same reason the second half of Corollary 3.4 (Cl□(C)≠C\mathrm{Cl}^{\square}(C) \neq CCl□(C)=C) is not included.
  • Theorem 4.4 (the family {A:f∈Cn(A)}\{A : f \in \mathrm{Cn}(A)\}{A:f∈Cn(A)} is a filter) is not included: closure under intersection fails in a four-element model satisfying all the axioms.
  • Results relying on informal notions (the practical topology of Section 2, Theorems 3.9–3.11 and 3.13, Theorems 4.1–4.2, and the "solution space" of Section 5) are not formalized.
  • CogCons.CognitiveSimilarityDistance C is independent of Cn\mathrm{Cn}Cn; sequences are indexed by N\mathbb NN starting at 000; ≈\approx≈ is an arbitrary relation constrained only by the listed axioms.

Selected references

  • S. Acharjee, U. Gogoi, The limit of human intelligence, arXiv:2310.10792v2 [math.GM], 2023. https://arxiv.org/abs/2310.10792
  • S. Smale, Mathematical problems for the next century, The Mathematical Intelligencer 20(2), 7–15, 1998. https://doi.org/10.1007/BF03025291
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Captain: Lucas

Jech Set Theory I: Silver's Theorem on Singular CardinalsTextbook

Motivation

How large can the power set of an infinite set be? For a regular cardinal κ\kappaκ (one that is not the supremum of fewer than κ\kappaκ smaller ordinals) the answer is: almost anything. Easton's theorem (1970) shows that the function κ↦2κ\kappa \mapsto 2^{\kappa}κ↦2κ on regular cardinals can be prescribed arbitrarily in any model of ZFC, subject only to monotonicity and König's inequality cf⁡(2κ)>κ\operatorname{cf}(2^{\kappa}) > \kappacf(2κ)>κ. For a long time it was expected that singular cardinals — those that are such a supremum, like ℵω\aleph_{\omega}ℵω​ — would behave the same way.

They do not. In 1974 Jack Silver proved that the Generalized Continuum Hypothesis cannot fail for the first time at a singular cardinal of uncountable cofinality: if 2α=α+2^{\alpha} = \alpha^{+}2α=α+ for every infinite α<κ\alpha < \kappaα<κ and cf⁡κ>ω\operatorname{cf}\kappa > \omegacfκ>ω, then 2κ=κ+2^{\kappa} = \kappa^{+}2κ=κ+. This was the first ZFC theorem constraining the continuum function at singular cardinals, and it opened the area now called the singular cardinal problem.

A short timeline:

  • 1970 — Easton: the continuum function on regular cardinals is essentially arbitrary.
  • 1974 — Silver (ICM Vancouver): GCH cannot first fail at a singular cardinal of uncountable cofinality; more generally the Singular Cardinal Hypothesis is decided at cofinality ω\omegaω.
  • 1975 — Galvin and Hajnal: elementary inequalities for cardinal powers at singular cardinals of uncountable cofinality.
  • 1976–77 — Baumgartner and Prikry, and independently Jensen, give elementary (non-forcing, non-ultrapower) proofs of Silver's theorem; the proof reproduced in Jech's Chapter 8 is of this kind.
  • 1977 — Magidor: it is consistent, relative to large cardinals, that GCH holds below ℵω\aleph_{\omega}ℵω​ while 2ℵω>ℵω+12^{\aleph_{\omega}} > \aleph_{\omega+1}2ℵω​>ℵω+1​ — so Silver's restriction to uncountable cofinality is necessary.
  • 1980s onwards — Shelah's pcf theory, whose flagship result ℵωℵ0<ℵω4\aleph_{\omega}^{\aleph_0} < \aleph_{\omega_4}ℵωℵ0​​<ℵω4​​ (when ℵω\aleph_\omegaℵω​ is a strong limit) grows out of exactly the stationary-set machinery assembled here.

This mission is the first in a series formalizing Thomas Jech, Set Theory (Third Millennium Edition, Springer 2003). It covers Chapter 8, "Stationary Sets" (pp. 91–98).

Setting

Fix a regular uncountable cardinal κ\kappaκ and regard it as the well-ordered set of ordinals below it. A set C⊆κC \subseteq \kappaC⊆κ is closed unbounded, or a club, if it is unbounded in κ\kappaκ and contains all of its limit points below κ\kappaκ (an ordinal α>0\alpha > 0α>0 is a limit point of CCC when sup⁡(C∩α)=α\sup(C \cap \alpha) = \alphasup(C∩α)=α). A set S⊆κS \subseteq \kappaS⊆κ is stationary if S∩C≠∅S \cap C \neq \emptysetS∩C=∅ for every club CCC. Clubs are closed under intersections of fewer than κ\kappaκ of them, so they generate a κ\kappaκ-complete filter, the club filter; its dual is the nonstationary ideal.

The club filter has a second closure property with no analogue for ordinary filters. The diagonal intersection of a κ\kappaκ-indexed family is

△α<κXα  =  {ξ<κ  :  ξ∈⋂α<ξXα},\mathop{\triangle}_{\alpha<\kappa} X_\alpha \;=\; \Bigl\{ \xi < \kappa \;:\; \xi \in \bigcap_{\alpha<\xi} X_\alpha \Bigr\},△α<κ​Xα​={ξ<κ:ξ∈α<ξ⋂​Xα​},

and a filter closed under diagonal intersections is called normal. A function fff defined on S⊆κS \subseteq \kappaS⊆κ is regressive if f(α)<αf(\alpha) < \alphaf(α)<α for all nonzero α∈S\alpha \in Sα∈S.

For cardinal arithmetic, cf⁡κ\operatorname{cf}\kappacfκ denotes the cofinality of κ\kappaκ (the least length of an unbounded sequence in κ\kappaκ), κ+\kappa^{+}κ+ the cardinal successor, and κ\kappaκ is singular when cf⁡κ<κ\operatorname{cf}\kappa < \kappacfκ<κ. The Singular Cardinal Hypothesis (SCH) is the assertion that κcf⁡κ=κ+\kappa^{\operatorname{cf}\kappa} = \kappa^{+}κcfκ=κ+ for every singular κ\kappaκ with 2cf⁡κ<κ2^{\operatorname{cf}\kappa} < \kappa2cfκ<κ. A sequence of cardinals is normal if it is strictly increasing and continuous at limits.

Formalization targets

Goal — Silver's theorem (Jech 8.12)

κ singular, cf⁡κ>ω, (∀α ℵ0≤α<κ⇒2α=α+)  ⟹  2κ=κ+.\kappa \text{ singular},\ \operatorname{cf}\kappa > \omega,\ \bigl(\forall \alpha \ \aleph_0 \le \alpha < \kappa \Rightarrow 2^{\alpha} = \alpha^{+}\bigr) \;\Longrightarrow\; 2^{\kappa} = \kappa^{+}.κ singular, cfκ>ω, (∀α ℵ0​≤α<κ⇒2α=α+)⟹2κ=κ+.

This is the weakest statement of the chapter that still needs the full machinery: it fixes no particular κ\kappaκ and no particular cofinality, and it stays correct no matter how the singular cardinal problem develops above it.

Milestones, in dependency order

  1. Lemma 8.4 — the diagonal intersection of κ\kappaκ clubs is a club; equivalently the club filter is normal.
  2. Theorem 8.7 (Fodor) — a regressive function on a stationary set is constant on a stationary subset.
  3. Theorem 8.10 (Solovay) — every stationary subset of κ\kappaκ is the union of κ\kappaκ pairwise disjoint stationary sets.
  4. Lemma 8.14 — if ⟨κα⟩\langle \kappa_\alpha\rangle⟨κα​⟩ is normal with limit κ\kappaκ, λcf⁡κ<κ\lambda^{\operatorname{cf}\kappa} < \kappaλcfκ<κ for λ<κ\lambda < \kappaλ<κ, and {α:καcf⁡κα=κα+}\{\alpha : \kappa_\alpha^{\operatorname{cf}\kappa_\alpha} = \kappa_\alpha^{+}\}{α:καcfκα​​=κα+​} is stationary in cf⁡κ\operatorname{cf}\kappacfκ, then κcf⁡κ=κ+\kappa^{\operatorname{cf}\kappa} = \kappa^{+}κcfκ=κ+.
  5. Theorem 8.13 (Silver) — SCH at every cardinal of cofinality ω\omegaω implies SCH everywhere.

Significance

Silver's theorem is the boundary between the two halves of cardinal arithmetic. Above it sit the ZFC theorems of pcf theory; below it sit the consistency results (Magidor, Prikry, Radin forcing) that show how much freedom is left, and they are confined to cofinality ω\omegaω precisely because Theorems 8.12 and 8.13 close off everything else. Its proof also packages tools used throughout set theory: the normality of the club filter, Fodor's pressing-down lemma, and the technique of bounding almost disjoint families of functions by a stationary-set argument.

Formalization status: Mathlib already has clubs and stationary sets in an arbitrary well-ordered type (IsClub, IsStationary), with the finite and <κ<\kappa<κ-indexed intersection lemmas — Jech's Lemma 8.2 and Theorem 8.3. It does not have diagonal intersections, Fodor's theorem, Solovay's splitting theorem, or any singular cardinal arithmetic beyond the definitions of regular and singular cardinals. Each milestone below is therefore a genuine addition, and the first three are reusable well outside this mission.

Difficulty

The naive route to the goal — induct on α<κ\alpha < \kappaα<κ and pass to the limit — fails immediately: 2κ2^{\kappa}2κ for singular κ\kappaκ is not determined by the values 2α2^{\alpha}2α for α<κ\alpha < \kappaα<κ in any elementary way; that is exactly the content of the independence results. What the proof must do instead is bound the number of functions on cf⁡κ\operatorname{cf}\kappacfκ, and it gets that bound from a stationary set rather than from a club: uncountable cofinality is what makes the set of relevant stages stationary, and stationarity is what survives the diagonal argument. Cofinality ω\omegaω breaks this at the first step, since every subset of ω\omegaω that is unbounded is already a club and Fodor's theorem is empty.

The two hard pieces are Lemma 8.14 and, inside it, Lemma 8.16: given an almost disjoint family FFF of functions with f(α)∈Aαf(\alpha) \in A_\alphaf(α)∈Aα​ and ∣Aα∣≤ℵα|A_\alpha| \le \aleph_\alpha∣Aα​∣≤ℵα​ on a stationary set of α\alphaα, one must show ∣F∣≤ℵω1|F| \le \aleph_{\omega_1}∣F∣≤ℵω1​​ by assigning to each fff a pair (stationary set, bounded restriction) and checking the assignment is injective. That argument uses Fodor's theorem on a set of functions, and the bookkeeping does not simplify.

Formalization scope

The ambient order is Mathlib's type of ordinals below a cardinal, k.ord.ToType (written Below k in the mission's definition bundle); clubs and stationary sets are Mathlib's IsClub and IsStationary on that type, so a set is closed in the sense of being closed under suprema of directed subsets — equivalent, for a well-order with the order topology, to Jech's "contains its limit points". Families indexed by "α<κ\alpha < \kappaα<κ" are functions out of that same type, which is what makes the diagonal intersection typecheck without a side condition. Cardinal exponentiation, cofinality (Ordinal.cof of k.ord) and the successor cardinal (Order.succ) are Mathlib's.

One trivializing formalization is ruled out explicitly: the GCH hypothesis of the goal is stated for infinite cardinals α<κ\alpha < \kappaα<κ only. Quantified over all cardinals it would be unsatisfiable — 22=4≠3=2+2^{2} = 4 \neq 3 = 2^{+}22=4=3=2+ — and Silver's theorem would become vacuous.

A complete development needs: diagonal intersections and the normality of the club filter; Fodor's theorem; the sets Eλκ={α<κ:cf⁡α=λ}E^{\kappa}_{\lambda} = \{\alpha < \kappa : \operatorname{cf}\alpha = \lambda\}Eλκ​={α<κ:cfα=λ} and their stationarity; Solovay's splitting theorem via Lemmas 8.8 and 8.9; almost disjoint families of ordinal functions and the counting Lemmas 8.15 and 8.16; and Theorem 5.22(ii) on cardinal powers, which Jech's proof of Theorem 8.13 cites. Contributions of any of these as separate reductions are welcome, as are alternative proofs of the goal (for instance via a generic elementary embedding) that bypass some of the chain.

Selected references

  • Thomas Jech, Set Theory, The Third Millennium Edition, revised and expanded. Springer Monographs in Mathematics, Springer, 2003 (ISBN 3-540-44085-2). Chapter 8, "Stationary Sets", pp. 91–98 — the source of every statement in this mission.
  • Jack Silver, On the singular cardinals problem. Proceedings of the International Congress of Mathematicians (Vancouver, 1974), vol. 1, pp. 265–268.
  • William B. Easton, Powers of regular cardinals. Annals of Mathematical Logic 1 (1970), pp. 139–178.
  • Fred Galvin and András Hajnal, Inequalities for cardinal powers. Annals of Mathematics 101 (1975), pp. 491–498.
  • James E. Baumgartner and Karel Prikry, Singular cardinals and the generalized continuum hypothesis. American Mathematical Monthly 84 (1977), pp. 108–113.
  • Menachem Magidor, On the singular cardinals problem I. Israel Journal of Mathematics 28 (1977), pp. 1–31.
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Model Theory·Captain: Lucas

Continuum Hypothesis: Independence from ZFCTextbook

Motivation

The continuum hypothesis (CH) asks how large the set of real numbers is. Cantor proved that ∣Z∣=∣Q∣<∣R∣|\mathbb Z|=|\mathbb Q|<|\mathbb R|∣Z∣=∣Q∣<∣R∣, and in 1878 he conjectured that no set has cardinality strictly between that of the integers and that of the reals. Settling CH was the first of Hilbert's problems at the 1900 International Congress of Mathematicians. The question was answered by showing that it cannot be answered from the standard axioms: CH is independent of Zermelo–Fraenkel set theory with choice (ZFC).

Timeline (as recorded in the source):

  • 1878 — Cantor formulates the (weak) continuum hypothesis; in 1882 he states CH in its modern form.
  • 1900 — Hilbert lists CH as the first of his problems.
  • 1906 — Kőnig proves a theorem that, combined with Hausdorff's notion of cofinality (1908), shows 2ℵ0≠ℵα2^{\aleph_0}\neq\aleph_\alpha2ℵ0​=ℵα​ whenever α\alphaα has cofinality ω\omegaω.
  • 1930 — Patai: if 2κ2^\kappa2κ is the γ\gammaγ-th successor of κ\kappaκ for every infinite κ\kappaκ, then γ\gammaγ is finite.
  • 1938/1940 — Gödel proves that AC and GCH hold in the constructible universe LLL, so Con(ZF)⇒Con(ZFC+GCH)\mathrm{Con}(\mathrm{ZF})\Rightarrow\mathrm{Con}(\mathrm{ZFC}+\mathrm{GCH})Con(ZF)⇒Con(ZFC+GCH).
  • 1947 — Sierpiński proves that ZF + GCH implies AC.
  • 1963 — Cohen invents forcing and proves Con(ZFC)⇒Con(ZFC+¬CH)\mathrm{Con}(\mathrm{ZFC})\Rightarrow\mathrm{Con}(\mathrm{ZFC}+\neg\mathrm{CH})Con(ZFC)⇒Con(ZFC+¬CH).
  • 1986 — Freiling shows that ¬\neg¬CH is equivalent to his axiom of symmetry.

Setting

A cardinal measures the size of a set; ∣A∣≤∣B∣|A|\le|B|∣A∣≤∣B∣ means there is an injection A→BA\to BA→B, and ∣A∣=∣B∣|A|=|B|∣A∣=∣B∣ a bijection. ℵ0=∣N∣\aleph_0=|\mathbb N|ℵ0​=∣N∣, 2κ=∣P(X)∣2^\kappa=|\mathcal P(X)|2κ=∣P(X)∣ for ∣X∣=κ|X|=\kappa∣X∣=κ, so ∣R∣=2ℵ0|\mathbb R|=2^{\aleph_0}∣R∣=2ℵ0​. ℵα\aleph_\alphaℵα​ is the α\alphaα-th infinite cardinal, ℶα\beth_\alphaℶα​ is defined by ℶ0=ℵ0\beth_0=\aleph_0ℶ0​=ℵ0​, ℶα+1=2ℶα\beth_{\alpha+1}=2^{\beth_\alpha}ℶα+1​=2ℶα​ and suprema at limits, and cf\mathrm{cf}cf denotes cofinality.

  • CH: there is no S⊆RS\subseteq\mathbb RS⊆R with ℵ0<∣S∣<2ℵ0\aleph_0<|S|<2^{\aleph_0}ℵ0​<∣S∣<2ℵ0​ (CH).
  • GCH: for every infinite cardinal κ\kappaκ there is no cardinal strictly between κ\kappaκ and 2κ2^\kappa2κ (GCH).

Independence is a statement about a first-order theory. The language of set theory has one binary relation symbol ∈\in∈ (setLang). ZF (ZF) consists of extensionality, foundation, pairing, union, power set, infinity, and all instances of the separation and replacement schemas; ZFC (ZFC) adds the axiom of choice (acSentence). CH and GCH are also written as first-order sentences (chSentence, gchSentence): CH says that every subset of P(ω)\mathcal P(\omega)P(ω) injects into ω\omegaω or receives an injection from P(ω)\mathcal P(\omega)P(ω). A theory is consistent if it has a model, and T⊨σT\models\sigmaT⊨σ means σ\sigmaσ holds in every model of TTT; by Gödel's completeness theorem these coincide with the proof-theoretic notions.

Formalization targets

Goal: independence of CH from ZFC

Con(ZFC) ⟹ ZFC⊭CH  and  ZFC⊭¬CH.\mathrm{Con}(\mathrm{ZFC})\ \Longrightarrow\ \mathrm{ZFC}\not\models\mathrm{CH}\ \text{ and }\ \mathrm{ZFC}\not\models\neg\mathrm{CH}.Con(ZFC) ⟹ ZFC⊨CH  and  ZFC⊨¬CH.

Milestones

  1. ∣Z∣=∣Q∣<∣R∣|\mathbb Z|=|\mathbb Q|<|\mathbb R|∣Z∣=∣Q∣<∣R∣.
  2. CH  ⟺  2ℵ0=ℵ1\mathrm{CH}\iff 2^{\aleph_0}=\aleph_1CH⟺2ℵ0​=ℵ1​.
  3. (Gödel) Con(ZF)⇒Con(ZFC+GCH)\mathrm{Con}(\mathrm{ZF})\Rightarrow\mathrm{Con}(\mathrm{ZFC}+\mathrm{GCH})Con(ZF)⇒Con(ZFC+GCH).
  4. ZFC+GCH⊨CH\mathrm{ZFC}+\mathrm{GCH}\models\mathrm{CH}ZFC+GCH⊨CH.
  5. (Cohen) Con(ZFC)⇒Con(ZFC+¬CH)\mathrm{Con}(\mathrm{ZFC})\Rightarrow\mathrm{Con}(\mathrm{ZFC}+\neg\mathrm{CH})Con(ZFC)⇒Con(ZFC+¬CH).
  6. (Kőnig) 2ℵ0≠ℵα2^{\aleph_0}\neq\aleph_\alpha2ℵ0​=ℵα​ when cf(α)=ω\mathrm{cf}(\alpha)=\omegacf(α)=ω; and cf(2ℵ0)>ℵ0\mathrm{cf}(2^{\aleph_0})>\aleph_0cf(2ℵ0​)>ℵ0​.
  7. (Sierpiński) ZF+GCH⊨AC\mathrm{ZF}+\mathrm{GCH}\models\mathrm{AC}ZF+GCH⊨AC.
  8. GCH  ⟺  ∀α, 2ℵα=ℵα+1  ⟺  ∀α, ℵα=ℶα\mathrm{GCH}\iff\forall\alpha,\ 2^{\aleph_\alpha}=\aleph_{\alpha+1}\iff\forall\alpha,\ \aleph_\alpha=\beth_\alphaGCH⟺∀α, 2ℵα​=ℵα+1​⟺∀α, ℵα​=ℶα​.
  9. Under GCH: A<B⇒2A<2BA<B\Rightarrow 2^A<2^BA<B⇒2A<2B for infinite AAA, and the three-case formula for ℵαℵβ\aleph_\alpha^{\aleph_\beta}ℵαℵβ​​.
  10. (Patai) ∀α, 2ℵα=ℵα+γ\forall\alpha,\ 2^{\aleph_\alpha}=\aleph_{\alpha+\gamma}∀α, 2ℵα​=ℵα+γ​ implies γ<ω\gamma<\omegaγ<ω.
  11. (Freiling) the axiom of symmetry is equivalent to ¬\neg¬CH.

Significance

The independence theorem shows that the ZFC axioms do not determine the size of the continuum, and it was among the first independence results for a mathematical statement. Its two halves introduced the two main tools of modern set theory: inner models (Gödel's LLL) and forcing (Cohen). Many statements in analysis, topology and measure theory were subsequently shown to be independent by the same methods.

For formalization, the goal requires a model theory of ZF inside Lean: interpreting the first-order theory, building LLL inside an arbitrary model of ZF, and building forcing extensions of an arbitrary model of ZFC. The independence of CH was formalized in Lean 3 by the Flypitch project (Han and van Doorn, 2019–2020); no Lean 4 / Mathlib formalization is known to this mission. The cardinal-arithmetic milestones (2, 6, 8–11) are theorems of ZFC that can be proved directly in Lean's cardinal library.

Difficulty

Neither half follows from cardinal arithmetic in Lean: Lean's foundations prove that models of ZFC exist, but (assuming they are consistent) they prove neither CH nor its negation about Lean's own Cardinal, so an argument inside Lean's cardinal arithmetic alone cannot produce the required models. Gödel's half needs relativization of formulas to a definable inner class and absoluteness of the construction of LLL; Cohen's half needs a generic extension in which a new family of ℵ2\aleph_2ℵ2​ reals is added while cardinals are preserved (the countable chain condition). Both require a theory of definability inside arbitrary, possibly non-well-founded, models of a first-order theory.

Formalization scope

  • Syntax and semantics are Mathlib's first-order logic (FirstOrder.Language, Theory.IsSatisfiable, Theory.ModelsBoundedFormula written ⊨ᵇ). Formulas name their variables by natural numbers; close φ is the universal closure. Equality is the built-in equality of first-order logic.
  • Consistency is satisfiability and provability is semantic consequence over models with universe in Type. By the completeness theorem these agree with the syntactic notions.
  • The first-order CH uses P(ω)\mathcal P(\omega)P(ω) in place of R\mathbb RR (they are equinumerous in ZF). The first-order GCH is the choice-free form "∣S∣≤∣T∣≤∣P(S)∣⇒∣T∣=∣S∣|S|\le|T|\le|\mathcal P(S)|\Rightarrow|T|=|S|∣S∣≤∣T∣≤∣P(S)∣⇒∣T∣=∣S∣ or ∣T∣=∣P(S)∣|T|=|\mathcal P(S)|∣T∣=∣P(S)∣, for SSS infinite".
  • The consistency results carry their hypotheses Con(ZF)\mathrm{Con}(\mathrm{ZF})Con(ZF) or Con(ZFC)\mathrm{Con}(\mathrm{ZFC})Con(ZFC) explicitly, even though these are provable in Lean. A solution may not replace a model of ZF by Lean's own universe in a way that assumes CH or its negation.
  • Results that need large-cardinal hypotheses or class forcing (Easton, Solovay, Foreman–Woodin, Merimovich) are not part of this mission.

Reusable infrastructure: the language of set theory with standard abbreviations (pairs, functions, ω\omegaω, power set), satisfaction lemmas for the abbreviations, relativization, absoluteness, the constructible hierarchy, and forcing over models of ZFC.

Selected references

  • G. Cantor, Ein Beitrag zur Mannigfaltigkeitslehre, J. Reine Angew. Math. 84 (1878) 242–258. https://doi.org/10.1515/crll.1878.84.242
  • K. Gödel, The Consistency of the Axiom of Choice and of the Generalized Continuum-Hypothesis, PNAS 24 (1938) 556–557. https://doi.org/10.1073/pnas.24.12.556
  • K. Gödel, The Consistency of the Axiom of Choice and of the Generalized Continuum-Hypothesis with the Axioms of Set Theory, Annals of Math. Studies 3, Princeton, 1940.
  • P. J. Cohen, The independence of the Continuum Hypothesis I, II, PNAS 50 (1963) 1143–1148; 51 (1964) 105–110. https://doi.org/10.1073/pnas.50.6.1143, https://doi.org/10.1073/pnas.51.1.105
  • C. Freiling, Axioms of symmetry: throwing darts at the real number line, J. Symbolic Logic 51 (1986) 190–200. https://doi.org/10.2307/2273955
  • L. Gillman, Two classical surprises concerning the Axiom of Choice and the Continuum Hypothesis, Amer. Math. Monthly 109 (2002) 544–553. https://doi.org/10.2307/2695444
  • K. Kunen, Set Theory: An Introduction to Independence Proofs, North-Holland, 1980.
  • Wikipedia, Continuum hypothesis (revision 1376315189). https://en.wikipedia.org/w/index.php?title=Continuum_hypothesis&oldid=1376315189
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🏆Completed
Captain: Lucas

Singular Cardinals Hypothesis: Silver's reduction to countable cofinalityTextbook

Motivation

Cantor's continuum problem asks for the value of 2κ2^\kappa2κ, the cardinality of the power set of an infinite cardinal κ\kappaκ. For regular cardinals ZFC says almost nothing: by Easton's theorem (1970) the function κ↦2κ\kappa\mapsto 2^\kappaκ↦2κ on regular cardinals can be essentially any non-decreasing function with cf⁡(2κ)>κ\operatorname{cf}(2^\kappa)>\kappacf(2κ)>κ. For singular cardinals the picture is different, and the singular cardinals hypothesis (SCH) is the natural statement of "GCH-like behaviour at singular cardinals". It arose from the question whether the least cardinal at which GCH fails can be singular.

Timeline.

  • Scott (1961): if GCH holds on a set of measure one below a measurable cardinal κ\kappaκ, then 2κ=κ+2^\kappa=\kappa^+2κ=κ+.
  • Solovay (1974): above a strongly compact cardinal, SCH holds.
  • Silver (1975): if κ\kappaκ is singular of uncountable cofinality and 2λ=λ+2^\lambda=\lambda^+2λ=λ+ for all infinite λ<κ\lambda<\kappaλ<κ, then 2κ=κ+2^\kappa=\kappa^+2κ=κ+. Consequently the least failure of SCH has countable cofinality.
  • Silver, Prikry: from a supercompact cardinal, a model of ¬\neg¬SCH (Prikry forcing over a measurable κ\kappaκ with 2κ>κ+2^\kappa>\kappa^+2κ>κ+).
  • Gitik (1989–1991), building on Woodin: ZFC+¬SCH\mathrm{ZFC}+\neg\mathrm{SCH}ZFC+¬SCH is equiconsistent with a measurable cardinal of Mitchell order κ++\kappa^{++}κ++.
  • Todorcevic, Veličković (1980s): the proper forcing axiom implies 2ℵ0=ℵ22^{\aleph_0}=\aleph_22ℵ0​=ℵ2​.
  • Viale (2006): the proper forcing axiom implies SCH.

Setting

All cardinals live in one universe. For a cardinal κ\kappaκ, κ+\kappa^+κ+ is its successor cardinal and cf⁡κ\operatorname{cf}\kappacfκ its cofinality (the cofinality of the initial ordinal of κ\kappaκ). A cardinal is singular if it is infinite and cf⁡κ<κ\operatorname{cf}\kappa<\kappacfκ<κ, and a strong limit if κ≠0\kappa\neq0κ=0 and 2λ<κ2^\lambda<\kappa2λ<κ for all λ<κ\lambda<\kappaλ<κ.

  • GCH: 2κ=κ+2^\kappa=\kappa^+2κ=κ+ for every infinite κ\kappaκ.
  • SCH (first form, following Mitchell 1992): 2κ=κ+2^\kappa=\kappa^+2κ=κ+ for every singular strong limit κ\kappaκ.
  • SCH2_22​ (second form): for every singular κ\kappaκ, 2cf⁡κ<κ2^{\operatorname{cf}\kappa}<\kappa2cfκ<κ implies κcf⁡κ=κ+\kappa^{\operatorname{cf}\kappa}=\kappa^+κcfκ=κ+.

A normal measure on κ\kappaκ is a nonprincipal κ\kappaκ-complete ultrafilter on κ\kappaκ under which every regressive function is constant on a measure-one set; κ\kappaκ is measurable if it is uncountable and carries one. κ\kappaκ is strongly compact if it is uncountable and every proper κ\kappaκ-complete filter on any set extends to a κ\kappaκ-complete ultrafilter. The proper forcing axiom (PFA) asserts that for every proper poset PPP and every ℵ1\aleph_1ℵ1​ dense subsets of PPP there is a filter meeting all of them.

Formalization targets

Goal: Silver's reduction

(∀λ singular strong limit, cf⁡λ=ℵ0⇒2λ=λ+) ⟹ SCH.\Bigl(\forall\lambda\ \text{singular strong limit},\ \operatorname{cf}\lambda=\aleph_0\Rightarrow 2^\lambda=\lambda^+\Bigr)\ \Longrightarrow\ \mathrm{SCH}.(∀λ singular strong limit, cfλ=ℵ0​⇒2λ=λ+) ⟹ SCH.

SCH itself is independent of ZFC, so it cannot be a goal; the goal is the ZFC theorem that SCH can only first fail at countable cofinality.

Milestones (in order)

  1. GCH ⇒\Rightarrow⇒ SCH.
  2. κcf⁡κ=2κ\kappa^{\operatorname{cf}\kappa}=2^\kappaκcfκ=2κ for every singular strong limit κ\kappaκ.
  3. SCH2_22​ ⇒\Rightarrow⇒ SCH.
  4. Silver's theorem: κ\kappaκ singular, cf⁡κ>ℵ0\operatorname{cf}\kappa>\aleph_0cfκ>ℵ0​, GCH below κ\kappaκ ⇒\Rightarrow⇒ 2κ=κ+2^\kappa=\kappa^+2κ=κ+ (already published and proved on the platform as JechSetTheory.silver).
  5. The second-form reduction: SCH2_22​ at all λ\lambdaλ of countable cofinality implies SCH2_22​ everywhere (published, open, as JechSetTheory.sch_of_sch_cof_aleph0).
  6. The least counterexample to SCH has cofinality ℵ0\aleph_0ℵ0​.
  7. Scott's theorem: DDD a normal measure on κ\kappaκ and {α<κ:2∣α∣=∣α∣+}∈D\{\alpha<\kappa: 2^{|\alpha|}=|\alpha|^+\}\in D{α<κ:2∣α∣=∣α∣+}∈D imply 2κ=κ+2^\kappa=\kappa^+2κ=κ+.
  8. Solovay's theorem: if κ\kappaκ is strongly compact, then 2μ=μ+2^\mu=\mu^+2μ=μ+ for every singular strong limit μ>κ\mu>\kappaμ>κ.
  9. PFA ⇒2ℵ0=ℵ2\Rightarrow 2^{\aleph_0}=\aleph_2⇒2ℵ0​=ℵ2​.
  10. PFA ⇒\Rightarrow⇒ SCH.

Significance

The goal locates the first possible failure of SCH: any model of ¬\neg¬SCH must have a counterexample of countable cofinality, which is why Prikry-type forcing (which adds an ω\omegaω-sequence) is the standard tool for producing failures, and why ℵω\aleph_\omegaℵω​ is the canonical test case. Milestones 1, 7, 8 and 10 show that very different hypotheses (GCH, large cardinals, forcing axioms) all imply SCH, while milestone 9 shows PFA does so without GCH.

These results are classical and have complete published proofs (Jech, Set Theory, Chapters 8, 20, 24, 31). Silver's theorem is already machine-checked on this platform; the rest, including the stationary-set form of Silver's theorem needed for the goal, are not yet formalized here.

Difficulty

The naive attempt at the goal, "apply Silver's theorem at the least counterexample", fails: Silver's theorem as stated needs GCH at all infinite cardinals below κ\kappaκ, but SCH below κ\kappaκ says nothing about regular cardinals or non-strong-limits. The proof needs the stronger, stationary-set version of Silver's theorem (Galvin–Hajnal style), applied to the stationary set of singular strong limits of countable cofinality below κ\kappaκ. Scott's and Solovay's theorems need ultrapower arguments, and PFA ⇒\Rightarrow⇒ SCH needs reflection and covering-type combinatorics; none of these exist in Mathlib.

Formalization scope

  • Cardinals are Mathlib Cardinal.{u}; singularity is Cardinal.IsSingular, strong limit is Cardinal.IsStrongLimit, cofinality is κ.ord.cof, successor is Order.succ κ. The second form reuses the published predicate JechSetTheory.SCHAt.
  • A measure on κ\kappaκ is an ultrafilter on Ordinal.{u} containing the ordinals below κ\kappaκ; completeness quantifies over index types of universe u. "Proper filter" in the definition of strong compactness is Filter.NeBot; without it no cardinal would be strongly compact and Solovay's theorem would be vacuous.
  • Properness of a poset is encoded by the Gray–Shelah game characterization (names for ordinals in maximal-antichain form, forcing unfolded into predensity), because Mathlib has no forcing relation; this encoding deserves particular scrutiny. Posets are nonempty preorders with q≤pq\le pq≤p meaning qqq is stronger.
  • Not formalized: the consistency and equiconsistency results (Gitik, Prikry forcing, Silver's supercompact construction, inner-model lower bounds) and the strictness of "SCH2_22​ is stronger than SCH"; these are statements about models of ZFC rather than theorems of ZFC.
  • Reusable infrastructure welcome: clubs and stationary sets on cardinals (partly available in Def_JechStationary), ultrapowers by normal measures, Fodor's lemma, Galvin–Hajnal bounds.

Selected references

  • W. J. Mitchell, On the singular cardinal hypothesis, Trans. Amer. Math. Soc. 329(2), 507–530, 1992.
  • J. Silver, On the singular cardinals problem, Proc. ICM Vancouver 1974, 265–268, 1975.
  • T. Jech, Set Theory, 3rd millennium ed., Springer, 2003.
  • M. Gitik, The strength of the failure of the singular cardinal hypothesis, Ann. Pure Appl. Logic 51, 215–240, 1991.
  • M. Viale, The proper forcing axiom and the singular cardinal hypothesis, J. Symbolic Logic 71(2), 473–479, 2006.
  • S. Shelah, Proper and Improper Forcing, Springer, 1998.
  • Wikipedia, Singular cardinals hypothesis, revision 1349753131. https://en.wikipedia.org/w/index.php?title=Singular_cardinals_hypothesis&oldid=1349753131
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Complexity Theory·Captain: Lucas

F_d versus F_{d+1}: fixed-depth separation of bounded-depth FregeOpen Problem

Motivation

Bounded-depth Frege systems are the proof systems of propositional logic in which every formula of a proof has depth bounded by a fixed constant ddd; the resulting system is written FdF_dFd​. They are the proof-theoretic counterpart of the circuit class AC0\mathrm{AC}^0AC0 and of the bounded arithmetic theories IΔ0I\Delta_0IΔ0​ and its relatives: short FdF_dFd​-proofs of a family of tautologies correspond to proofs of the corresponding first-order statement in bounded arithmetic. Lower bounds for FdF_dFd​ are among the strongest lower bounds known for any natural proof system.

A basic structural question is whether the depth hierarchy F1⊆F2⊆F3⊆…F_1 \subseteq F_2 \subseteq F_3 \subseteq \dotsF1​⊆F2​⊆F3​⊆… is strict in a strong sense. It is known that one extra level of depth can give an exponential speed-up, but in all known separations the tautologies themselves have depth growing with ddd. J. Krajíček asked whether there is a single fixed depth kkk at which tautologies witnessing the separation of Fd+1F_{d+1}Fd+1​ from FdF_dFd​ exist for every ddd (Open Problem Garden, F_d versus F_{d+1}; [Krajíček 1995, p. 243]). The question is closely related to conservativity relations among bounded arithmetic theories.

Timeline.

  • 1987 — Buss: the pigeonhole principle PHPnn+1\mathrm{PHP}^{n+1}_nPHPnn+1​ has polynomial-size Frege proofs (of unbounded depth).
  • 1988 — Ajtai: PHPnn+1\mathrm{PHP}^{n+1}_nPHPnn+1​ has no polynomial-size FdF_dFd​-proofs for any fixed ddd.
  • 1993/1995 — Pitassi–Beame–Impagliazzo and Krajíček–Pudlák–Woods: FdF_dFd​-proofs of PHPnn+1\mathrm{PHP}^{n+1}_nPHPnn+1​ require size 2nεd2^{n^{\varepsilon_d}}2nεd​.
  • 1994 — Krajíček: for each ddd there are tautologies of depth ddd with polynomial-size Fd+1F_{d+1}Fd+1​-proofs requiring exponential-size FdF_dFd​-proofs, i.e. the problem is solved when kkk may depend on ddd (k:=dk := dk:=d).
  • 2002 — Impagliazzo–Krajíček relate the question to conservativity among bounded arithmetic theories.

Setting

Formulas. A formula is in negation normal form over variables p0,p1,…p_0, p_1, \dotsp0​,p1​,…: a literal pip_ipi​ or ¬pi\neg p_i¬pi​, a conjunction ⋀(A1,…,Am)\bigwedge(A_1,\dots,A_m)⋀(A1​,…,Am​) or a disjunction ⋁(A1,…,Am)\bigvee(A_1,\dots,A_m)⋁(A1​,…,Am​) of any finite number m≥0m \ge 0m≥0 of formulas. The depth of a literal is 000, and the depth of ⋀(A1,…,Am)\bigwedge(A_1,\dots,A_m)⋀(A1​,…,Am​) or ⋁(A1,…,Am)\bigvee(A_1,\dots,A_m)⋁(A1​,…,Am​) is 1+max⁡jdepth⁡(Aj)1+\max_j \operatorname{depth}(A_j)1+maxj​depth(Aj​) (so clauses have depth 111 and CNFs depth 222). The size ∣A∣|A|∣A∣ is the number of symbols (literals and connectives). A tautology is a formula true under every assignment. The negation A‾\overline{A}A of a formula is obtained by De Morgan's laws.

The system FdF_dFd​. Lines of a proof are sequents: finite lists Γ\GammaΓ of formulas, read as their disjunction. The rules are: (axiom) any Γ\GammaΓ containing both pip_ipi​ and ¬pi\neg p_i¬pi​; (∧\wedge∧) from A,ΓA, \GammaA,Γ for every A∈{A1,…,Am}A \in \{A_1,\dots,A_m\}A∈{A1​,…,Am​} infer ⋀(A1,…,Am),Γ\bigwedge(A_1,\dots,A_m), \Gamma⋀(A1​,…,Am​),Γ; (∨\vee∨) from Aj,ΓA_j, \GammaAj​,Γ infer ⋁(A1,…,Am),Γ\bigvee(A_1,\dots,A_m), \Gamma⋁(A1​,…,Am​),Γ; (weakening) from Γ\GammaΓ infer any Δ\DeltaΔ containing every member of Γ\GammaΓ; (cut) from A,ΓA,\GammaA,Γ and A‾,Γ\overline{A},\GammaA,Γ infer Γ\GammaΓ. An FdF_dFd​-proof of AAA is a sequence of sequents, each following from earlier ones by a rule (lines may be reused), in which every formula has depth at most ddd, and which contains the sequent consisting of AAA alone. Its size is the total size of all formulas in it.

Formalization targets

Write τ=(τm)m∈N\tau = (\tau_m)_{m\in\mathbb N}τ=(τm​)m∈N​ for a family of tautologies of depth at most kkk with ∣τm∣≥m|\tau_m| \ge m∣τm​∣≥m. It requires exponential FdF_dFd​-size if for some ε>0\varepsilon>0ε>0 and all large mmm every FdF_dFd​-proof of τm\tau_mτm​ has size at least 2∣τm∣ε2^{|\tau_m|^{\varepsilon}}2∣τm​∣ε.

Goal

∃k ∀d≥k ∃τ:τm has Fd+1-proofs of size 2C(log⁡2∣τm∣)C and τ requires exponential Fd-size.\exists k\ \forall d \ge k\ \exists \tau:\quad \tau_m \text{ has } F_{d+1}\text{-proofs of size } 2^{C(\log_2|\tau_m|)^C} \text{ and } \tau \text{ requires exponential } F_d\text{-size.}∃k ∀d≥k ∃τ:τm​ has Fd+1​-proofs of size 2C(log2​∣τm​∣)C and τ requires exponential Fd​-size.

The goal takes the weaker of the two upper bounds allowed by the problem (quasi-polynomial); the polynomial version implies it.

Milestones

  1. Soundness of FdF_dFd​: every formula with an FdF_dFd​-proof is a tautology.
  2. Completeness of FdF_dFd​ for depth-ddd formulas: every tautology of depth ≤d\le d≤d has an FdF_dFd​-proof.
  3. Exponential FdF_dFd​ lower bound for PHPnn+1\mathrm{PHP}^{n+1}_nPHPnn+1​: size ≥2nεd\ge 2^{n^{\varepsilon_d}}≥2nεd​.
  4. The depth-dependent separation: for all large ddd, depth-ddd tautologies with polynomial-size Fd+1F_{d+1}Fd+1​-proofs requiring exponential FdF_dFd​-size.

Significance

A positive answer would show that the extra power of depth d+1d+1d+1 over depth ddd is visible already on formulas of a fixed, small depth, in particular that the hierarchy of bounded-depth Frege systems does not collapse on (for example) CNF tautologies. Through the correspondence with bounded arithmetic it bears on conservativity of theories of increasing quantifier complexity over a fixed class of low-complexity statements. A negative answer would say that for low-depth tautologies the bounded-depth Frege hierarchy collapses up to quasi-polynomial size.

Status: the goal is open. Milestones 1–2 are standard; milestones 3–4 are published theorems that, to the best of current knowledge, have no machine-checked proofs. Formalizing them requires switching-lemma machinery that would be reusable across proof complexity.

Difficulty

The known lower bounds for FdF_dFd​ go through switching lemmas that simplify depth-ddd formulas under random restrictions. In the known separations of Fd+1F_{d+1}Fd+1​ from FdF_dFd​ the hard tautologies have depth growing with ddd. Keeping the formulas at one fixed depth kkk while ddd ranges over all values requires a lower-bound argument whose strength depends on the depth of the proof lines alone, uniformly in ddd; the switching-lemma arguments used for depth-dependent separations do not directly provide this.

Formalization scope

Formulas are the Lean inductive type FdVersusFdSucc.Formula over variables ℕ, with conjunctions and disjunctions of lists (so and [] is true and or [] is false). Depth counts literals as 000. FdF_dFd​ is the dag-like one-sided sequent calculus described above; sequents are lists, weakening is by list inclusion. Proof size counts symbols of all formulas in all lines. Exponential lower bounds are asked only for all sufficiently large indices. The goal quantifies over d≥kd \ge kd≥k: for d+1<kd+1 < kd+1<k no Fd+1F_{d+1}Fd+1​-proof of a depth-kkk formula exists, so "every ddd" cannot be meant literally. A family whose formulas have depth >d> d>d would satisfy the lower bound vacuously; this is excluded because the formulas must have depth at most k≤dk \le dk≤d. The size condition ∣τm∣≥m|\tau_m| \ge m∣τm​∣≥m rules out finite families.

Contributions welcome: proofs of soundness and completeness, a formalized switching lemma, and the classical lower bounds.

Selected references

  • J. Krajíček, Lower bounds to the size of constant-depth propositional proofs, J. Symbolic Logic 59(1), 1994, 73–86.
  • J. Krajíček, Bounded arithmetic, propositional logic, and complexity theory, Encyclopedia of Mathematics and Its Applications 60, Cambridge University Press, 1995 (p. 243).
  • Open Problem Garden, F_d versus F_{d+1}, http://www.openproblemgarden.org/op/f_d_versus_f_d_1
  • M. Ajtai, The complexity of the pigeonhole principle, FOCS 1988.
  • J. Krajíček, P. Pudlák, A. Woods, An exponential lower bound to the size of bounded depth Frege proofs of the pigeonhole principle, Random Structures & Algorithms 7(1), 1995.
  • T. Pitassi, P. Beame, R. Impagliazzo, Exponential lower bounds for the pigeonhole principle, Computational Complexity 3, 1993.
  • S. Buss, Polynomial size proofs of the propositional pigeonhole principle, J. Symbolic Logic 52, 1987.
  • R. Impagliazzo, J. Krajíček, A note on conservativity relations among bounded arithmetic theories, Math. Logic Quarterly 48, 2002.
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Theory of Computation·Captain: mikedeng1

On the (Im)possibility of Obfuscating Programs 2: Without the Size 1^|M| a Simulator with Oracle Access to ⟨M⟩ Cannot Decide Every Decidable Promise Problem Closed under Functional EquivalenceResearch Paper

Motivation

Barak, Goldreich, Impagliazzo, Rudich, Sahai, Vadhan and Yang (J. ACM 59(2), 2012) proved that virtual-black-box obfuscation is impossible in general: there are programs from which any obfuscated copy leaks something that oracle access to the program does not. The definition being refuted asks that whatever can be computed from the obfuscated code can also be computed by a simulator that only has oracle access to the program and is told an upper bound on its size.

Appendix A of the paper studies the computability analogue of this definition, where running time is unbounded. Rice's theorem says that no non-trivial property of the function computed by a Turing machine is decidable. The appendix asks a relative version: if a property of programs depends only on the computed function and is decidable from the code on a promise, can it also be decided from black-box access alone? Theorem A.2 answers yes when the simulator is given the size 1∣M∣1^{|M|}1∣M∣. Theorem A.4, the goal of this mission, shows that the size cannot be dropped. The result says which part of the simulator's interface is essential: an upper bound on program length is needed already in the computability setting, not only to pay for running time.

Setting

A machine MMM is a one-tape Turing machine over the tape alphabet {0,1}\{0,1\}{0,1}, with 000 the blank, and with states 0,1,…,q0,1,\dots,q0,1,…,q; state 000 is initial. In state sss reading symbol aaa the machine either halts or moves to a new state and performs one action (move left, move right, or write a symbol). One such transition is one step. The size ∣M∣=q+1|M|=q+1∣M∣=q+1 is the number of states.

Inputs are natural numbers in unary: input xxx is the tape 1x1^x1x, so ∣x∣=x|x|=x∣x∣=x. When MMM halts, its output is the number of consecutive 111s starting at the head. The computed function [M]:N⇀N[M]:\mathbb N\rightharpoonup\mathbb N[M]:N⇀N is defined where MMM halts. The step-bounded function is

⟨M⟩(1t,x)={yif M(x) halts with output y after at most t steps,⊥otherwise.\langle M\rangle(1^t,x)=\begin{cases}y&\text{if } M(x)\text{ halts with output } y \text{ after at most } t \text{ steps},\\ \bot&\text{otherwise.}\end{cases}⟨M⟩(1t,x)={y⊥​if M(x) halts with output y after at most t steps,otherwise.​

A promise problem Π=(ΠY,ΠN)\Pi=(\Pi_Y,\Pi_N)Π=(ΠY​,ΠN​) is a pair of disjoint sets of machines. It is closed under [⋅][\cdot][⋅] if [M]=[M′][M]=[M'][M]=[M′] implies M∈ΠY  ⟺  M′∈ΠYM\in\Pi_Y\iff M'\in\Pi_YM∈ΠY​⟺M′∈ΠY​ and M∈ΠN  ⟺  M′∈ΠNM\in\Pi_N\iff M'\in\Pi_NM∈ΠN​⟺M′∈ΠN​. It is decidable if a program TTT reading the description ⌜M⌝\ulcorner M\urcorner┌M┐ outputs 111 on ΠY\Pi_YΠY​ and 000 on ΠN\Pi_NΠN​; outside ΠY∪ΠN\Pi_Y\cup\Pi_NΠY​∪ΠN​, TTT may diverge.

An oracle machine SSS may query ⟨M⟩\langle M\rangle⟨M⟩: it asks (t,x)(t,x)(t,x) and receives ⟨M⟩(1t,x)\langle M\rangle(1^t,x)⟨M⟩(1t,x). S⟨M⟩()S^{\langle M\rangle}()S⟨M⟩() is its output with this oracle and no input. In particular SSS does not know ∣M∣|M|∣M∣.

Conjecture A.3 (Rice's theorem, third generalization) asserts: for every promise problem Π\PiΠ closed under [⋅][\cdot][⋅] and decidable, there is an oracle machine SSS with

M∈ΠY⇒S⟨M⟩()=1,M∈ΠN⇒S⟨M⟩()=0.M\in\Pi_Y\Rightarrow S^{\langle M\rangle}()=1,\qquad M\in\Pi_N\Rightarrow S^{\langle M\rangle}()=0.M∈ΠY​⇒S⟨M⟩()=1,M∈ΠN​⇒S⟨M⟩()=0.

The proof uses the description complexity of a partial function, KC(f)=min⁡{∣M∣:[M]=f}\mathrm{KC}(f)=\min\{|M|:[M]=f\}KC(f)=min{∣M∣:[M]=f}.

Formalization targets

Goal: Theorem A.4

¬ Conjecture A.3.\neg\,\text{Conjecture A.3}.¬Conjecture A.3.

There is a decidable promise problem closed under [⋅][\cdot][⋅] that no oracle machine without the size can decide. The witness is

ΠY={M:M always halts and ∃x<KC([M]), [M](x)=1},ΠN={M:M always halts and [M]≡0}.\Pi_Y=\{M: M\text{ always halts and } \exists x<\mathrm{KC}([M]),\ [M](x)=1\},\qquad \Pi_N=\{M: M\text{ always halts and } [M]\equiv 0\}.ΠY​={M:M always halts and ∃x<KC([M]), [M](x)=1},ΠN​={M:M always halts and [M]≡0}.

Milestones, in the order of the paper's proof

  1. Π\PiΠ is closed under [⋅][\cdot][⋅].
  2. KC([M])≤∣M∣\mathrm{KC}([M])\le|M|KC([M])≤∣M∣.
  3. Π\PiΠ is decidable.
  4. The machine ZZZ that reads its input and returns 000 satisfies ⟨Z⟩(1t,x)=⊥\langle Z\rangle(1^t,x)=\bot⟨Z⟩(1t,x)=⊥ for t<∣x∣t<|x|t<∣x∣ and 000 otherwise, and Z∈ΠNZ\in\Pi_NZ∈ΠN​.
  5. If S⟨M0⟩()=σS^{\langle M_0\rangle}()=\sigmaS⟨M0​⟩()=σ, there is nnn such that every MMM agreeing with M0M_0M0​ on all ⟨⋅⟩(1t,x)\langle\cdot\rangle(1^t,x)⟨⋅⟩(1t,x) with t,∣x∣≤nt,|x|\le nt,∣x∣≤n also has S⟨M⟩()=σS^{\langle M\rangle}()=\sigmaS⟨M⟩()=σ.
  6. For all n,rn,rn,r some machine computes Nn,r(x)=0,1,rN_{n,r}(x)=0,1,rNn,r​(x)=0,1,r (for ∣x∣≤n|x|\le n∣x∣≤n, ∣x∣=n+1|x|=n+1∣x∣=n+1, ∣x∣≥n+2|x|\ge n+2∣x∣≥n+2) and has the same step profile as ZZZ on all ∣x∣≤n|x|\le n∣x∣≤n.
  7. For every nnn some rrr has KC(Nn,r)>n+1\mathrm{KC}(N_{n,r})>n+1KC(Nn,r​)>n+1, so machines computing Nn,rN_{n,r}Nn,r​ lie in ΠY\Pi_YΠY​.

Significance

The result. Theorem A.4 separates two interfaces for a simulator. With 1∣M∣1^{|M|}1∣M∣, every decidable promise problem closed under [⋅][\cdot][⋅] is decided from black-box access (Theorem A.2). Without it, some are not. The size bound in the virtual-black-box definition is therefore doing real work even when complexity is ignored, and any computability-theoretic notion of obfuscation must hand the simulator a bound on the program's length.

Formalizing it. The theorem is proved in the paper in a few lines, and the proof rests on facts the paper does not argue: that a halting oracle computation reads only finitely many oracle values, that a machine with a prescribed function and a prescribed running time on short inputs exists, and that some Nn,rN_{n,r}Nn,r​ has description complexity above n+1n+1n+1. The paper says only "Let r be a string of Kolmogorov complexity 2n"; the counting argument behind it is left implicit. No machine-checked proof of Theorem A.4 or of Theorem A.2 is known to exist. This mission produces one in a model with literal step counts, together with reusable statements about step-bounded Turing machines and the use principle for oracle programs.

Difficulty

The obvious attempt at Conjecture A.3 is to search for a machine consistent with the observed values of ⟨M⟩\langle M\rangle⟨M⟩ and run TTT on it. Without a size bound the search space is infinite, and a finite prefix of ⟨M⟩\langle M\rangle⟨M⟩ never pins down [M][M][M]. The difficulty of the formal proof is elsewhere. It has to build Nn,rN_{n,r}Nn,r​ as an actual Turing machine with an exact step count on short inputs, and to prove that its description complexity exceeds n+1n+1n+1. Both need the step semantics to be exact. A fuel-based semantics in which larger programs need more fuel on the same input does not support the construction of Nn,rN_{n,r}Nn,r​.

Formalization scope

Machines are Mathlib's Post–Turing machines Turing.TM0.Machine Bool (Fin (q+1)), with the blank false and initial state 0. [M][M][M] is TM0.eval followed by the head read-out; ⟨M⟩(1t,x)\langle M\rangle(1^t,x)⟨M⟩(1t,x) iterates TM0.step at most ttt times. ∣M∣|M|∣M∣ is the number of states, so there are finitely many machines of each size. Inputs are unary, ∣x∣=x|x|=x∣x∣=x. This is forced, because the proof compares inputs both with numbers (x<KC([M])x<\mathrm{KC}([M])x<KC([M]), Nn,r(n+1)=1N_{n,r}(n+1)=1Nn,r​(n+1)=1) and with lengths (∣x∣≤n|x|\le n∣x∣≤n, t<∣x∣t<|x|t<∣x∣). The paper's string rrr is a natural number. The description ⌜M⌝\ulcorner M\urcorner┌M┐ is the encoding of qqq and the transition table, and the decider TTT is a Mathlib Nat.Partrec.Code acting on it. KC\mathrm{KC}KC is sInf over N\mathbb NN, which would be 000 for a function no machine computes; that case does not arise for f=[M]f=[M]f=[M], and milestone 7 asserts KC>n+1\mathrm{KC}>n+1KC>n+1 outright.

Oracle machines are the published uniform oracle programs FriedbergMuchnik.Program with semantics oracleEval. A query (t,x)(t,x)(t,x) is coded Nat.pair t x, and the answers ⊥\bot⊥ and yyy are coded 000 and y+1y+1y+1. S⟨M⟩()S^{\langle M\rangle}()S⟨M⟩() runs SSS on the fixed input 000.

The simulator is a single program chosen before MMM, and it only sees the values of ⟨M⟩\langle M\rangle⟨M⟩ it queries. A formalization in which SSS is an arbitrary function of the whole oracle, or receives ∣M∣|M|∣M∣, would make the goal false, and one in which the oracle is [M][M][M] changes the statement. Both are excluded here.

Needed infrastructure: computability of simulating a TM0 machine from its description (for milestone 3), the use principle for oracleEval with a total oracle (milestone 5), explicit construction and step analysis of ZZZ and Nn,rN_{n,r}Nn,r​ (milestones 4 and 6), and a counting bound on machines of bounded size (milestone 7). The first two are reusable beyond this mission. Proofs of individual milestones are welcome independently.

Selected references

  • B. Barak, O. Goldreich, R. Impagliazzo, S. Rudich, A. Sahai, S. Vadhan, K. Yang, On the (Im)possibility of Obfuscating Programs, Journal of the ACM 59(2), 2012. https://doi.org/10.1145/2160158.2160159 (Appendix A, Conjecture A.3 and Theorem A.4, p. A:42 of the author's copy).
  • H. G. Rice, Classes of recursively enumerable sets and their decision problems, Transactions of the AMS 74(2), 1953. https://doi.org/10.1090/S0002-9947-1953-0053041-6
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Theory of Computation·Captain: mikedeng1

On the (Im)possibility of Obfuscating Programs 1: A Decidable Promise Problem Closed under Functional Equivalence Is Decided by a Simulator Given Only Oracle Access to ⟨M⟩ and the Size 1^|M|Research Paper

Motivation

Program obfuscation asks for a compiler that turns a program into one computing the same function while revealing nothing beyond what black-box access to the function reveals. Barak, Goldreich, Impagliazzo, Rudich, Sahai, Vadhan and Yang (J. ACM 59(2), 2012; conference version CRYPTO 2001) showed that this "virtual black box" notion cannot be achieved in general. Part of their motivation is an informal reading of Rice's Theorem: since every non-trivial semantic property of programs is undecidable from the code, "the only useful thing you can do with a machine is run it". Section 5 of the paper asks whether a complexity-theoretic analogue holds, and Appendix A asks what the right computability-theoretic statement is once one passes from languages to promise problems.

Classical Rice's Theorem (Rice, 1953) concerns languages: a decidable set of machines closed under functional equivalence is empty or everything. For promise problems the naive analogue (Conjecture A.1 of the paper) is false: the problem "halts everywhere and outputs 1 on 0" versus "halts everywhere and outputs 0 on 0" is decidable, non-trivial, and closed under functional equivalence. Appendix A replaces "trivial" by "decidable by a simulator that only runs the machine", proves that version (Theorem A.2), and shows that the size bound given to the simulator cannot be dropped (Theorem A.4, a separate mission in this series). This mission formalizes Theorem A.2.

Setting

A machine MMM is a program for a partial recursive function, identified with its description, a natural number ⌜M⌝\ulcorner M\urcorner┌M┐. Its computed function [M]:N⇀N[M]:\mathbb N\rightharpoonup\mathbb N[M]:N⇀N is possibly partial, and [M]≡[M′][M]\equiv[M'][M]≡[M′] means equality of partial functions. The size ∣M∣|M|∣M∣ is the length of the binary description of MMM; there are at most 2s2^s2s machines of size sss.

The bounded run of MMM is

⟨M⟩(1t,x)={yif M(x) halts with output y within budget t,⊥otherwise.\langle M\rangle(1^t,x)=\begin{cases} y & \text{if } M(x)\text{ halts with output } y \text{ within budget } t,\\ \bot&\text{otherwise.}\end{cases}⟨M⟩(1t,x)={y⊥​if M(x) halts with output y within budget t,otherwise.​

An oracle machine SSS may query ⟨M⟩\langle M\rangle⟨M⟩ at any pair (t,x)(t,x)(t,x) of its choice; S⟨M⟩(z)S^{\langle M\rangle}(z)S⟨M⟩(z) is its output on input zzz. Oracle access to ⟨M⟩\langle M\rangle⟨M⟩ is what running MMM provides: one may choose inputs and budgets, but one never sees the code.

A promise problem Π=(ΠY,ΠN)\Pi=(\Pi_Y,\Pi_N)Π=(ΠY​,ΠN​) is a pair of disjoint sets of machines. A machine TTT decides Π\PiΠ if T(⌜M⌝)=1T(\ulcorner M\urcorner)=1T(┌M┐)=1 for M∈ΠYM\in\Pi_YM∈ΠY​ and T(⌜M⌝)=0T(\ulcorner M\urcorner)=0T(┌M┐)=0 for M∈ΠNM\in\Pi_NM∈ΠN​; nothing is asked of TTT outside ΠY∪ΠN\Pi_Y\cup\Pi_NΠY​∪ΠN​. Π\PiΠ is closed under [⋅][\cdot][⋅] if [M]≡[M′][M]\equiv[M'][M]≡[M′] implies M∈ΠY⇔M′∈ΠYM\in\Pi_Y\Leftrightarrow M'\in\Pi_YM∈ΠY​⇔M′∈ΠY​ and M∈ΠN⇔M′∈ΠNM\in\Pi_N\Leftrightarrow M'\in\Pi_NM∈ΠN​⇔M′∈ΠN​.

A machine NNN is nnn-compatible with MMM if ⟨N⟩(1t,x)=⟨M⟩(1t,x)\langle N\rangle(1^t,x)=\langle M\rangle(1^t,x)⟨N⟩(1t,x)=⟨M⟩(1t,x) for all t≤nt\le nt≤n and all xxx with ∣x∣≤n|x|\le n∣x∣≤n, and SnS_nSn​ denotes the set of machines of size ∣M∣|M|∣M∣ that are nnn-compatible with MMM.

Formalization targets

Goal: Theorem A.2 (p. A:41)

If Π\PiΠ is closed under [⋅][\cdot][⋅] and decidable, then there is one oracle machine SSS such that for every machine MMM

M∈ΠY⇒S⟨M⟩(1∣M∣)=1,M∈ΠN⇒S⟨M⟩(1∣M∣)=0.M\in\Pi_Y\Rightarrow S^{\langle M\rangle}(1^{|M|})=1,\qquad M\in\Pi_N\Rightarrow S^{\langle M\rangle}(1^{|M|})=0 .M∈ΠY​⇒S⟨M⟩(1∣M∣)=1,M∈ΠN​⇒S⟨M⟩(1∣M∣)=0.

Milestones (proof of Theorem A.2, pp. A:41–A:42)

  1. Fact (1): nnn-compatibility with MMM is decidable by one oracle machine with oracle ⟨M⟩\langle M\rangle⟨M⟩, uniformly in MMM, NNN, nnn.
  2. Fact (3): if [M]≢[N][M]\not\equiv[N][M]≡[N], there is n′n'n′ with NNN not nnn-compatible with MMM for every n>n′n>n'n>n′.
  3. There is n′′n''n′′ such that every N∈SnN\in S_nN∈Sn​ with n>n′′n>n''n>n′′ satisfies [N]≡[M][N]\equiv[M][N]≡[M].
  4. If some stage nnn finds TTT, run for nnn steps, halting on all of SnS_nSn​ with a common answer σ\sigmaσ, then σ=T(⌜M⌝)\sigma=T(\ulcorner M\urcorner)σ=T(┌M┐).
  5. For M∈ΠYM\in\Pi_YM∈ΠY​ (resp. ΠN\Pi_NΠN​) such a stage exists, with answer 111 (resp. 000).

Significance

The result. Theorem A.2 is the precise sense in which a semantic property of programs, decidable from the code on a promise, is decidable without the code: running the program on chosen inputs for chosen budgets, plus a bound on its length, suffices. It is the computability-theoretic counterpart of Conjecture 5.2, the complexity-theoretic statement whose failure (Theorem 5.3, under cryptographic assumptions) is one face of the impossibility of obfuscation. Together with Theorem A.4 it isolates the size bound as the ingredient that makes the generalization true.

Formalizing it. The theorem has a complete proof in the paper, about one page long. No machine-checked version is known: the Prove2Me catalogue has classical Rice's Theorem for index sets of Nat.Partrec.Code (FamousTheorems.rice_theorem, which this proof does not use) but nothing on promise problems, bounded-run oracles, or simulators. A formal proof produces reusable infrastructure: the oracle ⟨M⟩\langle M\rangle⟨M⟩ as a total function on query codes, decidability of compatibility relative to it, and a dovetailing simulator assembled as a uniform oracle program.

Difficulty

Each mathematical step is elementary; the work lies in the effective content. The obvious idea, "query ⟨M⟩\langle M\rangle⟨M⟩ until [M][M][M] is known and then decide", fails: [M][M][M] is an infinite object, and no finite set of queries pins it down. The paper's simulator instead uses the size bound to restrict attention to finitely many candidate codes, and needs TTT to agree on every surviving candidate at a common budget. Turning this into one oracle program requires enumerating all codes of a given size, deciding compatibility from finitely many oracle answers, comparing bounded runs of TTT, and searching over nnn, all uniformly in MMM and with no access to ⌜M⌝\ulcorner M\urcorner┌M┐. Termination holds only on promise instances, since TTT may diverge elsewhere, so the argument cannot assume a total decider.

Formalization scope

Machines are Mathlib's Nat.Partrec.Code; descriptions are Encodable.encode; [M][M][M] is Code.eval; ⟨M⟩(1t,x)\langle M\rangle(1^t,x)⟨M⟩(1t,x) is Code.evaln t M x : Option ℕ, with none for ⊥\bot⊥. The budget of evaln is fuel rather than a literal Turing-machine step count. It is monotone in ttt, converges to [M][M][M], and is primitive recursive in (t,M,x)(t,M,x)(t,M,x), and these are the only properties of step counts that the proof uses. Inputs are natural numbers, ∣x∣|x|∣x∣ is Nat.size x, and ∣M∣|M|∣M∣ is Nat.size (encode M). Oracle machines are the published uniform oracle programs FriedbergMuchnik.Program with semantics oracleEval (a single syntax tree with an oracle-query instruction, equivalent to relative partial recursiveness). The oracle ⟨M⟩\langle M\rangle⟨M⟩ is the total function sending Nat.pair t x to 0 for ⊥\bot⊥ and to y + 1 for an output yyy. The simulator's input 1∣M∣1^{|M|}1∣M∣ is the number ∣M∣|M|∣M∣, since unary versus binary is immaterial here. "S⟨M⟩(1∣M∣)=1S^{\langle M\rangle}(1^{|M|})=1S⟨M⟩(1∣M∣)=1" means that the run halts with output 111. Promise problems carry their disjointness, as in the paper, although no statement uses it. The decider TTT is a Code applied to descriptions and may diverge off the promise.

The trivializing readings are excluded. SSS is chosen before MMM (one program, not one per machine). It receives the size of MMM, never its description, and many machines share each size. It sees only the finitely many oracle answers a halting run queries. The oracle is the bounded-run function ⟨M⟩\langle M\rangle⟨M⟩, not [M][M][M] or the code.

Contributions welcome: proofs of the milestones; a primitive-recursive enumeration of the codes of a given size; a proof that oracle M is uniformly computable from M, which is useful beyond this mission.

Selected references

  • B. Barak, O. Goldreich, R. Impagliazzo, S. Rudich, A. Sahai, S. Vadhan, K. Yang, On the (Im)possibility of Obfuscating Programs, Journal of the ACM 59(2), 2012. https://doi.org/10.1145/2160158.2160159
  • H. G. Rice, Classes of recursively enumerable sets and their decision problems, Transactions of the AMS 74(2), 1953. https://doi.org/10.1090/S0002-9947-1953-0053041-6
  • S. Even, A. L. Selman, Y. Yacobi, The complexity of promise problems with applications to public-key cryptography, Information and Control 61(2), 1984. https://doi.org/10.1016/S0019-9958(84)80056-X
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Functional Analysis·Captain: marwahaha

Relative independence of the separable quotient problemOpen Problem

Motivation

The separable quotient problem asks whether every infinite-dimensional Banach space admits a separable infinite-dimensional quotient. The selected statement isolates a negative result under the continuum hypothesis. The pinned manuscript supplies the research context.

Setting

A quotient is encoded as a surjective bounded linear map between Banach spaces over the same scalar field. The target considers both real and complex scalars.

Formalization target

The selected goal is OAI.SeparableQuotient.negative_main. Its central assertion is

CH⟹¬SQ(R)∧¬SQ(C).\mathrm{CH}\quad\Longrightarrow\quad\neg\mathrm{SQ}(\mathbb R)\land\neg\mathrm{SQ}(\mathbb C).CH⟹¬SQ(R)∧¬SQ(C).

The theorem states that, under the continuum hypothesis CH (the defined proposition that the cardinality of ℝ equals ℵ₁), the separable quotient assertion fails over both the real and the complex numbers, at any universe level u. Here SQ over a scalar field 𝕜 (ℝ or ℂ) asserts that every complete normed space X over 𝕜 in Type u that is not finite-dimensional has a separable quotient in the following sense: there exist a complete, separable, infinite-dimensional normed space Y over 𝕜, also in Type u, and a bounded 𝕜-linear map T from X to Y that is surjective. The conclusion is the conjunction of ¬SQ(ℝ) and ¬SQ(ℂ), so for each field there is an infinite-dimensional Banach space with no such bounded surjection onto an infinite-dimensional separable Banach space. The statement is admitted without proof in the source.

Significance and status

CH is an explicit assumption. The published goal does not state relative consistency, a positive model, or the complete independence assertion in the manuscript title. The target is currently Open on Prove2Me. The manuscript's mathematical argument and a machine-checked proof of the selected statement are separate deliverables.

Difficulty

The obstruction must rule out every bounded surjection onto every separable infinite-dimensional target space, rather than one proposed quotient.

Formalization scope

The exact published goal, its hypotheses and its referenced definition blocks specify the requested formalization. The explanatory formula above is a summary; all quantifiers and additional clauses in the linked statement remain required.

The available published material supplies this goal and its necessary definitions. Additional manuscript lemmas are not represented as attached milestones.

Selected references

  • OpenAI, Relative independence of the separable quotient problem, preprint, 2026. Manuscript.
  • OpenAI, accompanying formal statements, commit adc7f1241b42. Selected goal source.
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Algebra·Captain: marwahaha

Finite congruence lattices: characterization and undecidabilityOpen Problem

Motivation

A finite algebra has a lattice of congruences, ordered by inclusion. The source studies which finite lattices arise in this way, and the available goal captures its colored-graph characterization. The source is OpenAI's September 2026 manuscript.

Setting

A congruence is an equivalence relation preserved by every basic operation of an algebra. A graph witness assigns lattice-valued colors to pairs of finitely many vertices, with symmetry, diagonal, triangle, separation, and seed-connectivity conditions.

Formalization targets

Representable⁡(L)⟺HasGraphWitness⁡(L).\operatorname{Representable}(L)\quad\Longleftrightarrow\quad\operatorname{HasGraphWitness}(L).Representable(L)⟺HasGraphWitness(L).

For any finite type Lat carrying a lattice structure with a least element ⊥, Lat is Representable if and only if it HasGraphWitness. Representable means that there is a positive size n and an algebra on Fin n (a finite list of operations, each of some finite arity, acting on the carrier) such that Lat is order-isomorphic to the poset of congruences of that algebra, where a congruence is an equivalence relation compatible with every operation: if corresponding arguments are related coordinatewise, the outputs are related; congruences are ordered by the inherited order on equivalence relations. HasGraphWitness means that there is a positive size n and a Lat-valued coloring color of ordered pairs from Fin n satisfying five conditions: color is symmetric; color(a,b)=⊥ exactly when a=b; color(a,b) ≤ color(a,m) ⊔ color(m,b) for all a, b, m; whenever lower ≰ upper in Lat, some pair has color ≤ lower but not ≤ upper; and for every nonempty finite set of seed pairs, any pair whose color is at most the join of the seed colors is connected to the seeds. Connected is the equivalence closure of the relation Marked, where a,b are Marked if some self-map p of the vertices satisfies color(p x,p y) ≤ color(x,y) for all x,y and sends some seed pair (s,t) to (a,b).

The goal is OAI.FiniteCongruence.graph_criterion.

Significance

The equivalence provides a finite combinatorial certificate for representation as an entire congruence lattice. It is the characterization component of the manuscript. The manuscript presents a mathematical argument; the attached goal currently has status Open, so the requested contribution is a Lean proof of the displayed formal target.

Difficulty

Representing a lattice inside a congruence lattice is weaker than obtaining all congruences. The witness’s seed-connectivity condition must capture this fullness and work in both directions.

Formalization scope

The lattice is finite and has a least element. The algebra carrier and graph witness carrier each have positive finite size. Operations have finite arities, and representation is an order isomorphism. The manuscript’s undecidability conclusions are not among the attached targets.

The shared definitions are supplied by FiniteCongruenceGraph. The target uses Lean 4.33.1 with the pinned Mathlib environment. Contributions should preserve the existing statement and develop reusable supporting lemmas in its definitions.

Selected references

  • OpenAI, Finite congruence lattices: characterization and undecidability, preprint, September 2026. Manuscript.
  • OpenAI, accompanying Lean statement. Pinned source.
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AlgebraGroup Theory·Captain: wurtle

Finite algebraic envelopes and the Boone–Higman conjectureResearch Paper

Motivation

The word problem for a group with a finite generating set asks whether a given word in the generators and their inverses represents the identity; it is decidable if one algorithm answers this for every word. Since Higman's embedding theorem (1961), a recurring theme in combinatorial group theory has been to characterize algorithmic properties of groups by embeddings into groups with good algebraic properties. Boone and Higman (1974) proved that a finitely generated group has decidable word problem if and only if it embeds in a simple subgroup of a finitely presented group, and asked whether the simple group can itself be taken finitely presented. This is the Boone–Higman conjecture.

This preprint states a proof of the conjecture: a finitely generated group has decidable word problem if and only if it embeds in a finitely presented simple group.

Background

  • 1961. Higman characterizes finitely generated subgroups of finitely presented groups as the recursively presented groups.
  • 1974. Boone and Higman prove the characterization G≤H≤KG \le H \le KG≤H≤K with HHH simple and KKK finitely presented, and ask whether HHH can be finitely presented.
  • 1980. Thompson embeds every finitely generated group with decidable word problem into a finitely generated simple group with decidable word problem; a later proof is given by Darbinyan and Steenbock (2022).
  • 2022. Belk and Zaremsky introduce twisted Brin–Thompson groups, which embed an acting group in a simple group; Zaremsky (2024) gives a finite-presentation criterion requiring a finitely presented acting group with finitely generated point stabilizers and finitely many orbits on unordered pairs.
  • 2023–2026. The conjecture is verified for hyperbolic groups and contracting self-similar groups (Belk, Bleak, Matucci, Zaremsky), Baumslag–Solitar and free-by-cyclic groups (Bux, Llosa Isenrich, Wu), and Aut(Fn)\mathrm{Aut}(F_n)Aut(Fn​) and mapping class groups of punctured surfaces (Belk, Fournier-Facio, Hyde, Zaremsky); see the survey of Belk, Bleak, Matucci and Zaremsky.
  • 2026. The OpenAI preprint Finite algebraic envelopes and the Boone–Higman conjecture (dated September 23, 2026) states a proof of the conjecture in general. It has not been peer reviewed; the Lean goal is open on this platform.

Setting

A group GGG is finitely generated if some finite tuple g1,…,gng_1, \dots, g_ng1​,…,gn​ generates it. A word is a finite list of letters gi±1g_i^{\pm1}gi±1​; GGG has decidable word problem if for some finite generating tuple the set of words evaluating to 111 is computable. A group is finitely presented if it has a presentation with finitely many generators and relations, and simple if it is nontrivial with no normal subgroups other than 111 and itself. An embedding is an injective homomorphism.

In Lean (OAI.FiniteAlgebraicEnvelopes, definitions BooneHigman), words are List (Fin n × Bool), evalWord multiplies generators or their inverses, HasDecidableWordProblem G asserts a generating tuple whose identity predicate is ComputablePred, and EmbedsInFinitelyPresentedSimpleGroup G asserts a group in the same universe that is Group.IsFinitelyPresented and IsSimpleGroup, with an injective homomorphism from GGG.

Formalization targets

Goal: Theorem 1.1 (Boone–Higman conjecture)

For every finitely generated group GGG,

G has decidable word problem  ⟺  G embeds in a finitely presented simple group.G \text{ has decidable word problem} \iff G \text{ embeds in a finitely presented simple group.}G has decidable word problem⟺G embeds in a finitely presented simple group.

The implication from right to left is classical (a finitely presented simple group has decidable word problem, and the property passes to finitely generated subgroups); the content is the forward implication.

Significance

The result. Theorem 1.1 gives a purely algebraic characterization of decidability of the word problem and settles a question open since 1974. Its intermediate result, Theorem 1.2, embeds every finitely generated group with decidable word problem in a finitely presented group with a faithful two-transitive action on a countably infinite set with finitely generated point stabilizers, which feeds Zaremsky's criterion for twisted Brin–Thompson groups.

Formalizing it. No machine-checked proof exists, and the source is an unrefereed preprint. Formalizing it requires computability of word problems, finite presentations (available in Mathlib as Group.IsFinitelyPresented), Steinberg groups over noncommutative rings, and twisted Brin–Thompson groups, which are not in Mathlib. The proof also uses Zaremsky's finite-presentation theorem (Theorem 2.3) from the literature, which a complete formalization must include.

Difficulty

Finite presentation does not pass to subgroups, so Boone and Higman's simple group HHH and finitely presented group KKK cannot simply be merged. Thompson's construction gives a finitely generated simple overgroup but provides no finite presentation. The new route needs a finitely presented group acting faithfully and two-transitively, built from a finitely presented algebra whose simplicity appears only after passing to a direct limit, and an exact calculation removing the kernel of the Steinberg group without assuming it is central.

Formalization scope

  • GroupType : Type u with [Group GroupType] [Group.FG GroupType]; the target group lives in the same universe.
  • Decidability is ComputablePred on words over a specified finite generating tuple (Subgroup.closure (range generators) = ⊤); since the word problem's decidability is independent of the finite generating set, this matches the source.
  • IsSimpleGroup includes nontriviality, matching the paper's convention.

The statement is not trivially true: both directions are required for every finitely generated group, including infinitely presented ones.

Needed infrastructure: computable group presentations, Steinberg groups, mapping tori of endomorphisms, twisted Brin–Thompson groups and Zaremsky's criterion. Contributions toward the classical backward implication, or toward Theorem 1.2, are welcome.

Selected references

  • OpenAI, Finite algebraic envelopes and the Boone–Higman conjecture, OpenAI Math Release preprint, September 23, 2026. https://github.com/openai/math/blob/adc7f1241b42e322a6451854ab7e4b4c146bf78a/preprints/Finite-algebraic-envelopes-and-the-Boone-Higman-conjecture-September-23-2026/paper.pdf
  • W. W. Boone, G. Higman, An algebraic characterization of groups with soluble word problem, J. Austral. Math. Soc. 18 (1974). https://doi.org/10.1017/S1446788700019108
  • G. Higman, Subgroups of finitely presented groups, Proc. Roy. Soc. London Ser. A 262 (1961). https://doi.org/10.1098/rspa.1961.0132
  • R. J. Thompson, Embeddings into finitely generated simple groups which preserve the word problem, in Word Problems II, North-Holland, 1980. https://doi.org/10.1016/S0049-237X(08)71348-X
  • J. Belk, M. C. B. Zaremsky, Twisted Brin–Thompson groups, Geom. Topol. 26 (2022). https://doi.org/10.2140/gt.2022.26.1189
  • M. C. B. Zaremsky, Finite presentability of twisted Brin–Thompson groups, Proc. Roy. Soc. Edinburgh Sect. A (2024). https://doi.org/10.1017/prm.2024.131
  • J. Belk, C. Bleak, F. Matucci, M. C. B. Zaremsky, Hyperbolic groups satisfy the Boone–Higman conjecture, Duke Math. J. 175 (2026). https://doi.org/10.1215/00127094-2025-0055
  • J. Belk, C. Bleak, F. Matucci, M. C. B. Zaremsky, Progress around the Boone–Higman conjecture, 2025. https://arxiv.org/abs/2306.16356
  • K.-U. Bux, C. Llosa Isenrich, X. Wu, On the Boone–Higman conjecture for groups acting on locally finite trees, 2025. https://arxiv.org/abs/2408.05673
  • J. Belk, F. Fournier-Facio, J. Hyde, M. C. B. Zaremsky, Boone–Higman embeddings of Aut(F_n) and mapping class groups of punctured surfaces, 2026. https://arxiv.org/abs/2503.21882
  • A. Darbinyan, M. Steenbock, Embeddings into left-orderable simple groups, J. London Math. Soc. 105 (2022). https://doi.org/10.1112/jlms.12552
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Complexity TheoryTheoretical Computer Science·Captain: wurtle

Choiceless polynomial time with counting does not capture polynomial timeResearch Paper

Motivation

Descriptive complexity asks whether complexity classes can be characterized by logics, without reference to machines. Immerman and Vardi (1982) showed that on ordered finite structures, first-order logic with a least fixed-point operator expresses exactly the polynomial-time queries. On unordered structures, such as graphs given only up to isomorphism, an algorithm must not depend on an arbitrary ordering of the input, and whether some logic captures polynomial time there is a central open question (Chandra–Harel 1982, Gurevich 1988).

Choiceless polynomial time with counting (CPT), introduced by Blass, Gurevich and Shelah (1999), was for decades the strongest candidate: it computes with hereditarily finite sets built over the input elements, can count, and is invariant under automorphisms by construction. Blass, Gurevich and Shelah conjectured that CPT is still a proper fragment of polynomial time. This mission concerns a proof of that conjecture.

Timeline

  • 1982. Immerman and Vardi capture polynomial time on ordered structures by fixed-point logic (doi:10.1145/800070.802187, doi:10.1145/800070.802186); Chandra and Harel raise the unordered question (doi:10.1016/0022-0000(82)90012-5).
  • 1988. Gurevich formulates the question of a logic capturing polynomial time and conjectures that none exists.
  • 1992. Cai, Fürer and Immerman construct polynomial-time graph queries beyond fixed-point logic with counting (doi:10.1007/BF01305232).
  • 1999, 2002. Blass, Gurevich and Shelah introduce choiceless polynomial time and study CFI and linear-algebra instances as potential separating examples (doi:10.1016/S0168-0072(99)00005-6, doi:10.2178/jsl/1190150152).
  • 2000. Shelah claims noncapture for counting extensions; later accounts continue to list the problem as open.
  • 2008. Dawar, Richerby and Rossman show that CFI queries over ordered base graphs are CPT-definable, and that bounded-rank CPT with counting fails on them (doi:10.1016/j.apal.2007.11.011).
  • 2010–2023. Rossman and Pago prove functional lower bounds for CPT (no construction of hyperplanes, no fine preorders on hypercubes) (doi:10.1007/978-3-642-15025-8_28, doi:10.4230/LIPIcs.CSL.2021.33); Grädel, Pakusa, Schalthöfer and Kaiser characterize CPT by iterated first-order interpretations (doi:10.1109/LICS.2015.68).
  • 2025. Dawar, Grädel, Kullmann and Pago still record the capture problem as open (doi:10.4230/LIPIcs.MFCS.2025.40).

The source of this mission is an OpenAI preprint dated September 23, 2026.

Setting

Fix the field F=F3\mathbb F=\mathbb F_3F=F3​ and the vocabulary of eight binary relation symbols τ={Ed,Cf,EB,VB,I,Z0,Z1,Z2}\tau=\{\mathrm{Ed},\mathrm{Cf},\mathrm{EB},\mathrm{VB},I,Z_0,Z_1,Z_2\}τ={Ed,Cf,EB,VB,I,Z0​,Z1​,Z2​}. For a finite τ\tauτ-structure AAA put

Y={y:Ed(y,y)},X={a:Cf(a,a)},Xt={a∈X:VB(t,a)} (t∈X),Y=\{y:\mathrm{Ed}(y,y)\},\quad X=\{a:\mathrm{Cf}(a,a)\},\quad X_t=\{a\in X:\mathrm{VB}(t,a)\}\ (t\in X),Y={y:Ed(y,y)},X={a:Cf(a,a)},Xt​={a∈X:VB(t,a)} (t∈X), C(y,a)=∑x∈Y, I(a,x) ∑δ∈F, Zδ(y,x)δ.C(y,a)=\sum_{x\in Y,\ I(a,x)}\ \sum_{\delta\in\mathbb F,\ Z_\delta(y,x)}\delta .C(y,a)=x∈Y, I(a,x)∑​ δ∈F, Zδ​(y,x)∑​δ.

With unknowns (λa)a∈X(\lambda_a)_{a\in X}(λa​)a∈X​ and (μy)y∈Y(\mu_y)_{y\in Y}(μy​)y∈Y​ (separate families), the system consists of

∑a∈Xtλa=1(t∈X),μy=∑a∈Xtλa C(y,a)\sum_{a\in X_t}\lambda_a=1\quad(t\in X),\qquad \mu_y=\sum_{a\in X_t}\lambda_a\,C(y,a)a∈Xt​∑​λa​=1(t∈X),μy​=a∈Xt​∑​λa​C(y,a)

for every (t,y)∈X×Y(t,y)\in X\times Y(t,y)∈X×Y such that some a∈Xta\in X_ta∈Xt​ and x∈Yx\in Yx∈Y satisfy I(a,x)I(a,x)I(a,x) and EB(y,x)\mathrm{EB}(y,x)EB(y,x). The query Q(A)Q(A)Q(A) says this system is consistent over F3\mathbb F_3F3​. No promise is imposed on the input.

A CPT program is a fixed finite set of rules that updates dynamic functions whose values are hereditarily finite sets over the atoms of AAA, using pairing, union, comprehension, equality, membership, the input relations, and the cardinality operation (which returns a von Neumann ordinal). A program defines QQQ if for fixed polynomials T,ST,ST,S, on every input of size nnn it halts within T(n)T(n)T(n) steps, touches at most S(n)S(n)S(n) hereditarily finite objects, and accepts exactly when Q(A)Q(A)Q(A) holds. No bound on set rank is imposed.

Formalization targets

Goal: Theorem 1

Q is isomorphism-invariant,Q∈P (from an ordered encoding),Q∉CPT.Q\ \text{is isomorphism-invariant},\qquad Q\in\mathrm{P}\ \text{(from an ordered encoding)},\qquad Q\notin\mathrm{CPT}.Q is isomorphism-invariant,Q∈P (from an ordered encoding),Q∈/CPT.

The Lean statement OAI.CPTSeparation.main is the conjunction of these three clauses. It is open on the platform.

Significance

The theorem confirms the Blass–Gurevich–Shelah noncapture conjecture for the full counting formalism: CPT does not capture polynomial time on unordered structures. It removes the leading candidate logic from Gurevich's question, which itself remains open. Because QQQ is expressible in fixed-point logic with a solvability or rank operator over F3\mathbb F_3F3​ (Corollary 3, p. 3), it also shows that neither of those logics is contained in CPT. The lower bound holds for programs of arbitrary set rank, unlike the bounded-rank lower bounds of Dawar–Richerby–Rossman.

The result is proved in an OpenAI preprint, which has not been peer reviewed. No machine-checked proof exists.

Difficulty

The polynomial-time upper bound is Gaussian elimination. The obstacle is the lower bound against every CPT program, including those that build sets of unbounded rank. Bounded-variable counting games handle fixed-point logic with counting but do not by themselves control CPT, which is strictly stronger (it decides CFI queries over ordered base graphs). Rank-dependent support arguments give lower bounds only for each fixed rank. The paper needs support bounds independent of rank, for which it uses a nonabelian automorphism group acting on grid structures, then a counting equivalence for all supported objects, and finally a transfer to complete computations via an interpretation characterization of CPT.

Formalization scope

  • Input A gives a Boolean-valued binary relation for each Symbol (Ed, Cf, EB, VB, I, Z δ with δ : ZMod 3). Input.query is the existence of λ, μ satisfying the normalization and consistency equations exactly as above; an empty system is consistent.
  • OrdinaryPolynomialTime asks for a function on ordered tables and a Turing.TM2ComputableInPolyTime machine with finite alphabets computing it from the cell-by-cell table encoding, agreeing with the query for every ordering of every structure.
  • FullCPT defines terms and rules over hereditarily finite sets (Lists quotient), a state of finitely many nonempty locations, synchronous consistent updates, halt/accept flags, and the set of occurring objects (active elements plus roots of evaluated terms). EvaluationDefinable Q asks for one program and polynomials time, space that accept exactly the inputs satisfying Q, halting within time and with at most space occurring objects.
  • This is the original operational formulation; the paper's proof passes through the interpretation characterization of CPT (Grädel–Pakusa–Schalthöfer–Kaiser), which a formal proof must either formalize or bypass.

A complete development needs: hereditarily finite sets and automorphism actions on them, support theory, counting equivalence games, the grid structures with their automorphism group, and a CPT-to-interpretation translation. Contributions formalizing Proposition 2 (polynomial-time decidability), Proposition 8 (small-orbit support bound), Proposition 13 (quantitative homogeneity) and Theorem 14 (transfer through small supports) are welcome.

Selected references

  • OpenAI, Choiceless polynomial time with counting does not capture polynomial time, preprint, September 23, 2026. https://github.com/openai/math/blob/adc7f1241b42e322a6451854ab7e4b4c146bf78a/preprints/Choiceless-polynomial-time-with-counting-does-not-capture-polynomial-time-September-23-2026/paper.pdf
  • A. Blass, Y. Gurevich, S. Shelah, Choiceless polynomial time, Ann. Pure Appl. Logic, 1999. https://doi.org/10.1016/S0168-0072(99)00005-6
  • A. Blass, Y. Gurevich, S. Shelah, On polynomial time computation over unordered structures, J. Symbolic Logic, 2002. https://doi.org/10.2178/jsl/1190150152
  • J.-Y. Cai, M. Fürer, N. Immerman, An optimal lower bound on the number of variables for graph identification, Combinatorica, 1992. https://doi.org/10.1007/BF01305232
  • A. Dawar, D. Richerby, B. Rossman, Choiceless polynomial time, counting and the Cai–Fürer–Immerman graphs, Ann. Pure Appl. Logic, 2008. https://doi.org/10.1016/j.apal.2007.11.011
  • E. Grädel, W. Pakusa, S. Schalthöfer, Ł. Kaiser, Characterising choiceless polynomial time with first-order interpretations, LICS, 2015. https://doi.org/10.1109/LICS.2015.68
  • B. Rossman, Choiceless computation and symmetry, 2010. https://doi.org/10.1007/978-3-642-15025-8_28
  • B. Pago, Choiceless computation and symmetry: limitations of definability, CSL, 2021. https://doi.org/10.4230/LIPIcs.CSL.2021.33
  • N. Immerman, Relational queries computable in polynomial time, STOC, 1982. https://doi.org/10.1145/800070.802187
  • A. K. Chandra, D. Harel, Structure and complexity of relational queries, J. Comput. System Sci., 1982. https://doi.org/10.1016/0022-0000(82)90012-5
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AlgebraTheoretical Computer Science·Captain: wurtle

Generalized Star Height at Most ThreeResearch Paper

Motivation: how many nested stars does a regular language need?

A regular expression describes a set of words using letters, union, concatenation and the Kleene star P∗P^*P∗ (any finite repetition of words from PPP). The star height of an expression is the depth of nesting of its stars, and it measures how many layers of unbounded repetition a regular language genuinely needs. For ordinary expressions this hierarchy is infinite. Allowing complement as an additional operation changes the picture: Boolean conditions can replace some layers of repetition, and the generalized star-height problem asks how far this goes. Whether every regular language has generalized star height at most some fixed number — and in particular whether one star always suffices — has been one of the longest-standing questions in the algebraic theory of automata.

Timeline

  • 1963 — Eggan connects ordinary star height with the cycle structure of transition graphs and raises the question of its unboundedness (Michigan Math. J. 1963).
  • 1965 — Schützenberger characterizes the star-free languages (generalized height 000) as those recognized by finite aperiodic monoids (Inform. Control 1965).
  • 1966 — Dejean and Schützenberger show that ordinary star height is unbounded already over a two-letter alphabet (Inform. Control 1966).
  • 1992 — Pin, Straubing and Thérien prove generalized height at most one for languages recognized by finite nilpotent groups of class two and for further monoid classes (Inform. Comput. 1992, Theorems 7.3 and 7.8).
  • 2002 — Straubing distinguishes the questions "is generalized star height bounded?" and "is it at most one?" in his account of the problem (LATIN 2002, p. 537).
  • 2016–2017 — Bourne and Ruškuc prove height-one bounds for subword-counting languages with factors of length at most three (Theor. Comput. Sci. 2016); Bourne extends this to arbitrary fixed factors in his thesis (St Andrews, 2017, hdl:10023/12024).
  • 2026 — Companion OpenAI preprints prove uniform bounds of thirteen (Finite Monoid Computations and a Uniform Generalized Star-Height Bound) and four (Generalized Star Height at Most Four).
  • 2026 — The OpenAI preprint Generalized Star Height at Most Three (OpenAI Math Release, September 25, 2026) claims the bound three. It has not been peer reviewed and its theorem is not formally verified. Whether height one always suffices remains open; no language of generalized height greater than one is known.

Setting

Fix a finite alphabet Σ\SigmaΣ and let Σ∗\Sigma^*Σ∗ be the set of finite words. A generalized regular expression over Σ\SigmaΣ is built from the constants 000 (denoting ∅\emptyset∅) and 111 (denoting {ε}\{\varepsilon\}{ε}), single letters a∈Σa\in\Sigmaa∈Σ, union P∪QP\cup QP∪Q, concatenation PQPQPQ, complement ¬P\neg P¬P (taken in Σ∗\Sigma^*Σ∗), and star P∗P^*P∗. Its height is defined by

h(0)=h(1)=h(a)=0,h(P∪Q)=h(PQ)=max⁡{h(P),h(Q)},h(¬P)=h(P),h(P∗)=1+h(P).h(0)=h(1)=h(a)=0,\quad h(P\cup Q)=h(PQ)=\max\{h(P),h(Q)\},\quad h(\neg P)=h(P),\quad h(P^*)=1+h(P).h(0)=h(1)=h(a)=0,h(P∪Q)=h(PQ)=max{h(P),h(Q)},h(¬P)=h(P),h(P∗)=1+h(P).

For a regular language L⊆Σ∗L\subseteq\Sigma^*L⊆Σ∗, the generalized star height hΣ(L)h_\Sigma(L)hΣ​(L) is the least height of an expression over Σ\SigmaΣ denoting LLL.

In Lean (namespace OAI.GeneralizedStarHeight), expressions are an inductive type Expression Alphabet with exactly these seven constructors, Expression.language interprets them in Mathlib's Language Alphabet (complement is Language complement, i.e. relative to all words), Expression.height is the recursion above, and HasHeightAtMost L n means some expression has language LLL and height at most nnn.

Formalization targets

Goal: generalized star height at most three (Theorem 1.1)

For every finite alphabet Σ\SigmaΣ and every regular language L⊆Σ∗L\subseteq\Sigma^*L⊆Σ∗,

hΣ(L)≤3.h_\Sigma(L)\le 3 .hΣ​(L)≤3.

The goal is published on the platform with status Open.

Significance

The result itself. Before 2026 it was not known whether generalized star height is bounded at all; all positive results covered restricted families of monoids or languages. Theorem 1.1 gives a uniform bound independent of the alphabet and of the recognizing automaton, settling the boundedness question. It leaves open the sharper question of whether height one always suffices: no language is known to require height two. A uniform bound also implies that the hierarchy of generalized heights has at most four nonempty levels, which narrows where any separating example must be sought.

Formalizing it. The statement is short and fully elementary, but the proof composes many constructions (finite-monoid recognition, prefix codes of episodes, split identities, marker scales) whose height accounting is error-prone. A formal proof would certify the boundedness theorem itself. Mathlib already has regular languages (Language.IsRegular) and ordinary regular expressions; this mission adds generalized expressions with complement and their height.

Difficulty

The classical lower bounds for ordinary star height do not transfer once complement is allowed, and complement can simulate many but not obviously all counting arguments. The natural approach — write each fiber of a morphism from Σ∗\Sigma^*Σ∗ to a finite monoid by an expression that simulates the monoid computation letter by letter — needs one star per level of the computation's dependence on earlier history, and the depth of this dependence grows with the monoid. Group components are the difficulty: aperiodic monoids give height zero, but counting modulo nnn and nonabelian group computations require stars, and a uniform bound must handle arbitrary finite groups nested inside arbitrary monoids without the number of stars growing with their size.

Formalization scope

  • The alphabet is any Type u with [Finite Alphabet]; regularity is Mathlib's Language.IsRegular (acceptance by a finite-state DFA).
  • Expressions must use letters of the same alphabet, so no auxiliary letters may be introduced; complement is relative to all words over that alphabet.
  • height charges 111 only for stars; complement, union and concatenation are free, exactly as in the source. Since HasHeightAtMost asks only for existence of an expression, no bound on expression size or computability is claimed, matching the source's remark that the construction bounds nesting only.
  • Needed infrastructure: recognition of regular languages by finite monoids, prefix codes and unique factorization, and closure of bounded-height expressions under Boolean operations. The generalized-expression layer is reusable for the companion bounds (four and thirteen) and for formal work on star-free languages.

Selected references

  • L. C. Eggan, Transition graphs and the star-height of regular events, Michigan Math. J. 10 (1963). https://doi.org/10.1307/mmj/1028998975
  • M.-P. Schützenberger, On finite monoids having only trivial subgroups, Inform. Control 8 (1965). https://doi.org/10.1016/S0019-9958(65)90108-7
  • F. Dejean and M.-P. Schützenberger, On a question of Eggan, Inform. Control 9 (1966). https://doi.org/10.1016/S0019-9958(66)90083-0
  • J.-É. Pin, H. Straubing and D. Thérien, Some results on the generalized star-height problem, Inform. Comput. 101 (1992). https://doi.org/10.1016/0890-5401(92)90063-L
  • H. Straubing, On logical descriptions of regular languages, LATIN 2002, LNCS 2286. https://doi.org/10.1007/3-540-45995-2_46
  • T. Bourne and N. Ruškuc, On the star-height of subword counting languages and their relationship to Rees zero-matrix semigroups, Theor. Comput. Sci. 653 (2016). https://doi.org/10.1016/j.tcs.2016.09.024
  • T. Bourne, Counting subwords and other results related to the generalised star-height problem for regular languages, PhD thesis, University of St Andrews, 2017. https://hdl.handle.net/10023/12024
  • OpenAI, Generalized Star Height at Most Three, OpenAI Math Release preprint, September 25, 2026 (Theorem 1.1, p. 2). https://github.com/openai/math/blob/adc7f1241b42e322a6451854ab7e4b4c146bf78a/preprints/Generalized-Star-Height-at-Most-Three-September-25-2026/article.pdf
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Formal VerificationTheoretical Computer Science·Captain: mikedeng1

Compositional Shape Analysis by Means of Bi-Abduction 1: The Best Abductive Solution for Points-To Symbolic Heaps Is a Computable Disjunction of Symbolic HeapsResearch Paper

Motivation

Automatic verification of programs that manipulate pointers needs procedure specifications: preconditions saying which part of the heap a procedure touches and postconditions saying what it leaves behind. Writing them by hand does not scale to millions of lines of code. Calcagno, Distefano, O'Hearn and Yang (J. ACM 58(6), 2011) infer them compositionally with bi-abduction: when symbolic execution reaches a call whose precondition is not implied by the current state, it asks for the missing memory MMM that would make it hold. This abduction question in separation logic,

Δ∗M⊨H,\Delta * M \models H,Δ∗M⊨H,

is the engine behind the Abductor tool and, later, Facebook Infer.

The question has many solutions (M=falseM = \mathsf{false}M=false is always one), so its quality matters. The paper defines an order ≾\precsim≾ on solutions that prefers less memory and then logical strength, shows that a best solution always exists as a semantic object, and then asks whether it can be written down and computed in the language that the analysis actually manipulates: symbolic heaps. For the fragment whose only spatial predicate is points-to, it answers yes. That answer, Theorem 3.25, is the goal of this mission.

Setting

Values are natural numbers; 000 is nil and the nonzero naturals are the locations. A heap hhh is a finite partial map from locations to values; a stack sss assigns a value to every variable; a state is a pair (s,h)(s,h)(s,h). Write h=h0⊎h1h = h_0 \uplus h_1h=h0​⊎h1​ when h0,h1h_0,h_1h0​,h1​ have disjoint domains and hhh is their union. A predicate is an arbitrary set of states. On predicates:

  • s,h⊨F∗Gs,h \models F * Gs,h⊨F∗G iff h=h0⊎h1h = h_0 \uplus h_1h=h0​⊎h1​ with s,h0⊨Fs,h_0 \models Fs,h0​⊨F and s,h1⊨Gs,h_1 \models Gs,h1​⊨G;
  • s,h⊨F− ⁣ ⁣∗Gs,h \models F \mathbin{-\!\!*} Gs,h⊨F−∗G iff for every h′h'h′ disjoint from hhh with s,h′⊨Fs,h' \models Fs,h′⊨F, we have s,h⊎h′⊨Gs,h \uplus h' \models Gs,h⊎h′⊨G;
  • emp\mathsf{emp}emp holds in the empty heap, true\mathsf{true}true everywhere, and ¬\neg¬, ∨\vee∨, ∧\wedge∧, ∃X\exists X∃X, ∀X\forall X∀X are the Boolean and quantifier operations; F⊨GF \models GF⊨G is inclusion.

Symbolic heaps of the Points-to Instantiation. Expressions EEE are variables or natural-number constants. A pure formula Π\PiΠ is a conjunction of E=E′E = E'E=E′ and E≠E′E \neq E'E=E′. A spatial formula Σ\SigmaΣ is a ∗*∗-conjunction of points-to facts E↦E′E \mapsto E'E↦E′, possibly with ∗ true*\,\mathsf{true}∗true; here s,h⊨E↦E′s,h \models E \mapsto E's,h⊨E↦E′ iff [ ⁣[E] ⁣]s≠0[\![E]\!]s \neq 0[[E]]s=0 and hhh is the one-cell heap [ [ ⁣[E] ⁣]s↦[ ⁣[E′] ⁣]s ][\,[\![E]\!]s \mapsto [\![E']\!]s\,][[[E]]s↦[[E′]]s]. A symbolic heap is H=∃X⃗. Π∧ΣH = \exists \vec X.\, \Pi \wedge \SigmaH=∃X.Π∧Σ, and a disjunction D=H1∨⋯∨HnD = H_1 \vee \dots \vee H_nD=H1​∨⋯∨Hn​. In the abduction question (5), Δ∗D⊨H\Delta * D \models HΔ∗D⊨H, the left-hand side Δ=Π∧L1↦R1∗⋯∗Ln↦Rn\Delta = \Pi \wedge L_1 \mapsto R_1 * \dots * L_n \mapsto R_nΔ=Π∧L1​↦R1​∗⋯∗Ln​↦Rn​ is quantifier-free.

Orders. M≤M′M \le M'M≤M′ iff M′⊨M∗trueM' \models M * \mathsf{true}M′⊨M∗true (spatially smaller), and

M≾M′  ⟺  (M≤M′∧M′≰M)∨(M≤M′∧M′≤M∧M⊨M′).M \precsim M' \iff (M \le M' \wedge M' \not\le M) \vee (M \le M' \wedge M' \le M \wedge M \models M').M≾M′⟺(M≤M′∧M′≤M)∨(M≤M′∧M′≤M∧M⊨M′).

min⁡(F)\min(F)min(F) keeps the states of FFF no proper subheap of which (with the same stack) is in FFF. Elsewhere(F)=¬(F− ⁣ ⁣∗false)\mathrm{Elsewhere}(F) = \neg(F \mathbin{-\!\!*} \mathsf{false})Elsewhere(F)=¬(F−∗false) says that some separate heap satisfies FFF, and the compatible preorder is D≤cD′  ⟺  (D∧Elsewhere(Δ))≤(D′∧Elsewhere(Δ))D \le_c D' \iff (D \wedge \mathrm{Elsewhere}(\Delta)) \le (D' \wedge \mathrm{Elsewhere}(\Delta))D≤c​D′⟺(D∧Elsewhere(Δ))≤(D′∧Elsewhere(Δ)). Subtraction is F−G=F∧¬((¬emp)∗G)F - G = F \wedge \neg((\neg\mathsf{emp}) * G)F−G=F∧¬((¬emp)∗G).

Formalization targets

Goal: Theorem 3.25

There is a single algorithm fff such that for every quantifier-free Δ\DeltaΔ and symbolic heap HHH of the Points-to Instantiation, the disjunction D=f(Δ,H)D = f(\Delta, H)D=f(Δ,H) satisfies

Δ∗D⊨HandD≾M  for every predicate M with Δ∗M⊨H.\Delta * D \models H \quad\text{and}\quad D \precsim M \ \text{ for every predicate } M \text{ with } \Delta * M \models H.Δ∗D⊨HandD≾M  for every predicate M with Δ∗M⊨H.

Both halves are part of the goal: the best solution is expressible as a disjunction of symbolic heaps, and it is computable.

Milestones, in the order the paper's argument uses them

  1. Lemma 3.16. If DDD is a minimal solution w.r.t. ≤c\le_c≤c​, then min⁡(D∨¬Elsewhere(Δ))\min(D \vee \neg\mathrm{Elsewhere}(\Delta))min(D∨¬Elsewhere(Δ)) is the minimal solution w.r.t. ≾\precsim≾.
  2. Figure 2 derives solutions (§3.4.2). If Δ∗[D]▹H\Delta * [D] \triangleright HΔ∗[D]▹H is derivable, then Δ∗D⊨H\Delta * D \models HΔ∗D⊨H.
  3. Lemma 3.18. For all Δ,H\Delta, HΔ,H some DDD with Δ∗[D]▹H\Delta * [D] \triangleright HΔ∗[D]▹H is derivable, and the proof search terminates.
  4. Lemma 3.19. If Δ∗[D]▹H\Delta * [D] \triangleright HΔ∗[D]▹H is derivable and Δ∗F⊨H\Delta * F \models HΔ∗F⊨H, then D≤cFD \le_c FD≤c​F.
  5. Lemma 3.20. Incompat(Δ)\mathrm{Incompat}(\Delta)Incompat(Δ) and ¬Elsewhere(Δ)\neg\mathrm{Elsewhere}(\Delta)¬Elsewhere(Δ) are semantically equivalent.
  6. Lemma 3.23. Six algebraic properties of subtraction, including min⁡(F1∨⋯∨Fn)=G1∨⋯∨Gn\min(F_1 \vee \dots \vee F_n) = G_1 \vee \dots \vee G_nmin(F1​∨⋯∨Fn​)=G1​∨⋯∨Gn​ with Gi=((Fi−F1)⋯−Fn)G_i = ((F_i - F_1)\cdots - F_n)Gi​=((Fi​−F1​)⋯−Fn​).
  7. Quantifier removal for pure formulae (§3.4.4). ∃X.Π\exists X.\Pi∃X.Π is equivalent to a computable pure formula without XXX.
  8. Subtraction with quantifiers (§3.4.4). (∃X⃗.Δ)−(∃Y⃗.Δ′)(\exists\vec X.\Delta) - (\exists\vec Y.\Delta')(∃X.Δ)−(∃Y.Δ′) is a computable disjunction of symbolic heaps.

Significance

The result. Semantically the best solution is min⁡(Δ− ⁣ ⁣∗H)\min(\Delta \mathbin{-\!\!*} H)min(Δ−∗H) (Theorem 3.13 of the paper), a formula that nests − ⁣ ⁣∗-\!\!*−∗ and negation around ∗*∗, which no automatic separation-logic prover handles and which says nothing directly about the shape of the missing memory. Theorem 3.25 shows that in the points-to fragment this object collapses to a finite disjunction of symbolic heaps, the format in which an analysis draws "pictures of memory", and that it can be computed. It gives an optimal abduction procedure against which the paper's heuristic procedure can be measured.

Formalizing it. The theorem is proved in the paper, but several steps are given as sketches ("the proof of each of these properties is straightforward", "a straightforward exercise"), Lemma 3.18 has only a proof sketch, and the quantifier-free subtraction display on p. 33 is incomplete as printed (it inspects only the points-to facts and not the pure part of the subtracted heap). A machine-checked development would supply the missing arguments, fix the subtraction step, and give a verified reference implementation of optimal abduction for points-to symbolic heaps. No part of it is known to have been formalized.

Difficulty

The obvious route is to compute min⁡(Δ− ⁣ ⁣∗H)\min(\Delta \mathbin{-\!\!*} H)min(Δ−∗H) directly, but − ⁣ ⁣∗-\!\!*−∗ quantifies over all heaps, and symbolic heaps have no syntax for it. The paper splits the problem into three phases (a ≤c\le_c≤c​-minimal solution from a proof system, the incompatible solutions, and min⁡\minmin), and each phase carries its own difficulty. The proof system's minimality (Lemma 3.19) must hold against arbitrary semantic predicates, not only symbolic heaps, and it fails as soon as Δ\DeltaΔ may denote more than one heap for a given stack (the paper notes that adding list segments breaks it). Computing min⁡\minmin needs subtraction of symbolic heaps with quantifiers, which in turn needs quantifier elimination for Boolean combinations of equalities over an infinite domain, and the paper's quantifier-free subtraction step has to be repaired. Showing that the final disjunction is ≾\precsim≾-least, not merely a solution, needs all three phases.

Formalization scope

All statements live in the namespace BiAbduction.Systematic. Values are ℕ with nil 0, locations are the nonzero naturals, variables are ℕ (program and logical variables form one sort), and heaps are maps ℕ → Option ℕ with 0 unallocated and finite domain. Finiteness of heaps is essential, since min⁡\minmin relies on well-foundedness of the subheap order. Predicates are arbitrary Set (Stack × Heap). The syntax is built from ℕ, Bool, Sum, Prod and List by abbreviations, so "computed by an algorithm" is Mathlib's Computable₂ with the standard Primcodable encodings. The left-hand side Δ\DeltaΔ has no true conjunct, as in the paper's proofs; the right-hand side may have one. Figure 2 is an inductive relation in which psto may match any right-hand points-to fact. The soundness and ≤c\le_c≤c​-minimality statements assume the paper's bound-variable convention (no variable bound in HHH occurs in Δ\DeltaΔ); the goal does not need it.

The goal cannot be met by a degenerate reading. The competitors MMM range over all predicates, not only over disjunctions, and computability refers to fixed standard encodings. The points-to semantics excludes nil addresses, and heaps are finite.

A complete development needs the separation-logic algebra of predicates (associativity and commutativity of ∗*∗, the adjunction between ∗*∗ and − ⁣ ⁣∗-\!\!*−∗), properties of min⁡\minmin under finite heaps, the soundness and minimality of the proof system, quantifier elimination for equality logic, and a computability argument for list-manipulating procedures. The semantic layer and the quantifier-elimination result are reusable beyond this mission. Contributions to any milestone, and to the supporting lemmas of §3.3 (Lemmas 3.10–3.11, Theorem 3.13), are welcome.

Theorem numbers and page references are those of the accepted manuscript of the J. ACM article (72 pages), not of the typeset journal version.

Selected references

  • C. Calcagno, D. Distefano, P. W. O'Hearn, H. Yang, Compositional Shape Analysis by means of Bi-Abduction, Journal of the ACM 58(6), Article 26, 2011. https://doi.org/10.1145/2049697.2049700
  • C. Calcagno, D. Distefano, P. W. O'Hearn, H. Yang, Compositional Shape Analysis by means of Bi-Abduction, POPL 2009. https://doi.org/10.1145/1480881.1480917
  • J. Berdine, C. Calcagno, P. W. O'Hearn, Symbolic Execution with Separation Logic, APLAS 2005. https://doi.org/10.1007/11575467_5
  • J. C. Reynolds, Separation Logic: A Logic for Shared Mutable Data Structures, LICS 2002. https://doi.org/10.1109/LICS.2002.1029817
  • D. Pym, P. W. O'Hearn, H. Yang, Possible Worlds and Resources: The Semantics of BI, Theoretical Computer Science 315(1), 2004. https://doi.org/10.1016/j.tcs.2003.11.020
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Combinatorics·Captain: Lucas

Erdős Problem 592: which ω^β are partition ordinals?Open Problem

Motivation

Ramsey's theorem says that every red/blue colouring of the pairs of an infinite set has an infinite monochromatic subset. For well-ordered sets one can ask for more: the monochromatic set should have the same order type as the whole set. Erdős and Rado introduced the partition relation α→(β,c)2\alpha \to (\beta, c)^2α→(β,c)2 to measure exactly this, and asked which countable ordinals α\alphaα satisfy α→(α,3)2\alpha \to (\alpha, 3)^2α→(α,3)2 — every colouring either has a red copy of the whole order or a blue triangle. Such ordinals are called partition ordinals. Every partition ordinal α>1\alpha>1α>1 is a power of ω\omegaω, so the question becomes: for which countable β\betaβ is ωβ\omega^\betaωβ a partition ordinal? This is Erdős Problem 592.

The question is a basic test case for ordinal Ramsey theory: it is the smallest nontrivial "unbalanced" relation (a whole order type against a finite clique), and progress on it has repeatedly required new combinatorial methods.

Timeline (as recorded on erdosproblems.com/592):

  • 1957 — Specker. ω2→(ω2,3)2\omega^2 \to (\omega^2,3)^2ω2→(ω2,3)2, and ωn↛(ωn,3)2\omega^n \not\to (\omega^n,3)^2ωn→(ωn,3)2 for every finite n≥3n \ge 3n≥3.
  • 1972 — Chang. ωω→(ωω,3)2\omega^\omega \to (\omega^\omega,3)^2ωω→(ωω,3)2 (the subject of Erdős Problem 590). Milner extended this to ωω→(ωω,m)2\omega^\omega \to (\omega^\omega,m)^2ωω→(ωω,m)2 for all finite mmm; Larson (1973) gave a short proof.
  • 1974 — Galvin and Larson. If β≥3\beta \ge 3β≥3 and ωβ\omega^\betaωβ is a partition ordinal then β\betaβ is additively indecomposable, so β=ωγ\beta=\omega^\gammaβ=ωγ. They conjectured that every such β≥3\beta\ge3β≥3 works.
  • 2010 — Schipperus. Writing β=ωγ\beta=\omega^\gammaβ=ωγ: the relation holds when γ\gammaγ is a sum of one or two indecomposable ordinals, and fails when γ\gammaγ is a sum of four or more. This refutes the Galvin–Larson conjecture in general.

The case where γ\gammaγ is a sum of exactly three indecomposable ordinals appears to be the remaining open case.

Setting

An ordinal α\alphaα is identified with a well-ordered set XαX_\alphaXα​ of order type α\alphaα. A red/blue colouring of the complete graph KαK_\alphaKα​ on XαX_\alphaXα​ assigns to every pair of distinct vertices exactly one of two colours; equivalently, it is a pair of complementary simple graphs (red, blue) on XαX_\alphaXα​.

For ordinals α,β\alpha,\betaα,β and a cardinal ccc, the partition relation α→(β,c)2\alpha \to (\beta,c)^2α→(β,c)2 holds when every red/blue colouring of KαK_\alphaKα​ has

  • a set S⊆XαS \subseteq X_\alphaS⊆Xα​, all of whose pairs are red, whose order type (with the order inherited from XαX_\alphaXα​) is exactly β\betaβ, or
  • a set T⊆XαT \subseteq X_\alphaT⊆Xα​, all of whose pairs are blue, with ∣T∣=c|T| = c∣T∣=c.

The Lean predicate is Erdos592.OrdinalCardinalRamsey α β c, following the encoding used by the Formal Conjectures project. A partition ordinal is an α\alphaα with α→(α,3)2\alpha \to (\alpha,3)^2α→(α,3)2.

An ordinal is additively indecomposable if it is nonzero and a+b<βa+b<\betaa+b<β for all a,b<βa,b<\betaa,b<β; the additively indecomposable ordinals are exactly the powers ωδ\omega^\deltaωδ. An ordinal γ\gammaγ is the sum of kkk indecomposable ordinals when

γ=ωδ1+⋯+ωδk,δ1≥⋯≥δk,\gamma = \omega^{\delta_1}+\cdots+\omega^{\delta_k}, \qquad \delta_1 \ge \cdots \ge \delta_k,γ=ωδ1​+⋯+ωδk​,δ1​≥⋯≥δk​,

i.e. its Cantor normal form has kkk terms counted with multiplicity. The Lean predicate is Erdos592.IsSumOfIndecomposables k γ.

Formalization targets

Goal: the three-term case

γ countable, γ=ωδ1+ωδ2+ωδ3 (δ1≥δ2≥δ3)  ⟹  ωωγ→(ωωγ,3)2.\gamma \text{ countable},\ \gamma=\omega^{\delta_1}+\omega^{\delta_2}+\omega^{\delta_3}\ (\delta_1\ge\delta_2\ge\delta_3) \;\Longrightarrow\; \omega^{\omega^\gamma} \to \left(\omega^{\omega^\gamma}, 3\right)^2 .γ countable, γ=ωδ1​+ωδ2​+ωδ3​ (δ1​≥δ2​≥δ3​)⟹ωωγ→(ωωγ,3)2.

This is the positive answer in the open case, as predicted by the Galvin–Larson conjecture. Because the truth is unknown, a formal disproof (exhibiting a countable γ\gammaγ with three Cantor-normal-form terms for which the relation fails) is an equally valid resolution of the goal. Together with the milestones below, a proof of the goal gives a complete answer to Problem 592: for countable β\betaβ, ωβ\omega^\betaωβ is a partition ordinal iff β≤2\beta\le2β≤2 or β=ωγ\beta=\omega^\gammaβ=ωγ with γ\gammaγ a sum of at most three indecomposables.

Milestones (known results)

  1. Specker: ω2→(ω2,3)2\omega^2 \to (\omega^2,3)^2ω2→(ω2,3)2.
  2. Specker: ωn↛(ωn,3)2\omega^n \not\to (\omega^n,3)^2ωn→(ωn,3)2 for 3≤n<ω3 \le n < \omega3≤n<ω.
  3. Chang: ωω→(ωω,3)2\omega^\omega \to (\omega^\omega,3)^2ωω→(ωω,3)2.
  4. Galvin–Larson: β≥3\beta \ge 3β≥3 countable and ωβ→(ωβ,3)2\omega^\beta \to (\omega^\beta,3)^2ωβ→(ωβ,3)2 imply that β\betaβ is additively indecomposable.
  5. Schipperus: γ\gammaγ countable and a sum of one or two indecomposables imply ωωγ→(ωωγ,3)2\omega^{\omega^\gamma} \to (\omega^{\omega^\gamma},3)^2ωωγ→(ωωγ,3)2.
  6. Schipperus: γ\gammaγ countable and a sum of k≥4k \ge 4k≥4 indecomposables imply ωωγ↛(ωωγ,3)2\omega^{\omega^\gamma} \not\to (\omega^{\omega^\gamma},3)^2ωωγ→(ωωγ,3)2.

Significance

The result itself. A resolution of the three-term case would, together with the results above, finish the classification of countable partition ordinals of the form ωβ\omega^\betaωβ asked for in Problem 592. Either answer is informative: a positive answer shows the threshold between the positive and negative cases lies between three and four terms, and a negative answer shows it lies between two and three.

Formalizing it. The drafter is not aware of any of the milestone results in Mathlib. The Formal Conjectures entry for Problem 590 links an external Lean formalization of Chang's theorem; the other results (Specker's positive and negative theorems, Galvin–Larson, Schipperus) have, to the best of the drafter's knowledge, no public machine-checked proofs. Formalizing them is a substantial project on its own, independent of the open case, and the goal itself is an open research problem.

Difficulty

The property is not monotone in β\betaβ: it holds for β=2\beta=2β=2, fails for every finite β≥3\beta\ge3β≥3, holds again for β=ω\beta=\omegaβ=ω, and, by Schipperus, both holds and fails for various larger β=ωγ\beta=\omega^\gammaβ=ωγ depending on the number of terms in the Cantor normal form of γ\gammaγ. So no induction on β\betaβ can settle the question, and a naive transfer of the argument for a smaller exponent to a larger one can fail. The known positive and negative results use different arguments, and the three-term case lies exactly on the boundary between the ranges they cover.

Formalization scope

  • Ordinals and cardinals are Mathlib's Ordinal.{u} and Cardinal.{u} in an arbitrary universe u; "countable" is γ.card ≤ ℵ₀.
  • The graph lives on α.ToType, the canonical well-ordered type of order type α; a colouring is a pair of complementary SimpleGraphs (IsCompl red blue). A red KβK_\betaKβ​ is a red clique s with typeLT s = β; a blue K3K_3K3​ is a blue clique of cardinality exactly 3.
  • IsSumOfIndecomposables k γ requires a non-increasing list of exponents of length exactly k; without the ordering requirement, "sum of kkk" would not be well defined, since for instance ω+ω2=ω2\omega+\omega^2=\omega^2ω+ω2=ω2.
  • ω ^ ω ^ γ means ω(ωγ)\omega^{(\omega^\gamma)}ω(ωγ).
  • The Galvin–Larson milestone states additive indecomposability directly as ∀a,b<β, a+b<β\forall a,b<\beta,\ a+b<\beta∀a,b<β, a+b<β (for β≥3\beta \ge 3β≥3 this is equivalent to β=ωγ\beta=\omega^\gammaβ=ωγ).

The goal is not trivially satisfiable: the hypotheses hold, for example, for γ=3\gamma=3γ=3 and γ=ω2+ω+1\gamma=\omega^2+\omega+1γ=ω2+ω+1, and the conclusion is a genuine partition relation on an infinite ordinal.

Useful reusable infrastructure includes Cantor-normal-form combinatorics for countable ordinals, order-type calculations for subsets of ωβ\omega^\betaωβ, and a library of the classical colourings (Specker-type constructions). Contributions formalizing any milestone are welcome.

Selected references

  • T. F. Bloom, Erdős Problem #592, erdosproblems.com. https://www.erdosproblems.com/592 (this page lists the original references [Sp57], [Ch72], [GaLa74], [Sc10] cited below).
  • E. Specker, Teilmengen von Mengen mit Relationen, Comment. Math. Helv., 1957.
  • C. C. Chang, A partition theorem for the complete graph on ωω\omega^\omegaωω, J. Combinatorial Theory Ser. A, 1972.
  • J. A. Larson, A short proof of a partition theorem for the ordinal ωω\omega^\omegaωω, Ann. Math. Logic, 1973/74.
  • F. Galvin and J. Larson, Pinning countable ordinals, Fund. Math., 1974/75.
  • R. Schipperus, Countable partition ordinals, Ann. Pure Appl. Logic, 2010.
  • Formal Conjectures (Google DeepMind), Erdős Problems 590–592. https://github.com/google-deepmind/formal-conjectures
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Category TheoryFormal VerificationLogic+3·Captain: Alien60

Barrett–Halvorson 2016: Morita Equivalence Implies Categorical EquivalenceResearch Paper

Motivation

When are two scientific theories "the same theory in different words"? Philosophers of science have used two formal criteria. Definitional equivalence (Glymour) is strict: it requires a common definitional extension and therefore cannot relate theories with different numbers of sorts. Categorical equivalence (Halvorson, Weatherall) is liberal: it asks only that the categories of models be equivalent. It has been applied to classical mechanics, electromagnetism and Newtonian gravitation (Weatherall 2016; survey in Weatherall 2019).

Barrett and Halvorson (2016) proposed Morita equivalence as a syntactic criterion between these two. It allows definitions of new sorts (products, coproducts, subsorts, quotients) in addition to the usual definitions of new predicates and function symbols. Their main results place the new criterion strictly between the other two:

  • Morita equivalence implies categorical equivalence (their Theorem 5.1).
  • The converse fails (their Theorem 5.2).

This mission formalizes that comparison: the semantic core of Morita extensions, the implication, the separating example, and the paper's motivating "two kinds of things" example.

Setting

A signature Σ\SigmaΣ (Sig) consists of a type of sorts, function symbols f:σ1×⋯×σm→σf : \sigma_1\times\dots\times\sigma_m\to\sigmaf:σ1​×⋯×σm​→σ and predicate symbols p⊆σ1×⋯×σmp\subseteq\sigma_1\times\dots\times\sigma_mp⊆σ1​×⋯×σm​. Formulas (Formula L Γ) are built in a typed context Γ\GammaΓ from ⊥\bot⊥, ===, predicates, →\to→ and ∀\forall∀. A structure (Str) assigns a nonempty set to each sort and interprets the symbols. A theory TTT is a set of formulas in context, each read as its universal closure; a model of TTT is a structure satisfying them all.

The category Mod(T)\mathrm{Mod}(T)Mod(T) (Model T) has the models of TTT as objects and elementary embeddings as arrows: maps on every sort that preserve and reflect every formula. Theories T1,T2T_1,T_2T1​,T2​ are categorically equivalent (CategoricallyEquivalent) if Mod(T1)≃Mod(T2)\mathrm{Mod}(T_1)\simeq\mathrm{Mod}(T_2)Mod(T1​)≃Mod(T2​).

A Morita extension of a Σ\SigmaΣ-theory TTT (MoritaData L T) adds to Σ\SigmaΣ:

  • new sorts, each with its defining symbols:
    • a product σ\sigmaσ of σ1,σ2\sigma_1,\sigma_2σ1​,σ2​, with projections π1,π2\pi_1,\pi_2π1​,π2​;
    • a coproduct, with injections ρ1,ρ2\rho_1,\rho_2ρ1​,ρ2​;
    • a subsort cut out by a Σ\SigmaΣ-formula φ(x)\varphi(x)φ(x), with an injection iii;
    • a quotient by a Σ\SigmaΣ-formula φ(x1,x2)\varphi(x_1,x_2)φ(x1​,x2​), with a projection ε\varepsilonε;
  • new predicates and function symbols, each explicitly defined by a Σ\SigmaΣ-formula.

Each new symbol gets one defining axiom, and the symbols of one new sort share a single axiom. The extension T+T^+T+ (D.ext) is TTT together with these axioms. Each definition must satisfy its admissibility condition in every model of TTT:

DefinitionCondition
subsort∃x φ(x)\exists x\,\varphi(x)∃xφ(x)
quotientφ\varphiφ is an equivalence relation
function symbol∀x ∃!y φ(x,y)\forall x\,\exists!y\,\varphi(x,y)∀x∃!yφ(x,y)

The reduct gives a functor Π:Mod(T+)→Mod(T)\Pi:\mathrm{Mod}(T^+)\to\mathrm{Mod}(T)Π:Mod(T+)→Mod(T) (D.proj).

T1T_1T1​ and T2T_2T2​ are Morita equivalent (MoritaEquivalent) if two finite chains of Morita extensions T1⇝T1′T_1\rightsquigarrow T_1'T1​⇝T1′​ and T2⇝T2′T_2\rightsquigarrow T_2'T2​⇝T2′​ (MoritaReach) end in logically equivalent theories, that is, theories with the same models, over a common signature.

Formalization targets

Goal: Morita equivalence implies categorical equivalence

T1 Morita equivalent to T2  ⟹  Mod(T1)≃Mod(T2).T_1 \text{ Morita equivalent to } T_2 \;\Longrightarrow\; \mathrm{Mod}(T_1)\simeq\mathrm{Mod}(T_2).T1​ Morita equivalent to T2​⟹Mod(T1​)≃Mod(T2​).

Milestones

  1. Unique expansion. Every model MMM of TTT has an expansion N⊨T+N\models T^+N⊨T+ with ΠN≅M\Pi N\cong MΠN≅M, and any two such expansions are isomorphic.
  2. Π\PiΠ is faithful.
  3. Π\PiΠ is full.
  4. Conservativity. For every Σ\SigmaΣ-formula α\alphaα: T⊨αT\models\alphaT⊨α iff T+⊨αT^+\models\alphaT+⊨α.
  5. Translation. Every Σ+\Sigma^+Σ+-formula whose free variables all have old sorts is T+T^+T+-equivalent to a Σ\SigmaΣ-formula.
  6. Two kinds of things. T1={∃x p(x), ∃x q(x), ∀x (p(x)∨q(x)), ∀x ¬(p(x)∧q(x))}T_1=\{\exists x\,p(x),\ \exists x\,q(x),\ \forall x\,(p(x)\vee q(x)),\ \forall x\,\neg(p(x)\wedge q(x))\}T1​={∃xp(x), ∃xq(x), ∀x(p(x)∨q(x)), ∀x¬(p(x)∧q(x))} over one sort is Morita equivalent to the empty theory over two sorts.
  7. Separating example, first half. T1={∃=1x (x=x)}T_1=\{\exists^{=1}x\,(x=x)\}T1​={∃=1x(x=x)} and T2=T1∪{∀y (p0(y)→pi(y)):i∈N}T_2=T_1\cup\{\forall y\,(p_0(y)\to p_i(y)) : i\in\mathbb N\}T2​=T1​∪{∀y(p0​(y)→pi​(y)):i∈N}, over one sort and countably many unary predicates, are categorically equivalent.
  8. Separating example, second half. These two theories are not Morita equivalent.

Milestones 1–3 together say that Π\PiΠ is an equivalence of categories. Milestones 6–8 are not needed for the goal; they record the paper's examples, which show that Morita equivalence relates theories with different numbers of sorts and is strictly stronger than categorical equivalence.

Significance

The result itself. The goal makes Morita equivalence a syntactic sufficient condition for categorical equivalence: an explicit dictionary of definitions yields an equivalence of model categories. Milestone 8 shows the converse fails, so categorical equivalence is strictly weaker. Milestone 6 shows that Morita equivalence, unlike definitional equivalence, can relate theories with different numbers of sorts. Together these locate the criterion that later philosophy-of-physics work uses to compare formulations with different ontologies, such as different numbers of kinds of objects.

Formalizing it. The results are proved on paper and, as far as is known, have no machine-checked proofs. Mathlib's ModelTheory library is single-sorted, so none of this is available there. The mission produces a reusable many-sorted first-order logic with:

  • syntax, semantics and signature maps;
  • elementary embeddings and the category of models;
  • Morita extensions with their reduct functors.

The paper's arguments for conservativity and translation are informal in places; for example, the translation in the appendix uses coding of new-sorted variables. A formal proof fixes these details.

Difficulty

The obvious route to the goal composes the reduct functors along the chains. The work lies in showing that each reduct is an equivalence when the arrows are elementary embeddings.

  • Fullness. An elementary embedding of reducts extends uniquely to the new sorts. Showing that the extension is again elementary requires the translation of Milestone 5: every Σ+\Sigma^+Σ+-formula has to be pulled back to a Σ\SigmaΣ-formula, including formulas that quantify over the new sorts.
  • Coproduct and quotient sorts. Quantifiers over these sorts must be eliminated in favour of quantifiers over the old sorts.

For Milestone 8 the difficulty is reversed. One must rule out every pair of chains, including chains that introduce arbitrary new sorts. The paper's proof argues with a single common Morita extension. The definition, however, allows chains, and the paper notes that a chain need not collapse to a single extension. The argument must therefore be run through conservativity and translation along each chain.

Formalization scope

  • Self-contained logic. Signatures have dependent arities (Fin m → Srt). Terms and formulas use de Bruijn variables in typed contexts Γ : Fin n → Srt; Formula.all binds variable 0.
  • Theories. A theory is a Set (Axiom L), each axiom being a formula in context read as its universal closure. Because sorts are nonempty, this agrees with the sentence version.
  • Nonempty sorts. Every structure has nonempty carriers. This is the paper's convention, and without it Milestone 6 is false: the empty theory over two sorts would have models with an empty sort.
  • Morita extensions are built, not recognized. MoritaData L T lists the new sorts and symbols and their defining Σ\SigmaΣ-formulas. The signature Σ+\Sigma^+Σ+ (D.sig) and the axioms are generated from it, and admissibility is stated semantically, over all models of TTT. Any Morita extension in the paper's sense is of this form up to renaming of the new symbols.
  • A common signature up to renaming. MoritaEquivalent asks for a bijective signature map F between the two chain ends such that a structure is a model of T2′T_2'T2′​ iff its F-reduct is a model of T1′T_1'T1′​.
  • Universes. Models live in Type.
  • Separating example. The paper writes T2T_2T2​ with fresh symbols qiq_iqi​; here T1T_1T1​ and T2T_2T2​ share the signature with predicates pip_ipi​, which only renames symbols.
  • No trivialization. None of the targets holds vacuously: Morita extensions require genuine admissibility proofs, and the arrows are elementary embeddings, not arbitrary maps. Milestone 6 needs explicit chains and an explicit renaming, with arity casts.

A complete development will need:

  • the substitution and weakening lemmas for de Bruijn formulas;
  • the realization of signature maps;
  • the construction of expansions;
  • a translation function from Σ+\Sigma^+Σ+-formulas to Σ\SigmaΣ-formulas.

These pieces are reusable for any later work on interpretations or theoretical equivalence. Alternative proofs of Milestone 8, for example by counting finitely axiomatized subtheories, are welcome.

Selected references

  • T. W. Barrett, H. Halvorson, Morita Equivalence, Rev. Symb. Logic 9 (2016) 556–582. arXiv:1506.04675. Source of all targets: §3–§5 and the appendix; Example 4 and Theorems 4.1–4.3, 5.1, 5.2.
  • J. O. Weatherall, Are Newtonian Gravitation and Geometrized Newtonian Gravitation Theoretically Equivalent?, Erkenntnis 81 (2016) 1073–1091. arXiv:1411.5757.
  • J. O. Weatherall, Theoretical Equivalence in Physics, Philos. Compass 14 (2019). arXiv:1810.08192.
  • H. Halvorson, The Logic in Philosophy of Science, Cambridge University Press, 2019. doi:10.1017/9781316275603.
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CombinatoricsGraph Theory·Captain: mikedeng1

Twin-width I: Tractable FO Model Checking 5: First-Order Interpretations of a Class of Graphs of Bounded Twin-width Have Bounded Twin-widthResearch Paper

Why interpretations

Twin-width is a graph parameter introduced by Bonnet, Kim, Thomassé and Watrigant in Twin-width I: Tractable FO Model Checking (J. ACM 69(1), Article 3, 2021; arXiv:2004.14789). Its main algorithmic consequence is that first-order model checking is fixed-parameter tractable on graphs given with a contraction sequence of bounded width. A parameter of this kind is useful only if the classes it bounds are closed under the constructions that arise in practice. Many such constructions are first-order definable: the complement of a graph, its square (join vertices at distance at most two), the map graph of a planar map. Section 8 of the paper shows that twin-width is stable under every such construction: a first-order interpretation of a class of bounded twin-width has bounded twin-width. Bounded VC-dimension, by contrast, is not preserved (interval graphs have VC-dimension at most two but interpret every graph, p. 3:41).

The result has since become the basis of the "model-theoretic" view of twin-width: bounded twin-width classes are closed under first-order transductions (Theorem 8.1 of the paper), and later work identifies them, among ordered structures, with the monadically dependent classes (Twin-width IV, Bonnet, Giocanti, Ossona de Mendez, Simon, Thomassé and Toruńczyk).

Setting

Let GGG be a finite simple graph on a vertex set VVV.

Two vertex sets X,YX,YX,Y are homogeneous if all pairs in X×YX\times YX×Y are edges or none is. For a partition P\mathcal PP of VVV, the red graph GPG_{\mathcal P}GP​ has the parts as vertices, two distinct parts being adjacent when they are not homogeneous. P\mathcal PP is a ddd-partition if GPG_{\mathcal P}GP​ has maximum degree at most ddd. The graph GGG has twin-width at most ddd, tww⁡(G)≤d\operatorname{tww}(G)\le dtww(G)≤d, if there is a sequence of ddd-partitions from the partition into singletons to a partition with at most one part, each obtained from the previous one by merging two parts (p. 3:32).

A prenex formula of depth ℓ\ellℓ with two free variables is

φ(x,y)=Q1x1 Q2x2⋯Qℓxℓ φ∗,Qi∈{∀,∃},\varphi(x,y)=Q_1x_1\,Q_2x_2\cdots Q_\ell x_\ell\ \varphi^*,\qquad Q_i\in\{\forall,\exists\},φ(x,y)=Q1​x1​Q2​x2​⋯Qℓ​xℓ​ φ∗,Qi​∈{∀,∃},

with φ∗\varphi^*φ∗ a Boolean combination of atoms u=vu=vu=v and E(u,v)E(u,v)E(u,v), u,v∈{x1,…,xℓ,x,y}u,v\in\{x_1,\dots,x_\ell,x,y\}u,v∈{x1​,…,xℓ​,x,y}. The interpretation φ(G)\varphi(G)φ(G) is the graph on VVV in which distinct u,vu,vu,v are adjacent iff G⊨φ(u,v)∧φ(v,u)G\models\varphi(u,v)\wedge\varphi(v,u)G⊨φ(u,v)∧φ(v,u). For a class G\mathcal GG, φ(G)\varphi(\mathcal G)φ(G) is the class of all induced subgraphs of the graphs φ(G)\varphi(G)φ(G), G∈GG\in\mathcal GG∈G.

The proof works with morphism-trees. The complete tree MTk(V)MT_k(V)MTk​(V) has as nodes all tuples (v1,…,vi)(v_1,\dots,v_i)(v1​,…,vi​) of vertices with i≤ki\le ki≤k, the parent of a tuple being its prefix. Two sibling nodes are equivalent in (G,P)(G,\mathcal P)(G,P) if an automorphism of the tree swaps them, preserving equalities and adjacencies among the entries of every tuple and the part of P\mathcal PP of every entry. A reduction deletes one of two equivalent siblings with its subtree, repeatedly; a reduct admits no further reduction. Two vertices u,u′u,u'u,u′ are adjacent in Eℓ+2(G,P)E_{\ell+2}(G,\mathcal P)Eℓ+2​(G,P) when (u)(u)(u) and (u′)(u')(u′) are equivalent in some reduction of MTℓ+2(V)MT_{\ell+2}(V)MTℓ+2​(V), and Iℓ+2(G,P)I_{\ell+2}(G,\mathcal P)Iℓ+2​(G,P) is the partition into the connected components of Eℓ+2(G,P)E_{\ell+2}(G,\mathcal P)Eℓ+2​(G,P).

Formalization targets

Goal: Theorem 8.3 (graphs)

For every prenex formula φ(x,y)\varphi(x,y)φ(x,y) and every ddd there is DDD such that

tww⁡(G)≤d ⟹ tww⁡(φ(G)[S])≤Dfor every graph G and every S⊆V(G).\operatorname{tww}(G)\le d\ \Longrightarrow\ \operatorname{tww}\bigl(\varphi(G)[S]\bigr)\le D\quad\text{for every graph } G \text{ and every } S\subseteq V(G).tww(G)≤d ⟹ tww(φ(G)[S])≤Dfor every graph G and every S⊆V(G).

The bound DDD depends on φ\varphiφ and ddd only. No explicit function is asserted: the paper gives none (the proof goes through a tower-type bound on reducts).

Milestones

  1. Lemma 5.2 in graph form (p. 3:18, applied on p. 3:43): rrr-refining hhh-partitions from the finest to the coarsest partition give tww⁡(G)≤r(h+1)\operatorname{tww}(G)\le r(h+1)tww(G)≤r(h+1).
  2. Lemma 7.3 (p. 3:32): reducts of MTℓ(G)MT_\ell(G)MTℓ​(G) have size bounded by a function of ℓ\ellℓ.
  3. Lemmas 7.11–7.14 (pp. 3:35–3:38): restriction to connected tuples rooted at a part and the pruned shuffle commute with reductions; the pruned shuffle of all MTℓ(G,P,X)MT_\ell(G,\mathcal P,X)MTℓ​(G,P,X) is MTℓ(G,P)MT_\ell(G,\mathcal P)MTℓ​(G,P), and that of reducts is a reduction of it.
  4. Lemma 8.4 (p. 3:41) and its extension to reductions (p. 3:42): equivalent nodes (u,v),(u,v′)(u,v),(u,v')(u,v),(u,v′) of MTℓ+2MT_{\ell+2}MTℓ+2​ satisfy the same prenex formulas of depth ℓ\ellℓ.
  5. Iℓ+2I_{\ell+2}Iℓ+2​ refines P\mathcal PP and is monotone under coarsening (p. 3:42).
  6. Lemma 8.5 (p. 3:42): a part of a ddd-partition meets boundedly many components of Eℓ+2(G,P)E_{\ell+2}(G,\mathcal P)Eℓ+2​(G,P).
  7. Lemma 8.6 (p. 3:43): parts of Iℓ+2(G,P)I_{\ell+2}(G,\mathcal P)Iℓ+2​(G,P) inside parts at distance at least 3ℓ+23^{\ell+2}3ℓ+2 in GPG_{\mathcal P}GP​ are homogeneous in φ(G)\varphi(G)φ(G).

Significance

Theorem 8.3 turns every first-order definable graph construction into a source of bounded-twin-width classes: squares and fixed powers of planar graphs, complements, map graphs (as transductions, via Theorem 8.1), kkk-planar graphs and bounded-degree string graphs. Combined with the model-checking algorithm it gives FPT first-order model checking on all these classes, provided a contraction sequence is available. The theorem is also the first step of the model-theoretic characterisations of bounded twin-width in the later papers of the series.

The result is proved in the paper; no machine-checked proof of it, or of any twin-width bound, exists to our knowledge, and Mathlib has no notion of twin-width or of first-order interpretations of graphs. The mission asks for a formal proof of the graph case. The morphism-tree milestones are also the combinatorial core of the paper's linear-time model-checking algorithm (Theorems 7.5, 7.15), whose running-time statements are not part of this mission.

Difficulty

The obvious approach refines the ddd-partitions Pi\mathcal P_iPi​ of GGG by the "type" of each vertex with respect to φ\varphiφ. The number of such types is not bounded: a vertex's behaviour depends on the whole graph through the quantifiers. The paper replaces types by the components of Eℓ+2(G,Pi)E_{\ell+2}(G,\mathcal P_i)Eℓ+2​(G,Pi​), defined through reductions of morphism-trees that respect Pi\mathcal P_iPi​. Two facts must then be shown, and neither is local in an obvious way. First, the number of components inside a part is bounded (Lemma 8.5); this needs the pruned shuffle, which assembles a reduction of the whole tree MTℓ+2(G,P)MT_{\ell+2}(G,\mathcal P)MTℓ+2​(G,P) from reductions of the local trees MTℓ+2(G,P,X)MT_{\ell+2}(G,\mathcal P,X)MTℓ+2​(G,P,X), and the fact that it commutes with reductions (Lemma 7.12), where an automorphism of one local tree must be extended to the shuffle without breaking adjacency between non-homogeneous parts. Second, far-apart components are homogeneous in φ(G)\varphi(G)φ(G) (Lemma 8.6), which needs the evaluation argument of Lemma 8.4 inside a reduction of the partitioned tree.

Formalization scope

  • Graphs are SimpleGraph V on a Fintype with decidable equality. Twin-width is the predicate TwinWidthLE G d in the paper's partition form (pp. 3:12, 3:32); trigraphs are not formalized, and twin-width is never an infimum, so no junk value of sInf can make a bound vacuous. Merge steps merge exactly two distinct parts, so a sequence cannot jump to one part.
  • Formulas are Mathlib's FirstOrder.Language.graph.Formula (Fin 2), evaluated in G.structure; variable 0 is xxx, variable 1 is yyy. Prenex is Mathlib's IsPrenex; quantifier depth qdepth is defined in the mission (∃=¬∀¬\exists=\neg\forall\neg∃=¬∀¬ counts once).
  • The interpretation requires u≠vu\neq vu=v for an edge (a simple graph has no loops). The goal is the graph case of Theorem 8.3, which the page states for augmented binary structures; Section 8 itself says it works with undirected graphs (p. 3:40).
  • In the goal the bound DDD is chosen after φ\varphiφ and ddd and before the graph: a bound chosen after the graph would make the statement trivial. The hereditary closure (induced subgraphs on every SSS) is part of the statement.
  • Morphism-trees are sets of tuples (Set (List V)), the paper's identification of a node with its current path (p. 3:30), which is exact for reductions of complete trees. Equivalent siblings are distinct. Distances in GPG_{\mathcal P}GP​ are in N∪{∞}\mathbb N\cup\{\infty\}N∪{∞}, so disconnected parts are far apart. The sequence graph uses 0-based indices, hence the exponent 3ℓ−(k+1)3^{\ell-(k+1)}3ℓ−(k+1). The pruned shuffle is defined by its characterisation (each component of the sequence graph of a tuple lies in the tree of its local root). Lemmas of §7.3 carry its standing assumption ℓ>0\ell>0ℓ>0.
  • Lemma 5.2 is posed for graphs with the bound r(h+1)r(h+1)r(h+1) (the matrix lemma's rtrtrt with t=h+1t=h+1t=h+1).
  • Not in the mission: Theorem 8.1 (transductions), Lemma 8.2 and Lemma 5.1 (augmented binary structures), and all running-time statements (Theorems 1.1, 7.1, 7.5, 7.15, Corollaries 7.16–7.17).

Reusable beyond this mission: the partition form of twin-width, quantifier depth of first-order formulas, and the morphism-tree machinery (reductions, pruned shuffles), which formalizes the algorithmic core of first-order model checking on bounded twin-width. Contributions welcome: proofs of the §7 lemmas, of Lemma 8.4 by induction on the quantifier prefix, and the closure of twin-width under induced subgraphs.

Selected references

  • É. Bonnet, E. J. Kim, S. Thomassé, R. Watrigant, Twin-width I: Tractable FO Model Checking, J. ACM 69(1), Article 3, 2021. https://doi.org/10.1145/3486655 (arXiv:2004.14789, https://arxiv.org/abs/2004.14789)
  • É. Bonnet, U. Giocanti, P. Ossona de Mendez, P. Simon, S. Thomassé, S. Toruńczyk, Twin-width IV: Ordered Graphs and Matrices, J. ACM, 2024. https://arxiv.org/abs/2102.03117
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Formal VerificationTheory of Computation·Captain: mikedeng1

Monitoring Metric First-Order Temporal Properties: The Monitor M_Φ Outputs Exactly the Violations (¬Φ)^(D̄,τ̄,i) of a Bounded MFOTL Formula Φ, as Regular Sets, at Every Time Point iResearch Paper

Motivation

Runtime monitoring checks a running system against a specification by inspecting its trace of events as it is produced. Policies on data, such as "every published report was approved within the last 10 days" or "each value stored in in reaches out within 5 time units", quantify over data values and constrain the time between events. Metric first-order temporal logic (MFOTL) expresses such policies: it extends first-order logic with past and future temporal operators that carry intervals of allowed time differences.

Basin, Klaedtke, Müller and Zălinescu (J. ACM 62(2), Article 15, 2015) give a monitor for MFOTL properties □Φ\square\Phi□Φ with Φ\PhiΦ bounded. It allows negation and quantification over the infinite domain N\mathbb NN without restriction, because every relation it manipulates is represented by a finite automaton: the input structures are automatic structures. The paper implements it in the prototype tool MonPoly-Reg (§6). Its correctness rests on Theorem 3.9: the monitor reports exactly the violations of Φ\PhiΦ, at every time point.

Setting

A signature S=(C,R,ι)S=(C,R,\iota)S=(C,R,ι) has finitely many constants and predicates with arities. Formulas are built from t≈t′t\approx t't≈t′ and r(t1,…,tι(r))r(t_1,\dots,t_{\iota(r)})r(t1​,…,tι(r)​) by ¬\neg¬, ∨\vee∨, ∃x\exists x∃x and the temporal operators ∙I\bullet_I∙I​ (previous), ∘I\circ_I∘I​ (next), SI\mathsf S_ISI​ (since) and UI\mathsf U_IUI​ (until), where I=[b,b′)I=[b,b')I=[b,b′) is a nonempty interval of N\mathbb NN with b′∈N∪{∞}b'\in\mathbb N\cup\{\infty\}b′∈N∪{∞}. A formula is bounded if every UI\mathsf U_IUI​ in it has b′<∞b'<\inftyb′<∞.

A temporal structure (Dˉ,τˉ)(\bar{\mathcal D},\bar\tau)(Dˉ,τˉ) assigns to each time point i∈Ni\in\mathbb Ni∈N relations rDi⊆Nι(r)r^{\mathcal D_i}\subseteq\mathbb N^{\iota(r)}rDi​⊆Nι(r) and a time stamp τi∈N\tau_i\in\mathbb Nτi​∈N; constants are rigid, τ0≤τ1≤⋯\tau_0\le\tau_1\le\cdotsτ0​≤τ1​≤⋯, and τˉ\bar\tauτˉ exceeds every bound. Satisfaction (Dˉ,τˉ,v,i)⊨φ(\bar{\mathcal D},\bar\tau,v,i)\models\varphi(Dˉ,τˉ,v,i)⊨φ is Definition 2.2; for example ψ UI ψ′\psi\,\mathsf U_I\,\psi'ψUI​ψ′ holds at iii if ψ′\psi'ψ′ holds at some j≥ij\ge ij≥i with τj−τi∈I\tau_j-\tau_i\in Iτj​−τi​∈I and ψ\psiψ holds at all k∈[i,j)k\in[i,j)k∈[i,j). For φ\varphiφ with free variables x1<⋯<xnx_1<\dots<x_nx1​<⋯<xn​, the satisfying set is

φ(Dˉ,τˉ,i)={dˉ∈Nn∣(Dˉ,τˉ,v[xˉ↦dˉ],i)⊨φ for some v}.\varphi^{(\bar{\mathcal D},\bar\tau,i)}=\{\bar d\in\mathbb N^n\mid(\bar{\mathcal D},\bar\tau,v[\bar x\mapsto\bar d],i)\models\varphi\ \text{for some }v\}.φ(Dˉ,τˉ,i)={dˉ∈Nn∣(Dˉ,τˉ,v[xˉ↦dˉ],i)⊨φ for some v}.

A domain representation is a regular language L\mathcal LL over a finite alphabet with a surjection ν:L→N\nu:\mathcal L\to\mathbb Nν:L→N whose equality relation is regular. A relation A⊆NkA\subseteq\mathbb N^kA⊆Nk is regular if the words u1⊗⋯⊗uku_1\otimes\cdots\otimes u_ku1​⊗⋯⊗uk​ (the letter-by-letter convolution, padded with #\##) with (ν(u1),…,ν(uk))∈A(\nu(u_1),\dots,\nu(u_k))\in A(ν(u1​),…,ν(uk​))∈A form a regular language.

The monitor MΦ\mathsf M_\PhiMΦ​ (Fig. 2) reads (D0,τ0),(D1,τ1),…(\mathcal D_0,\tau_0),(\mathcal D_1,\tau_1),\dots(D0​,τ0​),(D1​,τ1​),… one element per loop iteration. It extends each Dj\mathcal D_jDj​ to D^j\hat{\mathcal D}_jD^j​ with auxiliary relations pαp_\alphapα​ (and rαr_\alpharα​, sαs_\alphasα​ for since and until) for every temporal subformula α\alphaα, built incrementally by the constructions of §3.4 once a list QQQ of pending triples (α,j,S)(\alpha,j,S)(α,j,S) says that enough of the future has been read. When D^i\hat{\mathcal D}_iD^i​ is complete it outputs the set (¬Φ^)D^i(\neg\hat\Phi)^{\hat{\mathcal D}_i}(¬Φ^)D^i​ with τi\tau_iτi​, where Φ^\hat\PhiΦ^ is Φ\PhiΦ with each top-level temporal subformula α\alphaα replaced by the atom pα(xˉ)p_\alpha(\bar x)pα​(xˉ).

Formalization targets

Goal: Theorem 3.9

For every input satisfying the restrictions of §3.1 and every bounded Φ\PhiΦ:

(i)(n,O,t) output by line 9 ⟹ O=(¬Φ)(Dˉ,τˉ,n), O regular, t=τn;\text{(i)}\quad (n,O,t)\ \text{output by line 9}\ \Longrightarrow\ O=(\neg\Phi)^{(\bar{\mathcal D},\bar\tau,n)},\ O\ \text{regular},\ t=\tau_n;(i)(n,O,t) output by line 9 ⟹ O=(¬Φ)(Dˉ,τˉ,n), O regular, t=τn​; (ii)∀n∈N,n=0 or some output at index n−1 occurs,\text{(ii)}\quad \forall n\in\mathbb N,\quad n=0\ \text{or some output at index }n-1\text{ occurs},(ii)∀n∈N,n=0 or some output at index n−1 occurs,

where each output is followed by line 11 setting iii to the next index. This records counter values reached inside the while loop, including values that may never appear at a loop entry. Part (ii) excludes the vacuous reading of (i).

Milestones

  1. Lemma 3.4: φ^D^i=φ(Dˉ,τˉ,i)\hat\varphi^{\hat{\mathcal D}_i}=\varphi^{(\bar{\mathcal D},\bar\tau,i)}φ^​D^i​=φ(Dˉ,τˉ,i) once the auxiliary relations of tsub(φ)\mathit{tsub}(\varphi)tsub(φ) are correct, and regularity is preserved.
  2. Lemmas 3.5–3.8: the constructions for ∙I\bullet_I∙I​, ∘I\circ_I∘I​, SI\mathsf S_ISI​ and bounded UI\mathsf U_IUI​ produce regular relations equal to the satisfying sets, with the stated characterizations of rαr_\alpharα​ and sαs_\alphasα​.
  3. Observations (1)–(3) of the proof of Theorem 3.9 on the list QQQ and the counter iii.
  4. The claim that line 7 can always be executed, in the form: if (α,j,∅)∈Qk(\alpha,j,\emptyset)\in Q_k(α,j,∅)∈Qk​ then the build stores pαD^j=α(Dˉ,τˉ,j)p_\alpha^{\hat{\mathcal D}_j}=\alpha^{(\bar{\mathcal D},\bar\tau,j)}pαD^j​​=α(Dˉ,τˉ,j), which is regular.

Significance

Theorem 3.9 makes MΦ\mathsf M_\PhiMΦ​ a sound and complete monitor for the safety fragment □Φ\square\Phi□Φ, Φ\PhiΦ bounded: every violation at every time point is reported, nothing else is, and every output is a finitely represented (regular) set even though it may be infinite. It covers unrestricted negation and quantification, which finite-relation monitors (§4 of the paper) cannot handle. The incremental constructions of §3.4 are the template reused by later MFOTL monitors.

The paper gives a mathematical proof of its monitor's correctness. This mission asks for a Lean proof that checks the MFOTL semantics, automatic relations over N\mathbb NN, and the correctness of this particular monitor, including the until construction whose printed version is defective (see Formalization scope).

Difficulty

Two parts are hard. The regularity claims need the fundamental property of automatic structures: first-order definable relations are regular, which requires closure of padded convolution languages under Boolean operations, cylindrification, permutation of tracks and projection, the last with removal of trailing padding. The arithmetic constraints of the since and until constructions additionally need successor and offset relations definable from ≺\prec≺.

The scheduling argument is the other part. Correctness of a build at (α,j)(\alpha,j)(α,j) requires that every input relation, at time points up to j+ℓjj+\ell_jj+ℓj​ for until, was built before and not discarded since. Following only the syntax of α\alphaα does not settle this: a needed relation for a subformula at a later time point may be built in the same iteration as α\alphaα's.

Formalization scope

The domain is N\mathbb NN, as §3.1 assumes. Relations are sets of lists of natural numbers, ordered by the sorted free variables. Interval upper bounds are in N∞\mathbb N_\inftyN∞​; intervals carry the proof that they are nonempty. The representation (L,ν)(\mathcal L,\nu)(L,ν) is a single parameter shared by the hypothesis "every rDir^{\mathcal D_i}rDi​ is regular" and the regularity claims of the conclusions; it is never chosen existentially. The binary predicate ≺\prec≺ interpreted as <<< is a hypothesis.

Two hypotheses on Φ\PhiΦ besides boundedness are the paper's own "without loss of generality" assumptions: each temporal subformula occurs once in Φ\PhiΦ (§3.6), and the direct subformulas of each SI\mathsf S_ISI​, UI\mathsf U_IUI​ subformula have the same free variables (§3.4; achieved by padding with x≈xx\approx xx≈x, as in §3.5). No other restriction is placed on Φ\PhiΦ.

The printed until construction (§3.4.4) is repaired: the guard of UrU_rUr​ is "aˉ∈β^D^i+k\bar a\in\hat\beta^{\hat{\mathcal D}_{i+k}}aˉ∈β^​D^i+k​ for all ℓi−1≤k≤ℓi\ell_{i-1}\le k\le\ell_iℓi−1​≤k≤ℓi​", since the printed guard empties rαr_\alpharα​ in the paper's own example of §3.5; UsU_sUs​ keeps only tuples with j′≥1j'\ge1j′≥1, since otherwise a γ\gammaγ-witness at time point i−1i-1i−1 survives to time point iii when τi=τi−1\tau_i=\tau_{i-1}τi​=τi−1​; and time-stamp equations are encoded additively. Lemma 3.8(i) reads aˉ∈Nn\bar a\in\mathbb N^naˉ∈Nn and adds j≤ℓij\le\ell_ij≤ℓi​. "Effectively computable" in Theorem 3.9(i) is not formalized. Part (ii) uses the output at index n−1n-1n−1 as the record of the line-11 increment to nnn; for n=0n=0n=0, the counter is initialized by line 2. Loop-entry states alone can skip counter values reached within one iteration.

The monitor is a state machine whose builds and outputs are computed only from its store of relations; it never consults the satisfaction relation or satisfying sets. A formalization in which line 7 reads the semantics would turn Theorem 3.9(i) into a restatement of Lemma 3.4, and is ruled out. A build whose input is missing stores nothing, so a monitor that silently fails would violate (ii).

Contributions welcome: the theory of automatic relations (reusable beyond this mission), the semantic halves of Lemmas 3.4–3.8, and the scheduling invariants.

Selected references

  • D. Basin, F. Klaedtke, S. Müller, E. Zălinescu, Monitoring metric first-order temporal properties, J. ACM 62(2), Article 15, 2015. https://doi.org/10.1145/2699444
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AnalysisOptimization·Captain: mikedeng1

Clarke Subgradients of Stratifiable Functions II: The Nonsmooth Kurdyka–Łojasiewicz Inequality for Lower Semicontinuous Functions Definable in an O-minimal StructureResearch Paper

Motivation

The Kurdyka–Łojasiewicz (KL) inequality is the analytic engine behind most global convergence proofs for descent methods on nonconvex problems: proximal point and proximal gradient schemes, alternating minimization, ADMM variants and subgradient flows all reach a critical point with finite trajectory length once the objective satisfies a KL inequality. Its origin is Łojasiewicz's gradient inequality for real-analytic functions ([Łojasiewicz, 1963]); Kurdyka extended it to differentiable functions definable in an o-minimal structure, with a reparametrization ψ of the values, on bounded sets (Kurdyka, Ann. Inst. Fourier 48 (1998)).

Optimization objectives are rarely smooth or finite everywhere: indicator functions of constraint sets, ℓ₁ penalties, rank functions and maxima take the value +∞ or have kinks. Bolte, Daniilidis, Lewis and Shiota (SIAM J. Optim. 18(2) (2007), 556–572) proved that every lower semicontinuous function definable in an o-minimal structure satisfies a KL inequality for Clarke subgradients, globally in space. This mission formalizes that result (Theorem 14) and the chain of statements in §4 of the paper on which its proof rests.

Timeline. Łojasiewicz (1963): gradient inequality for real-analytic functions near a critical point. Kurdyka (1998): the inequality ‖∇(ψ∘f)‖ ≥ 1 for C¹ definable functions on bounded sets. Bolte, Daniilidis and Lewis (SIAM J. Optim. 17(4) (2007)): nonsmooth Łojasiewicz inequality for lower semicontinuous subanalytic functions with the limiting subdifferential. Bolte, Daniilidis, Lewis and Shiota (2007): this paper, for definable functions, Clarke subgradients, and unbounded sets.

Setting

Write ℝⁿ for Euclidean space with norm ‖·‖ and inner product ⟨·,·⟩, and let f : ℝⁿ → ℝ ∪ {+∞} be lower semicontinuous, with domain dom f = {x : f(x) < +∞} and graph Graph f = {(x, f(x)) : x ∈ dom f} ⊆ ℝⁿ⁺¹. Π : ℝⁿ⁺¹ → ℝⁿ forgets the last coordinate.

Subdifferentials. A vector x* is a Fréchet subgradient of f at x ∈ dom f if liminf_{y→x, y≠x} [f(y) − f(x) − ⟨x*, y − x⟩]/‖y − x‖ ≥ 0. The limiting subdifferential ∂f(x) collects the limits of Fréchet subgradients x_k at points x_k → x with f(x_k) → f(x); the singular subdifferential ∂^∞f(x) collects the limits of t_k x_k with t_k ↘ 0⁺. The Clarke subdifferential is

∂∘f(x)=co‾ (∂f(x)+∂∞f(x))  for x∈dom⁡f,∂∘f(x)=∅ otherwise,\partial^\circ f(x)=\overline{\mathrm{co}}\,\big(\partial f(x)+\partial^\infty f(x)\big)\ \text{ for } x\in\operatorname{dom} f,\qquad \partial^\circ f(x)=\emptyset \text{ otherwise},∂∘f(x)=co(∂f(x)+∂∞f(x))  for x∈domf,∂∘f(x)=∅ otherwise,

with co‾\overline{\mathrm{co}}co the closed convex hull. It may be empty even on dom f, for example for −‖x‖^{1/2} at 0.

O-minimal structures. An o-minimal structure 𝒪 on (ℝ, +, ·) is a sequence of Boolean algebras 𝒪ₙ of subsets of ℝⁿ (the definable sets) that is stable under A ↦ A × ℝ, A ↦ ℝ × A and the projection Π, contains every algebraic set {p = 0}, and whose one-dimensional sets are exactly the finite unions of intervals and points. Semialgebraic sets, the globally subanalytic sets and the sets definable with the exponential each form such a structure. A function is definable if its graph is.

Stratifications. A C^p stratification of a set X is a locally finite partition of X into C^p submanifolds (strata) such that a stratum meeting the closure of another lies in its frontier. It is Whitney-(a) if tangent spaces T_{x_k}X_i converging to 𝒯 along x_k → x ∈ X_j satisfy T_xX_j ⊆ 𝒯, and a stratification of a set in ℝⁿ⁺¹ is nonvertical if e_{n+1} is tangent to no stratum. For x in a stratum X_i, ∇_R f(x) is the Riemannian gradient of f restricted to X_i.

Formalization targets

Goal: Theorem 14 (nonsmooth Kurdyka–Łojasiewicz inequality)

For every lower semicontinuous definable f there are ρ > 0, a strictly increasing continuous definable ψ : [0, ρ) → ℝ, C¹ on (0, ρ) with ψ(0) = 0, and a continuous definable χ : ℝ₊ → (0, ρ) such that

∥x∗∥ ≥ 1ψ′(∣f(x)∣)whenever 0<∣f(x)∣≤χ(∥x∥), x∗∈∂∘f(x).(22)\|x^*\|\ \ge\ \frac{1}{\psi'(|f(x)|)}\qquad\text{whenever } 0<|f(x)|\le\chi(\|x\|),\ x^*\in\partial^\circ f(x). \tag{22}∥x∗∥ ≥ ψ′(∣f(x)∣)1​whenever 0<∣f(x)∣≤χ(∥x∥), x∗∈∂∘f(x).(22)

Neither ρ, ψ nor χ is fixed: the goal asserts only their existence, so it is independent of any choice of exponent.

Milestones

  1. Lemma 8 — a nonvertical definable C^p-Whitney stratification of Graph f whose projection stratifies dom f compatibly with given definable sets.
  2. Corollary 9, (15) and (i) — on a definable stratification of dom f, Proj⁡TxXx∂∘f(x)⊂{∇Rf(x)}\operatorname{Proj}_{T_xX_x}\partial^\circ f(x)\subset\{\nabla_R f(x)\}ProjTx​Xx​​∂∘f(x)⊂{∇R​f(x)}, so ∥∇Rf(x)∥≤∥x∗∥\|\nabla_R f(x)\|\le\|x^*\|∥∇R​f(x)∥≤∥x∗∥.
  3. Proposition 10 — a definable ψ with ψ(t) ≥ φ(t, s) uniformly in s ∈ [a, +∞), for t ∈ (0, χ(s)).
  4. Theorem 11 — the smooth KL inequality ‖∇(ψ∘f)(x)‖ ≥ 1 for 0 < f(x) ≤ χ(‖x‖) on an unbounded definable submanifold.

Corollary 9 (ii)–(iii) (finitely many Clarke critical and asymptotic critical values) and Corollary 12 (the KL inequality around the zero set) are included as further statements.

Significance

The result. Theorem 14 makes the KL property available, with no further verification, for every lower semicontinuous objective built from semialgebraic, globally subanalytic or exp-definable pieces. Global convergence theorems for proximal alternating minimization, proximal gradient methods, PALM and nonconvex ADMM take a KL inequality as hypothesis; definability is how that hypothesis is checked in applications. The inequality is relative to the value 0, holds globally through χ(‖x‖), and controls every Clarke subgradient rather than only the one of least norm. Corollary 9 is a definable, nonsmooth Morse–Sard theorem.

Formalizing it. The theorem has been proved since 2007; it has not been formalized. Lean's Mathlib has no o-minimal geometry: no cell decomposition, monotonicity theorem, definable choice or Whitney stratification. A complete development of this mission would supply the first machine-checked KL inequality for a general class of nonsmooth functions, and the o-minimal infrastructure it needs is reusable for every result in optimization and real algebraic geometry that cites "tame" functions.

Difficulty

The statements are short; the proofs rest on geometry absent from Lean. Lemma 8 is proved in the paper by citation of a stratification theorem for definable maps ([Shiota, Geometry of Subanalytic and Semialgebraic Sets, 1997, II.1.17]). Proposition 10 and Theorem 11 use the monotonicity lemma for definable functions of one variable and definable selection. The natural first idea — apply Kurdyka's inequality on each stratum and take a minimum — fails twice: Kurdyka's inequality is local on bounded sets, and the reparametrizations ψ_i of different strata must be compared near 0, which needs the monotonicity lemma once more. Passing from the strata to Clarke subgradients needs the projection formula of Corollary 9, which in turn depends on nonverticality and the Whitney-(a) condition.

Formalization scope

ℝⁿ is EuclideanSpace ℝ (Fin n); ℝ ∪ {+∞} is EReal, with f never equal to ⊥ as a standing hypothesis; ℝⁿ⁺¹ carries the value in the last coordinate. The Fréchet and limiting subdifferentials are the published NonconvexSplitting.Shared.IsRegularSubgrad and LimitingSubdiff. The Clarke subdifferential uses closedConvexHull. An o-minimal structure is a structure with a family O n of sets of subsets of ℝⁿ satisfying Definition 6 (Boolean algebra as: ∅, complements, binary unions; one-dimensional sets as finite unions of order-connected sets). Definability of ψ, χ and of the strata is part of every conclusion that asserts it. Tangent spaces are spans of Mathlib's tangentConeAt; C^p submanifolds are local graphs; the Riemannian gradient on a set U is a vector g in the tangent space with HasFDerivWithinAt h ⟪g, ·⟫ U x.

Inequality (22) is stated as ψ′(|f(x)|)·‖x*‖ ≥ 1, never as 1/ψ′ ≤ ‖x*‖, so that ψ′ = 0 cannot make it hold through 1/0 = 0. The goal mentions no strata, no auxiliary functions of the proof and no Clarke critical points; a formalization that adds such hypotheses, drops definability of ψ and χ, or restricts to bounded sets is a different theorem.

Welcome contributions: the elementary closure properties of o-minimal structures (definability of sums, compositions, images), the monotonicity theorem for definable functions of one variable, definable choice, cell decomposition and Whitney stratification — all reusable well beyond this mission.

Selected references

  • J. Bolte, A. Daniilidis, A. Lewis, M. Shiota, Clarke subgradients of stratifiable functions, SIAM J. Optim. 18(2) (2007), 556–572. https://doi.org/10.1137/060670080
  • K. Kurdyka, On gradients of functions definable in o-minimal structures, Ann. Inst. Fourier 48 (1998), 769–783. https://doi.org/10.5802/aif.1638
  • J. Bolte, A. Daniilidis, A. Lewis, The Łojasiewicz inequality for nonsmooth subanalytic functions with applications to subgradient dynamical systems, SIAM J. Optim. 17(4) (2007), 1205–1223. https://doi.org/10.1137/050644641
  • L. van den Dries, C. Miller, Geometric categories and o-minimal structures, Duke Math. J. 84 (1996), 497–540. https://doi.org/10.1215/S0012-7094-96-08416-1
  • M. Coste, An Introduction to O-minimal Geometry, Istituti Editoriali e Poligrafici Internazionali, Pisa, 2000. https://perso.univ-rennes1.fr/michel.coste/polyens/OMIN.pdf
7 thms1 active userReviewed
Functional Analysis·Captain: marwahaha

A counterexample to Naimark's problem in ZFCOpen Problem

Motivation

Uniqueness of irreducible representations is a strong condition on a C*-algebra. The target asks for a counterexample to the expected compact-operator characterization without added set-theoretic hypotheses. The pinned manuscript supplies the research context.

Setting

Representations are nonzero star homomorphisms on complete complex inner product spaces. Irreducibility is expressed by closed invariant subspaces, and equivalence by an intertwining linear isometry.

Formalization target

The selected goal is OAI.NaimarkZFC.main. Its central assertion is

∃A:simple(A)∧unique irreducible representation(A)∧A≇K(H).\exists A:\quad\mathrm{simple}(A)\land\mathrm{unique\ irreducible\ representation}(A)\land A\not\cong\mathcal K(H).∃A:simple(A)∧unique irreducible representation(A)∧A≅K(H).

The theorem states that, for every universe level choice, there exists a type A carrying a C*-algebra structure (over ℂ) with the following five properties. First, A is simple in the sense that it is nontrivial and its only closed two-sided ideals are {0} and A itself. Second, A is not finite-dimensional as a complex vector space. Third, there is a continuous linear functional τ : A → ℂ that is a faithful tracial state: τ has norm 1, τ(aa) is a nonnegative real number for every a, τ(ab) = τ(ba) for all a and b, and τ(aa) > 0 whenever a ≠ 0. Fourth, A has a unique irreducible representation up to unitary equivalence: whenever π and ρ are nonzero star-homomorphisms (non-unital algebra homomorphisms preserving star) from A into the bounded operators on complete complex inner product spaces H and K, taken in the stated universe levels, and each has no closed invariant subspace other than 0 and the whole space, there is a linear isometric isomorphism U : H → K with U(π(a)x) = ρ(a)(Ux) for all a and x. Fifth, for every complete complex inner product space H in the stated universe, A is not isomorphic to the algebra of compact operators on H, meaning there is no injective star-homomorphism φ : A → B(H) whose range is exactly the set of compact operators on H.

Significance and status

The target explicitly requires simplicity, infinite dimension, a faithful tracial state and failure of all compact-operator isomorphisms in the stated universes. No additional set-theoretic hypothesis appears in its binders. The target is currently Open on Prove2Me. The manuscript's mathematical argument and a machine-checked proof of the selected statement are separate deliverables.

Difficulty

The construction must combine infinite dimension, a faithful trace and uniqueness across the specified representation universes.

Formalization scope

The exact published goal, its hypotheses and its referenced definition blocks specify the requested formalization. The explanatory formula above is a summary; all quantifiers and additional clauses in the linked statement remain required.

The available published material supplies this goal and its necessary definitions. Additional manuscript lemmas are not represented as attached milestones.

Selected references

  • OpenAI, A counterexample to Naimark's problem in ZFC, preprint, 2026. Manuscript.
  • OpenAI, accompanying formal statements, commit adc7f1241b42. Selected goal source.
2 thms1 active userReviewed
Captain: marwahaha

The Partition Principle does not imply ChoiceOpen Problem

Motivation

The Partition Principle reverses the cardinal comparison provided by a surjection without requiring a section. The source studies whether this weaker form of comparison can hold while full Choice fails. The source is OpenAI's September 2026 manuscript.

Setting

The Partition Principle says that a surjection from XXX onto YYY implies an injection from YYY into XXX. Ordinal-indexed Choice provides choices for families indexed by internal ordinals. The formal consistency predicate rules out finite derivations of falsity in a specified first-order proof system.

Formalization targets

Con⁡(ZF)⟹Con⁡(ZF+PP+ACwo+¬AC).\operatorname{Con}(\mathrm{ZF})\quad\Longrightarrow\quad\operatorname{Con}(\mathrm{ZF}+\mathrm{PP}+\mathrm{AC}_{\mathrm{wo}}+\neg\mathrm{AC}).Con(ZF)⟹Con(ZF+PP+ACwo​+¬AC).

A relative consistency result in pure first-order set theory with membership and equality: if ZF, meaning the full axioms of extensionality, empty set, pairing, union, power set, infinity and foundation together with every instance of Separation and Replacement for arbitrary formulas with arbitrary parameters, and without Choice, is consistent, then so is the theory obtained by adding three sentences to ZF. Consistency is syntactic: no finite classical natural-deduction derivation of falsity exists from the theory, allowing any finite supply of free variables. The added sentences are the Partition Principle PP, which says that whenever there is a set surjection from X onto Y then there is a set injection from Y into X, with functions coded internally as sets of Kuratowski pairs and the injection not required to be a section of the surjection; the ordinal-indexed Choice axiom ACwo, which says that for every internal von Neumann ordinal (transitive and internally well-ordered by membership) every family of nonempty sets indexed by it has a choice function; and the negation of full Axiom of Choice AC, which says that every set of nonempty sets has a choice function.

The goal is OAI.PartitionConsistency.main_consistency. Supporting targets are listed below; they retain their individual hypotheses and are separate statements.

  • exists model partitionPrinciple without choice.

Significance

The chosen goal is the relative syntactic consistency claim. A separate attached model-construction target gives a transitive extension under its stronger ground-model assumptions. The manuscript presents a mathematical argument; the attached goal currently has status Open, so the requested contribution is a Lean proof of the displayed formal target.

Difficulty

The function and ordinal notions must be expressed internally in the set-theoretic language. An ambient use of classical choice cannot substitute for satisfaction of the internal Choice sentence.

Formalization scope

The source theory includes every Separation and Replacement instance with parameters. The auxiliary model target assumes a countable transitive ground and a standard inaccessible in that ground; it is not presented as an assumption-free construction over arbitrary countable transitive models.

The shared definitions are supplied by PartitionConsistency, PartitionPrinciple. The target uses Lean 4.33.1 with the pinned Mathlib environment. Contributions should preserve the existing statement and develop reusable supporting lemmas in its definitions.

Selected references

  • OpenAI, The Partition Principle does not imply Choice, preprint, September 2026. Manuscript.
  • OpenAI, accompanying Lean statement. Pinned source.
4 thms1 active userReviewed
Algebraic TopologyGroup Theory·Captain: wurtle

Simple F∞ overgroups of groups with decidable word problemResearch Paper

Motivation

A finitely generated group has decidable word problem if an algorithm decides, for every word in a fixed finite generating set and its inverses, whether it represents the identity. Embedding theorems translate this algorithmic property into algebra: Higman (1961) characterized subgroups of finitely presented groups, and Boone and Higman (1974) showed that decidable word problem is equivalent to embedding in a simple subgroup of a finitely presented group, asking whether the simple group can be finitely presented (the Boone–Higman conjecture).

This preprint proves a stronger statement in the direction of higher finiteness: every finitely generated group with decidable word problem embeds in a simple group of type F∞F_\inftyF∞​, i.e. one with a classifying space having finitely many cells in each dimension. Since type F∞F_\inftyF∞​ implies finite presentation, this implies the Boone–Higman conjecture.

Background

  • 1961. Higman's embedding theorem: finitely generated subgroups of finitely presented groups are exactly the recursively presented groups.
  • 1974. Boone and Higman characterize decidable word problem by embeddings into simple subgroups of finitely presented groups and pose the conjecture.
  • 1975, 1987. Brown gives homological criteria for finiteness and establishes type F∞F_\inftyF∞​ for the Thompson–Higman groups.
  • 1980. Thompson shows the simple overgroup can be chosen finitely generated with decidable word problem.
  • 2022. Belk and Zaremsky introduce twisted Brin–Thompson groups and construct a simple group of type F∞F_\inftyF∞​ containing all finitely generated right-angled Artin groups.
  • 2024–2026. Belk, Hyde and Matucci construct a simple F∞F_\inftyF∞​ group containing every countable abelian group; the Boone–Higman conjecture is verified for hyperbolic groups (Belk, Bleak, Matucci, Zaremsky) and for Aut(Fn)\mathrm{Aut}(F_n)Aut(Fn​) and mapping class groups (Belk, Fournier-Facio, Hyde, Zaremsky). The survey of Belk, Bleak, Matucci and Zaremsky raises the higher-finiteness question.
  • 2026. The OpenAI preprint Simple F∞F_\inftyF∞​ overgroups of groups with decidable word problem (dated September 23, 2026) proves the theorem below, building on the companion preprint Finite algebraic envelopes and the Boone–Higman conjecture. It has not been peer reviewed; the Lean goal is open on this platform.

Setting

A classifying space K(H,1)K(H,1)K(H,1) is a connected CW complex with fundamental group HHH and contractible universal cover. HHH has type FnF_nFn​ if some K(H,1)K(H,1)K(H,1) has finite nnn-skeleton and type F∞F_\inftyF∞​ if some K(H,1)K(H,1)K(H,1) has finitely many cells in each dimension. Type F1F_1F1​ is finite generation and type F2F_2F2​ is finite presentation. A group is simple if it is nontrivial and has no normal subgroups besides 111 and itself.

In Lean (OAI.SimpleFInftyOvergroups, definitions SimpleOvergroups), HasDecidableWordProblem G asks for a finite generating tuple with a ComputablePred identity predicate on words; HasTypeFInfty H asks for a Hausdorff connected space XXX with a CW structure of finite type (Topology.CWComplex.FiniteType), a basepoint with H≃π1(X,x)H \simeq \pi_1(X, x)H≃π1​(X,x), and a surjective covering map from a contractible space.

Formalization targets

Goal: Theorem 1.1

Every finitely generated group GGG with decidable word problem admits an injective homomorphism G↪HG \hookrightarrow HG↪H into a nontrivial simple group HHH of type F∞F_\inftyF∞​.

The input need not be finitely presented and its word-problem algorithm may have arbitrary running time; the classifying space of HHH may be infinite-dimensional.

Significance

The result. Theorem 1.1 answers the higher-finiteness strengthening of the Boone–Higman conjecture and implies the conjecture itself. With the classical converse (finitely presented simple groups have decidable word problem), it characterizes groups with decidable word problem as the finitely generated subgroups of simple groups of type F∞F_\inftyF∞​.

Formalizing it. No machine-checked proof exists; the source is an unrefereed preprint. A formalization requires CW complexes of finite type, classifying spaces and covering theory (partially in Mathlib), Brown's homological finiteness criterion, Steinberg groups, and twisted Brin–Thompson groups with the Belk–Zaremsky finiteness theorem (Theorem 5.1 in the preprint, from the literature).

Difficulty

Finite presentation already fails to pass to subgroups; higher finiteness additionally requires controlling relations among relations in every dimension. Known families of simple F∞F_\inftyF∞​ groups contain large classes of groups but give no general construction from a word-problem algorithm. The proof needs an oligomorphic faithful action of a group of type F∞F_\inftyF∞​ with stabilizers of type F∞F_\inftyF∞​, built from an algebraic envelope whose Steinberg-group kernel must be annihilated exactly (Theorem 4.6).

Formalization scope

  • GGG and HHH live in the same universe Type u; GGG carries [Group.FG G].
  • IsSimpleGroup H together with Nontrivial H encodes simplicity.
  • Type F∞F_\inftyF∞​ is encoded topologically via a finite-type CW structure on a connected Hausdorff space whose universal cover (a surjective covering with contractible total space) witnesses asphericity.
  • The word problem uses ComputablePred on lists of (generator, sign) pairs for a specified generating tuple.

The statement is not vacuous: the hypothesis holds for many groups (e.g. all finitely presented residually finite groups), and the conclusion requires a genuine embedding.

Needed infrastructure: finite-type CW complexes and K(π,1)K(\pi,1)K(π,1) spaces, Brown's criterion, group actions with finiteness properties, Steinberg groups over noncommutative algebras, and twisted Brin–Thompson groups. Contributions toward any of these are welcome.

Selected references

  • OpenAI, Simple F∞F_\inftyF∞​ overgroups of groups with decidable word problem, OpenAI Math Release preprint, September 23, 2026. https://github.com/openai/math/blob/adc7f1241b42e322a6451854ab7e4b4c146bf78a/preprints/Simple-F-infinity-overgroups-of-groups-with-decidable-word-problem-September-23-2026/paper.pdf
  • OpenAI, Finite algebraic envelopes and the Boone–Higman conjecture, OpenAI Math Release preprint, September 23, 2026. https://github.com/openai/math/blob/main/preprints/Finite-algebraic-envelopes-and-the-Boone-Higman-conjecture-September-23-2026/paper.pdf
  • W. W. Boone, G. Higman, An algebraic characterization of groups with soluble word problem, J. Austral. Math. Soc. 18 (1974). https://doi.org/10.1017/S1446788700019108
  • G. Higman, Subgroups of finitely presented groups, Proc. Roy. Soc. London Ser. A 262 (1961). https://doi.org/10.1098/rspa.1961.0132
  • R. J. Thompson, Embeddings into finitely generated simple groups which preserve the word problem, in Word Problems II, 1980. https://doi.org/10.1016/S0049-237X(08)71348-X
  • K. S. Brown, Homological criteria for finiteness, Comment. Math. Helv. 50 (1975). https://doi.org/10.1007/BF02565740
  • K. S. Brown, Finiteness properties of groups, J. Pure Appl. Algebra 44 (1987). https://doi.org/10.1016/0022-4049(87)90015-6
  • J. Belk, M. C. B. Zaremsky, Twisted Brin–Thompson groups, Geom. Topol. 26 (2022). https://doi.org/10.2140/gt.2022.26.1189
  • J. Belk, C. Bleak, F. Matucci, M. C. B. Zaremsky, Progress around the Boone–Higman conjecture, 2025. https://arxiv.org/abs/2306.16356
  • J. Belk, C. Bleak, F. Matucci, M. C. B. Zaremsky, Hyperbolic groups satisfy the Boone–Higman conjecture, Duke Math. J. 175 (2026). https://doi.org/10.1215/00127094-2025-0055
  • J. Belk, J. Hyde, F. Matucci, Finite germ extensions, 2024. https://arxiv.org/abs/2407.03149
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Theoretical Computer Science·Captain: wurtle

Weak and strong normalization in pure type systemsResearch Paper

Motivation: does having a normal form imply termination?

Pure type systems (PTSs) are a uniform framework covering the simply typed λ-calculus, System F, the calculus of constructions and the type theories behind proof assistants such as Coq and Lean. A system is specified by its sorts (like ∗\ast∗ and □\square□), axioms (∗:□\ast:\square∗:□) and product rules (which dependent function types may be formed). Normalization is the key metatheoretic property: it underlies logical consistency and decidability of type checking. There are two versions. A term is weakly normalizing if some sequence of β-reductions reaches a normal form, and strongly normalizing if every reduction sequence is finite. Strong normalization is what implementations need (any evaluation strategy terminates) but it is much harder to prove. The Barendregt–Geuvers–Klop conjecture asserts that, for pure type systems, the system-wide weak property already implies the strong one.

Timeline

  • 1989 — Girard's Proofs and Types presents the Tait–Girard reducibility method for strong normalization (book).
  • 1993 — Geuvers states the system-wide conjecture for β and βη reduction in his thesis (Logics and Type Systems, Nijmegen, Conjecture 8.1.2) (thesis).
  • 1997 — Sørensen proves the implication for generalized nondependent, clean, negatable systems via continuation-passing translations (thesis, Copenhagen, Theorem 3.5.20).
  • 2001 — Barthe, Hatcliff and Sørensen prove a uniform nondependent version (TCS 2001).
  • 2006 — Barthe and Coquand show that erasing λ-domain annotations can identify non-convertible terms in non-normalizing PTSs (JFP 2006), a constraint on erasure-based approaches.
  • 2014 — Roux and van Doorn develop a structural theory of PTSs, preserving weak normalization under sums and added rules (RTA–TLCA 2014).
  • 2023 — Mull reduces the conjecture within tiered systems to an irrelevance-eliminating translation (TYPES 2022, LIPIcs 2023) and extends the nondependent results in his thesis (Chicago, 2023).
  • 2025 — Roux proposes internal strong-normalization proofs for specific systems (TYPES 2025 abstract).
  • 2026 — An OpenAI preprint, Weak and strong normalization in pure type systems (OpenAI Math Release, September 25, 2026), claims the β-version of the conjecture for every PTS, with no functionality assumption. The preprint has not been peer reviewed and its theorem is not formally verified.

Setting

A specification consists of a type SSS of sorts, an axiom relation A⊆S2\mathcal A\subseteq S^2A⊆S2 and a rule relation R⊆S3\mathcal R\subseteq S^3R⊆S3; neither is assumed functional. Expressions are

M,N,A,B::=x∣s∣MN∣λx:A. M∣Πx:A. B,s∈S,M,N,A,B ::= x\mid s\mid MN\mid \lambda x{:}A.\,M\mid \Pi x{:}A.\,B,\qquad s\in S,M,N,A,B::=x∣s∣MN∣λx:A.M∣Πx:A.B,s∈S,

up to renaming of bound variables. β-reduction is the compatible closure of (λx:A.M)N→βM[x:=N](\lambda x{:}A.M)N\to_\beta M[x:=N](λx:A.M)N→β​M[x:=N], with steps allowed everywhere, including inside the annotation AAA of a λ and both parts of a Π. Typing Γ⊢M:A\Gamma\vdash M:AΓ⊢M:A is generated by the seven standard PTS rules (axiom, variable, weakening, product, abstraction, application, conversion up to =β=_\beta=β​). A context is valid if each declared type is typed by a sort in the preceding prefix, and MMM is legal in Γ\GammaΓ if Γ⊢M:A\Gamma\vdash M:AΓ⊢M:A or Γ⊢A:M\Gamma\vdash A:MΓ⊢A:M for some AAA.

The system is weakly (strongly) normalizing if every legal expression in every valid context is weakly (strongly) normalizing.

Formalization targets

Goal: β-Barendregt–Geuvers–Klop (Theorem 1)

For every specification P=(S,A,R)\mathcal P=(S,\mathcal A,\mathcal R)P=(S,A,R),

(∀Γ valid, ∀M legal in Γ: M is weakly β-normalizing) ⟹ (∀Γ valid, ∀M legal in Γ: M is strongly β-normalizing).\bigl(\forall \Gamma\ \text{valid},\ \forall M\ \text{legal in }\Gamma:\ M\ \text{is weakly }\beta\text{-normalizing}\bigr) \ \Longrightarrow\ \bigl(\forall \Gamma\ \text{valid},\ \forall M\ \text{legal in }\Gamma:\ M\ \text{is strongly }\beta\text{-normalizing}\bigr).(∀Γ valid, ∀M legal in Γ: M is weakly β-normalizing) ⟹ (∀Γ valid, ∀M legal in Γ: M is strongly β-normalizing).

Both quantifiers range over the whole system; the theorem does not claim that an individual weakly normalizing term is strongly normalizing, and it says nothing about η-reduction. The goal statement is published on the platform with status Open: no machine-checked proof exists yet.

Significance

The result itself. For any PTS, a proof that every legal term has a normal form now yields termination of every β-reduction strategy, including reductions inside types and annotations that may continue after the computational body is normal. Earlier results needed nondependent products, cleanliness, negatability, a tiered sort hierarchy, or a translation requiring extra product rules; the general case — arbitrary sorts and nonfunctional rules — was not covered by any earlier result since Geuvers' 1993 formulation.

Formalizing it. The statement is entirely syntactic and finitary, so it is a natural target for a proof assistant, and the relevant objects (de Bruijn syntax, substitution, typing derivations) are standard formalization material. The proof in the source is long (over 70 pages) and combines reducibility candidates with a fixed-point construction over sort profiles; a machine check would substantially raise confidence. A formal PTS metatheory library (substitution lemmas, confluence, subject reduction) would be reusable for other type-theory formalizations.

Difficulty

The obvious route — reducibility candidates à la Tait–Girard — needs a model of types built by recursion on sorts, which fails for impredicative or nonfunctional specifications where a type can have several sorts and no hierarchy is available. Translation methods (CPS or irrelevance elimination) need specific product rules to preserve typing, which a general specification may lack. Weak normalization of one term cannot be used alone; the argument must exploit normal forms of other open expressions. Annotations add a further complication: erasing them can change the equational theory (Barthe–Coquand), so reduction inside annotations must be tracked rather than erased.

Formalization scope

  • Expressions use de Bruijn indices (Expr S with var, sort, app, lam, pi), so α-equivalence is syntactic; substitution is parallel substitution with lifting.
  • Beta is full compatible closure, including lam_domain, pi_domain and pi_body steps; conversion is Relation.EqvGen Beta.
  • HasType has exactly the seven PTS rules; contexts are lists stored newest-first, each declaration typed in its prefix (ValidContext). Legal includes expressions occurring as types.
  • Weak normalization is reduction to a normal form; strong normalization is accessibility for the converse of Beta.
  • The sort type S is arbitrary (any universe), and axioms and rules are arbitrary relations: no functionality, finiteness or hierarchy assumptions.
  • Needed infrastructure: renaming/substitution lemmas, confluence of β, subject reduction for PTSs, reducibility candidates. All are reusable for any PTS formalization.

Selected references

  • J. H. Geuvers, Logics and Type Systems, PhD thesis, Katholieke Universiteit Nijmegen, 1993. https://www.cs.ru.nl/~herman/PUBS/Proefschrift.pdf
  • J.-Y. Girard, Proofs and Types, Cambridge University Press, 1989. https://www.paultaylor.eu/stable/prot.pdf
  • M. H. B. Sørensen, Normalization in λ-Calculus and Type Theory, PhD thesis, University of Copenhagen, 1997. https://di.ku.dk/forskning/Publikationer/tekniske_rapporter/tekniske-rapporter-1997/97-27.pdf
  • G. Barthe, J. Hatcliff and M. H. Sørensen, Weak normalization implies strong normalization in a class of non-dependent pure type systems, Theoret. Comput. Sci. (2001). https://doi.org/10.1016/S0304-3975(01)00012-3
  • G. Barthe and T. Coquand, Remarks on the equational theory of non-normalizing pure type systems, J. Funct. Programming (2006). https://doi.org/10.1017/S0956796803004726
  • C. Roux and F. van Doorn, The structural theory of pure type systems, RTA–TLCA 2014. https://doi.org/10.1007/978-3-319-08918-8_25
  • N. Mull, An irrelevancy-eliminating translation of pure type systems, TYPES 2022, LIPIcs, 2023. https://doi.org/10.4230/LIPIcs.TYPES.2022.7
  • OpenAI, Weak and strong normalization in pure type systems, OpenAI Math Release preprint, September 25, 2026 (source of the goal; Theorem 1, p. 3). https://github.com/openai/math/blob/adc7f1241b42e322a6451854ab7e4b4c146bf78a/preprints/Weak-and-strong-normalization-in-pure-type-systems-September-25-2026/paper.pdf
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