# Curry–Howard correspondence

In programming language theory and proof theory, the Curry–Howard correspondence (also known as the Curry–Howard isomorphism or equivalence, or the proofs-as-programs and propositions- or formulae-as-types interpretation) is the direct relationship between computer programs and mathematical proofs.

It is a generalization of a syntactic analogy between systems of formal logic and computational calculi that was first discovered by the American mathematician Haskell Curry and the logician William Alvin Howard.[1] It is the link between logic and computation that is usually attributed to Curry and Howard, although the idea is related to the operational interpretation of intuitionistic logic given in various formulations by L. E. J. Brouwer, Arend Heyting and Andrey Kolmogorov (see Brouwer–Heyting–Kolmogorov interpretation)[2] and Stephen Kleene (see Realizability). The relationship has been extended to include category theory as the three-way Curry–Howard–Lambek correspondence.[3][4][5]

## Origin, scope, and consequences

The beginnings of the Curry–Howard correspondence lie in several observations:

1. In 1934 Curry observes that the types of the combinators could be seen as axiom-schemes for intuitionistic implicational logic.[6]
2. In 1958 he observes that a certain kind of proof system, referred to as Hilbert-style deduction systems, coincides on some fragment to the typed fragment of a standard model of computation known as combinatory logic.[7]
3. In 1969 Howard observes that another, more "high-level" proof system, referred to as natural deduction, can be directly interpreted in its intuitionistic version as a typed variant of the model of computation known as lambda calculus.[8]

The Curry–Howard correspondence is the observation that there is an isomorphism between the proof systems, and the models of computation. It is the statement that these two families of formalisms can be considered as identical.

If one abstracts on the peculiarities of either formalism, the following generalization arises: a proof is a program, and the formula it proves is the type for the program. More informally, this can be seen as an analogy that states that the return type of a function (i.e., the type of values returned by a function) is analogous to a logical theorem, subject to hypotheses corresponding to the types of the argument values passed to the function; and that the program to compute that function is analogous to a proof of that theorem. This sets a form of logic programming on a rigorous foundation: proofs can be represented as programs, and especially as lambda terms, or proofs can be run.

The correspondence has been the starting point of a large spectrum of new research after its discovery, leading in particular to a new class of formal systems designed to act both as a proof system and as a typed functional programming language. This includes Martin-Löf's intuitionistic type theory and Coquand's Calculus of Constructions, two calculi in which proofs are regular objects of the discourse and in which one can state properties of proofs the same way as of any program. This field of research is usually referred to as modern type theory.

Such typed lambda calculi derived from the Curry–Howard paradigm led to software like Coq in which proofs seen as programs can be formalized, checked, and run.

A converse direction is to use a program to extract a proof, given its correctness—an area of research closely related to proof-carrying code. This is only feasible if the programming language the program is written for is very richly typed: the development of such type systems has been partly motivated by the wish to make the Curry–Howard correspondence practically relevant.

The Curry–Howard correspondence also raised new questions regarding the computational content of proof concepts that were not covered by the original works of Curry and Howard. In particular, classical logic has been shown to correspond to the ability to manipulate the continuation of programs and the symmetry of sequent calculus to express the duality between the two evaluation strategies known as call-by-name and call-by-value.

Because of the possibility of writing non-terminating programs, Turing-complete models of computation (such as languages with arbitrary recursive functions) must be interpreted with care, as naive application of the correspondence leads to an inconsistent logic. The best way of dealing with arbitrary computation from a logical point of view is still an actively debated research question, but one popular approach is based on using monads to segregate provably terminating from potentially non-terminating code (an approach that also generalizes to much richer models of computation,[9] and is itself related to modal logic by a natural extension of the Curry–Howard isomorphism[ext 1]). A more radical approach, advocated by total functional programming, is to eliminate unrestricted recursion (and forgo Turing completeness, although still retaining high computational complexity), using more controlled corecursion wherever non-terminating behavior is actually desired.

## General formulation

In its more general formulation, the Curry–Howard correspondence is a correspondence between formal proof calculi and type systems for models of computation. In particular, it splits into two correspondences. One at the level of formulas and types that is independent of which particular proof system or model of computation is considered, and one at the level of proofs and programs which, this time, is specific to the particular choice of proof system and model of computation considered.

At the level of formulas and types, the correspondence says that implication behaves the same as a function type, conjunction as a "product" type (this may be called a tuple, a struct, a list, or some other term depending on the language), disjunction as a sum type (this type may be called a union), the false formula as the empty type and the true formula as the unit type (whose sole member is the null object). Quantifiers correspond to dependent function space or products (as appropriate). This is summarized in the following table:

Logic side Programming side
universal quantification generalised product type (Π type)
existential quantification generalised sum type (Σ type)
implication function type
conjunction product type
disjunction sum type
true formula unit type or top type
false formula empty type or bottom type

At the level of proof systems and models of computations, the correspondence mainly shows the identity of structure, first, between some particular formulations of systems known as Hilbert-style deduction system and combinatory logic, and, secondly, between some particular formulations of systems known as natural deduction and lambda calculus.

Logic side Programming side
Hilbert-style deduction system type system for combinatory logic
natural deduction type system for lambda calculus

Between the natural deduction system and the lambda calculus there are the following correspondences:

Logic side Programming side
hypotheses free variables
implication elimination (modus ponens) application
implication introduction abstraction

## Corresponding systems

### Intuitionistic Hilbert-style deduction systems and typed combinatory logic

It was at the beginning a simple remark in Curry and Feys's 1958 book on combinatory logic: the simplest types for the basic combinators K and S of combinatory logic surprisingly corresponded to the respective axiom schemes α → (βα) and (α → (βγ)) → ((αβ) → (αγ)) used in Hilbert-style deduction systems. For this reason, these schemes are now often called axioms K and S. Examples of programs seen as proofs in a Hilbert-style logic are given below.

If one restricts to the implicational intuitionistic fragment, a simple way to formalize logic in Hilbert's style is as follows. Let Γ be a finite collection of formulas, considered as hypotheses. Then δ is derivable from Γ, denoted Γ ⊢ δ, in the following cases:

• δ is an hypothesis, i.e. it is a formula of Γ,
• δ is an instance of an axiom scheme; i.e., under the most common axiom system:
• δ has the form α → (βα), or
• δ has the form (α → (βγ)) → ((αβ) → (αγ)),
• δ follows by deduction, i.e., for some α, both αδ and α are already derivable from Γ (this is the rule of modus ponens)

This can be formalized using inference rules, as in the left column of the following table.

Typed combinatory logic can be formulated using a similar syntax: let Γ be a finite collection of variables, annotated with their types. A term T (also annotated with its type) will depend on these variables [Γ ⊢ T:δ] when:

• T is one of the variables in Γ,
• T is a basic combinator; i.e., under the most common combinator basis:
• T is K:α → (βα) [where α and β denote the types of its arguments], or
• T is S:(α → (βγ)) → ((αβ) → (αγ)),
• T is the composition of two subterms which depend on the variables in Γ.

The generation rules defined here are given in the right-column below. Curry's remark simply states that both columns are in one-to-one correspondence. The restriction of the correspondence to intuitionistic logic means that some classical tautologies, such as Peirce's law ((αβ) → α) → α, are excluded from the correspondence.

Hilbert-style intuitionistic implicational logic Typed combinatory logic
${\displaystyle {\frac {\alpha \in \Gamma }{\Gamma \vdash \alpha }}\qquad \qquad {\text{Assum}}}$ ${\displaystyle {\frac {x:\alpha \in \Gamma }{\Gamma \vdash x:\alpha }}}$
${\displaystyle {\frac {}{\Gamma \vdash \alpha \rightarrow (\beta \rightarrow \alpha )}}\qquad {\text{Ax}}_{K}}$ ${\displaystyle {\frac {}{\Gamma \vdash K:\alpha \rightarrow (\beta \rightarrow \alpha )}}}$
${\displaystyle {\frac {}{\Gamma \vdash (\alpha \!\rightarrow \!(\beta \!\rightarrow \!\gamma ))\!\rightarrow \!((\alpha \!\rightarrow \!\beta )\!\rightarrow \!(\alpha \!\rightarrow \!\gamma ))}}\;{\text{Ax}}_{S}}$ ${\displaystyle {\frac {}{\Gamma \vdash S:(\alpha \!\rightarrow \!(\beta \!\rightarrow \!\gamma ))\!\rightarrow \!((\alpha \!\rightarrow \!\beta )\!\rightarrow \!(\alpha \!\rightarrow \!\gamma ))}}}$
${\displaystyle {\frac {\Gamma \vdash \alpha \rightarrow \beta \qquad \Gamma \vdash \alpha }{\Gamma \vdash \beta }}\quad {\text{Modus Ponens}}}$ ${\displaystyle {\frac {\Gamma \vdash E_{1}:\alpha \rightarrow \beta \qquad \Gamma \vdash E_{2}:\alpha }{\Gamma \vdash E_{1}\;E_{2}:\beta }}}$

Seen at a more abstract level, the correspondence can be restated as shown in the following table. Especially, the deduction theorem specific to Hilbert-style logic matches the process of abstraction elimination of combinatory logic.

Logic side Programming side
assumption variable
axiom schemes combinators
modus ponens application
deduction theorem abstraction elimination

Thanks to the correspondence, results from combinatory logic can be transferred to Hilbert-style logic and vice versa. For instance, the notion of reduction of terms in combinatory logic can be transferred to Hilbert-style logic and it provides a way to canonically transform proofs into other proofs of the same statement. One can also transfer the notion of normal terms to a notion of normal proofs, expressing that the hypotheses of the axioms never need to be all detached (since otherwise a simplification can happen).

Conversely, the non provability in intuitionistic logic of Peirce's law can be transferred back to combinatory logic: there is no typed term of combinatory logic that is typable with type

((αβ) → α) → α.

Results on the completeness of some sets of combinators or axioms can also be transferred. For instance, the fact that the combinator X constitutes a one-point basis of (extensional) combinatory logic implies that the single axiom scheme

(((α → (βγ)) → ((αβ) → (αγ))) → ((δ → (εδ)) → ζ)) → ζ,

which is the principal type of X, is an adequate replacement to the combination of the axiom schemes

α → (βα) and
(α → (βγ)) → ((αβ) → (αγ)).

### Intuitionistic Natural deduction and typed lambda calculus

After Curry emphasized the syntactic correspondence between intuitionistic Hilbert-style deduction and typed combinatory logic, Howard made explicit in 1969 a syntactic analogy between the programs of simply typed lambda calculus and the proofs of natural deduction. Below, the left-hand side formalizes intuitionistic implicational natural deduction as a calculus of sequents (the use of sequents is standard in discussions of the Curry–Howard isomorphism as it allows the deduction rules to be stated more cleanly) with implicit weakening and the right-hand side shows the typing rules of lambda calculus. In the left-hand side, Γ, Γ1 and Γ2 denote ordered sequences of formulas while in the right-hand side, they denote sequences of named (i.e., typed) formulas with all names different.

Intuitionistic implicational natural deduction Lambda calculus type assignment rules
${\displaystyle {\frac {}{\Gamma _{1},\alpha ,\Gamma _{2}\vdash \alpha }}{\text{Ax}}}$ ${\displaystyle {\frac {}{\Gamma _{1},x:\alpha ,\Gamma _{2}\vdash x:\alpha }}}$
${\displaystyle {\frac {\Gamma ,\alpha \vdash \beta }{\Gamma \vdash \alpha \rightarrow \beta }}\rightarrow I}$ ${\displaystyle {\frac {\Gamma ,x:\alpha \vdash t:\beta }{\Gamma \vdash (\lambda x\!:\!\alpha .~t):\alpha \rightarrow \beta }}}$
${\displaystyle {\frac {\Gamma \vdash \alpha \rightarrow \beta \qquad \Gamma \vdash \alpha }{\Gamma \vdash \beta }}\rightarrow E}$ ${\displaystyle {\frac {\Gamma \vdash t:\alpha \rightarrow \beta \qquad \Gamma \vdash u:\alpha }{\Gamma \vdash t\;u:\beta }}}$

To paraphrase the correspondence, proving Γ ⊢ α means having a program that, given values with the types listed in Γ, manufactures an object of type α. An axiom/hypothesis corresponds to the introduction of a new variable with a new, unconstrained type, the → I rule corresponds to function abstraction and the → E rule corresponds to function application. Observe that the correspondence is not exact if the context Γ is taken to be a set of formulas as, e.g., the λ-terms λxy.x and λxy.y of type ααα would not be distinguished in the correspondence. Examples are given below.

Howard showed that the correspondence extends to other connectives of the logic and other constructions of simply typed lambda calculus. Seen at an abstract level, the correspondence can then be summarized as shown in the following table. Especially, it also shows that the notion of normal forms in lambda calculus matches Prawitz's notion of normal deduction in natural deduction, from which it follows that the algorithms for the type inhabitation problem can be turned into algorithms for deciding intuitionistic provability.

Logic side Programming side
axiom/hypothesis variable
introduction rule constructor
elimination rule destructor
normal deduction normal form
normalisation of deductions weak normalisation
provability type inhabitation problem
intuitionistic tautology universally inhabited type

Howard's correspondence naturally extends to other extensions of natural deduction and simply typed lambda calculus. Here is a non-exhaustive list:

• Girard-Reynolds System F as a common language for both second-order propositional logic and polymorphic lambda calculus,
• higher-order logic and Girard's System Fω
• inductive types as algebraic data type
• necessity ${\displaystyle \Box }$ in modal logic and staged computation[ext 2]
• possibility ${\displaystyle \Diamond }$ in modal logic and monadic types for effects[ext 1]
• The λI calculus (where abstraction is restricted to λx.E where x has at least one free occurrence in E) and CLI calculus correspond to relevant logic.[10]
• The local truth (∇) modality in Grothendieck topology or the equivalent "lax" modality (◯) of Benton, Bierman, and de Paiva (1998) correspond to CL-logic describing "computation types".[11]

### Classical logic and control operators

At the time of Curry, and also at the time of Howard, the proofs-as-programs correspondence concerned only intuitionistic logic, i.e. a logic in which, in particular, Peirce's law was not deducible. The extension of the correspondence to Peirce's law and hence to classical logic became clear from the work of Griffin on typing operators that capture the evaluation context of a given program execution so that this evaluation context can be later on reinstalled. The basic Curry–Howard-style correspondence for classical logic is given below. Note the correspondence between the double-negation translation used to map classical proofs to intuitionistic logic and the continuation-passing-style translation used to map lambda terms involving control to pure lambda terms. More particularly, call-by-name continuation-passing-style translations relates to Kolmogorov's double negation translation and call-by-value continuation-passing-style translations relates to a kind of double-negation translation due to Kuroda.

Logic side Programming side
Peirce's law: ((αβ) → α) → α call-with-current-continuation
double-negation translation continuation-passing-style translation

A finer Curry–Howard correspondence exists for classical logic if one defines classical logic not by adding an axiom such as Peirce's law, but by allowing several conclusions in sequents. In the case of classical natural deduction, there exists a proofs-as-programs correspondence with the typed programs of Parigot's λμ-calculus.

### Sequent calculus

A proofs-as-programs correspondence can be settled for the formalism known as Gentzen's sequent calculus but it is not a correspondence with a well-defined pre-existing model of computation as it was for Hilbert-style and natural deductions.

Sequent calculus is characterized by the presence of left introduction rules, right introduction rule and a cut rule that can be eliminated. The structure of sequent calculus relates to a calculus whose structure is close to the one of some abstract machines. The informal correspondence is as follows:

Logic side Programming side
cut elimination reduction in a form of abstract machine
right introduction rules constructors of code
left introduction rules constructors of evaluation stacks
priority to right-hand side in cut-elimination call-by-name reduction
priority to left-hand side in cut-elimination call-by-value reduction

## Related proofs-as-programs correspondences

### The role of de Bruijn

N. G. de Bruijn used the lambda notation for representing proofs of the theorem checker Automath, and represented propositions as "categories" of their proofs. It was in the late 1960s at the same period of time Howard wrote his manuscript; de Bruijn was likely unaware of Howard's work, and stated the correspondence independently (Sørensen & Urzyczyn [1998] 2006, pp 98–99). Some researchers tend to use the term Curry–Howard–de Bruijn correspondence in place of Curry–Howard correspondence.

### BHK interpretation

The BHK interpretation interprets intuitionistic proofs as functions but it does not specify the class of functions relevant for the interpretation. If one takes lambda calculus for this class of function, then the BHK interpretation tells the same as Howard's correspondence between natural deduction and lambda calculus.

### Realizability

Kleene's recursive realizability splits proofs of intuitionistic arithmetic into the pair of a recursive function and of a proof of a formula expressing that the recursive function "realizes", i.e. correctly instantiates the disjunctions and existential quantifiers of the initial formula so that the formula gets true.

Kreisel's modified realizability applies to intuitionistic higher-order predicate logic and shows that the simply typed lambda term inductively extracted from the proof realizes the initial formula. In the case of propositional logic, it coincides with Howard's statement: the extracted lambda term is the proof itself (seen as an untyped lambda term) and the realizability statement is a paraphrase of the fact that the extracted lambda term has the type that the formula means (seen as a type).

Gödel's dialectica interpretation realizes (an extension of) intuitionistic arithmetic with computable functions. The connection with lambda calculus is unclear, even in the case of natural deduction.

### Curry–Howard–Lambek correspondence

Joachim Lambek showed in the early 1970s that the proofs of intuitionistic propositional logic and the combinators of typed combinatory logic share a common equational theory, the theory of cartesian closed categories. The expression Curry–Howard–Lambek correspondence is now used by some people[by whom?] to refer to the relationships between intuitionistic logic, typed lambda calculus and cartesian closed categories, with objects being interpreted as types or propositions and morphisms as terms or proofs. Lambek's correspondence is a correspondence of equational theories, abstracting away from dynamics of computation such as beta reduction and term normalization, and is not the expression of a syntactic identity of structures as it is the case for each of Curry's and Howard's correspondences: i.e. the structure of a well-defined morphism in a cartesian-closed category is not comparable to the structure of a proof of the corresponding judgment in either Hilbert-style logic or natural deduction. For example, it is not possible to state or prove that a morphism is normalizing, establish a Church-Rosser type theorem, or speak of a "strongly normalizing" cartesian closed category. To clarify this distinction, the underlying syntactic structure of cartesian closed categories is rephrased below.

Objects (types) are defined by

• ${\displaystyle \top }$ is an object
• if α and β are objects then ${\displaystyle \alpha \times \beta }$ and ${\displaystyle \alpha \rightarrow \beta }$ are objects.

Morphisms (terms) are defined by

• ${\displaystyle id}$, ${\displaystyle \star }$, ${\displaystyle \operatorname {eval} }$, ${\displaystyle \pi _{1}}$ and ${\displaystyle \pi _{2}}$ are morphisms
• if t is a morphism, λt is a morphism
• if t and u are morphisms, ${\displaystyle (t,u)}$ and ${\displaystyle u\circ t}$ are morphisms.

Well-defined morphisms (typed terms) are defined by the following typing rules (in which the usual categorical morphism notation ${\displaystyle f:\alpha \to \beta }$ is replaced with sequent calculus notation ${\displaystyle f:\!\!-~~\alpha ~\vdash ~\beta }$).

Identity:

${\displaystyle {\frac {}{id:\!\!-~~\alpha ~\vdash ~\alpha }}}$

Composition:

${\displaystyle {\frac {t:\!\!-~~\alpha ~\vdash ~\beta \qquad u:\!\!-~~\beta ~\vdash ~\gamma }{u\circ t:\!\!-~\alpha ~\vdash ~\gamma }}}$
${\displaystyle {\frac {}{\star :\!\!-~~\alpha ~\vdash ~\top }}}$

Cartesian product:

${\displaystyle {\frac {t:\!\!-~~\alpha ~\vdash ~\beta \qquad u:\!\!-~~\alpha ~\vdash ~\gamma }{(t,u):\!\!-~~\alpha ~\vdash ~\beta \times \gamma }}}$

Left and right projection:

${\displaystyle {\frac {}{\pi _{1}:\!\!-~~\alpha \times \beta ~\vdash ~\alpha }}\qquad {\frac {}{\pi _{2}:\!\!-~~\alpha \times \beta ~\vdash ~\beta }}}$
${\displaystyle {\frac {t:\!\!-~~\alpha \times \beta ~\vdash ~\gamma }{\lambda t:\!\!-~~\alpha ~\vdash ~\beta \rightarrow \gamma }}}$
${\displaystyle {\frac {}{eval:\!\!-~~(\alpha \rightarrow \beta )\times \alpha ~\vdash ~\beta }}}$

Finally, the equations of the category are

• ${\displaystyle id\circ t=t}$
• ${\displaystyle t\circ id=t}$
• ${\displaystyle (v\circ u)\circ t=v\circ (u\circ t)}$
• ${\displaystyle \star =id}$ (if well-typed)
• ${\displaystyle \star \circ u=\star }$
• ${\displaystyle \pi _{1}\circ (t,u)=t}$
• ${\displaystyle \pi _{2}\circ (t,u)=u}$
• ${\displaystyle (\pi _{1},\pi _{2})=id}$
• ${\displaystyle (t_{1},t_{2})\circ u=(t_{1}\circ u,t_{2}\circ u)}$
• ${\displaystyle eval\circ (\lambda t\circ \pi _{1},\pi _{2})=t}$
• ${\displaystyle \lambda eval=id}$
• ${\displaystyle \lambda t\circ u=\lambda (t\circ (u\circ \pi _{1},\pi _{2}))}$

These equations imply the following ${\displaystyle \eta }$-laws:

• ${\displaystyle (\pi _{1}\circ t,\pi _{2}\circ t)=t}$
• ${\displaystyle \lambda (eval\circ (t\circ \pi _{1},\pi _{2}))=t}$

Now, there exists t such that ${\displaystyle t:\!\!-~\alpha _{1}\times \ldots \times \alpha _{n}\vdash \beta }$ iff ${\displaystyle \alpha _{1},\ldots ,\alpha _{n}\vdash \beta }$ is provable in implicational intuitionistic logic.

## Examples

Thanks to the Curry–Howard correspondence, a typed expression whose type corresponds to a logical formula is analogous to a proof of that formula. Here are examples.

### The identity combinator seen as a proof of α → α in Hilbert-style logic

As an example, consider a proof of the theorem αα. In lambda calculus, this is the type of the identity function I = λx.x and in combinatory logic, the identity function is obtained by applying S = λfgx.fx(gx) twice to K = λxy.x. That is, I = ((S K) K). As a description of a proof, this says that the following steps can be used to prove αα:

• instantiate the second axiom scheme with the formulas α, βα and α to obtain a proof of (α → ((βα) → α)) → ((α → (βα)) → (αα)),
• instantiate the first axiom scheme once with α and βα to obtain a proof of α → ((βα) → α),
• instantiate the first axiom scheme a second time with α and β to obtain a proof of α → (βα),
• apply modus ponens twice to obtain a proof of αα

In general, the procedure is that whenever the program contains an application of the form (P Q), these steps should be followed:

1. First prove theorems corresponding to the types of P and Q.
2. Since P is being applied to Q, the type of P must have the form αβ and the type of Q must have the form α for some α and β. Therefore, it is possible to detach the conclusion, β, via the modus ponens rule.

### The composition combinator seen as a proof of (β → α) → (γ → β) → γ → α in Hilbert-style logic

As a more complicated example, let's look at the theorem that corresponds to the B function. The type of B is (βα) → (γβ) → γα. B is equivalent to (S (K S) K). This is our roadmap for the proof of the theorem (βα) → (γβ) → γα.

The first step is to construct (K S). To make the antecedent of the K axiom look like the S axiom, set α equal to (αβγ) → (αβ) → αγ, and β equal to δ (to avoid variable collisions):

K : αβα
K[α = (αβγ) → (αβ) → αγ, β = δ] : ((αβγ) → (αβ) → αγ) → δ → (αβγ) → (αβ) → αγ

Since the antecedent here is just S, the consequent can be detached using Modus Ponens:

K S : δ → (αβγ) → (αβ) → αγ

This is the theorem that corresponds to the type of (K S). Now apply S to this expression. Taking S as follows

S : (αβγ) → (αβ) → αγ,

put α = δ, β = αβγ, and γ = (αβ) → αγ, yielding

S[α = δ, β = αβγ, γ = (αβ) → αγ] : (δ → (αβγ) → (αβ) → αγ) → (δ → (αβγ)) → δ → (αβ) → αγ

and then detach the consequent:

S (K S) : (δαβγ) → δ → (αβ) → αγ

This is the formula for the type of (S (K S)). A special case of this theorem has δ = (βγ):

S (K S)[δ = βγ] : ((βγ) → αβγ) → (βγ) → (αβ) → αγ

This last formula must be applied to K. Specialize K again, this time by replacing α with (βγ) and β with α:

K : αβα
K[α = βγ, β = α] : (βγ) → α → (βγ)

This is the same as the antecedent of the prior formula so, detaching the consequent:

S (K S) K : (βγ) → (αβ) → αγ

Switching the names of the variables α and γ gives us

(βα) → (γβ) → γα

which was what remained to prove.

### The normal proof of (β → α) → (γ → β) → γ → α in natural deduction seen as a λ-term

The diagram below gives proof of (βα) → (γβ) → γα in natural deduction and shows how it can be interpreted as the λ-expression λabg.(a (b g)) of type (βα) → (γβ) → γα.

                                     a:β → α, b:γ → β, g:γ ⊢ b : γ → β    a:β → α, b:γ → β, g:γ ⊢ g : γ
———————————————————————————————————  ————————————————————————————————————————————————————————————————————
a:β → α, b:γ → β, g:γ ⊢ a : β → α      a:β → α, b:γ → β, g:γ ⊢ b g : β
————————————————————————————————————————————————————————————————————————
a:β → α, b:γ → β, g:γ ⊢ a (b g) : α
————————————————————————————————————
a:β → α, b:γ → β ⊢ λ g. a (b g) : γ → α
————————————————————————————————————————
a:β → α ⊢ λ b. λ g. a (b g) : (γ → β) → γ → α
————————————————————————————————————
⊢ λ a. λ b. λ g. a (b g) : (β → α) → (γ → β) → γ → α


## Other applications

Recently, the isomorphism has been proposed as a way to define search space partition in genetic programming.[12] The method indexes sets of genotypes (the program trees evolved by the GP system) by their Curry–Howard isomorphic proof (referred to as a species).

As noted by INRIA research director Bernard Lang,[13] the Curry-Howard correspondence constitutes an argument against the patentability of software: since algorithms are mathematical proofs, patentability of the former would imply patentability of the latter. A theorem could be private property; a mathematician would have to pay for using it, and to trust the company that sells it but keeps its proof secret and rejects responsibility for any errors.

## Generalizations

The correspondences listed here go much farther and deeper. For example, cartesian closed categories are generalized by closed monoidal categories. The internal language of these categories is the linear type system (corresponding to linear logic), which generalizes simply-typed lambda calculus as the internal language of cartesian closed categories. Moreover, these can be shown to correspond to cobordisms,[14] which play a vital role in string theory.

An extended set of equivalences is also explored in homotopy type theory, which became a very active area of research around 2013 and as of 2018 still is.[15] Here, type theory is extended by the univalence axiom ("equivalence is equivalent to equality") which permits homotopy type theory to be used as a foundation for all of mathematics (including set theory and classical logic, providing new ways to discuss the axiom of choice and many other things). That is, the Curry–Howard correspondence that proofs are elements of inhabited types is generalized to the notion of homotopic equivalence of proofs (as paths in space, the identity type or equality type of type theory being interpreted as a path).[16]

## References

1. ^ The correspondence was first made explicit in Howard 1980. See, for example section 4.6, p.53 Gert Smolka and Jan Schwinghammer (2007-8), Lecture Notes in Semantics
2. ^ The Brouwer–Heyting–Kolmogorov interpretation is also called the 'proof interpretation': p. 161 of Juliette Kennedy, Roman Kossak, eds. 2011. Set Theory, Arithmetic, and Foundations of Mathematics: Theorems, Philosophies ISBN 978-1-107-00804-5
3. ^ Casadio, Claudia; Scott, Philip J. (2021). Joachim Lambek: The Interplay of Mathematics, Logic, and Linguistics. Springer. p. 184. ISBN 978-3-030-66545-6.
4. ^ Coecke, Bob; Kissinger, Aleks (2017). Picturing Quantum Processes. Cambridge University Press. p. 82. ISBN 978-1-107-10422-8.
5. ^ "Computational trilogy". nLab. Retrieved October 29, 2023.
6. ^
7. ^
8. ^
9. ^ Moggi, Eugenio (1991), "Notions of Computation and Monads" (PDF), Information and Computation, 93 (1): 55–92, doi:10.1016/0890-5401(91)90052-4
10. ^ Sørenson, Morten; Urzyczyn, Paweł (1998), Lectures on the Curry-Howard Isomorphism, CiteSeerX 10.1.1.17.7385
11. ^ Goldblatt, "7.6 Grothendieck Topology as Intuitionistic Modality" (PDF), Mathematical Modal Logic: A Model of its Evolution, pp. 76–81. The "lax" modality referred to is from Benton; Bierman; de Paiva (1998), "Computational types from a logical perspective", Journal of Functional Programming, 8 (2): 177–193, CiteSeerX 10.1.1.258.6004, doi:10.1017/s0956796898002998, S2CID 6149614
12. ^ F. Binard and A. Felty, "Genetic programming with polymorphic types and higher-order functions." In Proceedings of the 10th annual conference on Genetic and evolutionary computation, pages 1187 1194, 2008.[1]
13. ^ "Article". bat8.inria.fr. Archived from the original on 2013-11-17. Retrieved 2020-01-31.
14. ^ John c. Baez and Mike Stay, "Physics, Topology, Logic and Computation: A Rosetta Stone", (2009) ArXiv 0903.0340 in New Structures for Physics, ed. Bob Coecke, Lecture Notes in Physics vol. 813, Springer, Berlin, 2011, pp. 95–174.
16. ^ Homotopy Type Theory: Univalent Foundations of Mathematics. (2013) The Univalent Foundations Program. Institute for Advanced Study.

### Extensions of the correspondence

1. ^ a b Pfenning, Frank; Davies, Rowan (2001), "A Judgmental Reconstruction of Modal Logic" (PDF), Mathematical Structures in Computer Science, 11 (4): 511–540, CiteSeerX 10.1.1.43.1611, doi:10.1017/S0960129501003322, S2CID 16467268
2. ^ Davies, Rowan; Pfenning, Frank (2001), "A Modal Analysis of Staged Computation" (PDF), Journal of the ACM, 48 (3): 555–604, CiteSeerX 10.1.1.3.5442, doi:10.1145/382780.382785, S2CID 52148006
• Griffin, Timothy G. (1990), "The Formulae-as-Types Notion of Control", Conf. Record 17th Annual ACM Symp. on Principles of Programming Languages, POPL '90, San Francisco, CA, USA, 17–19 Jan 1990, pp. 47–57, doi:10.1145/96709.96714, ISBN 978-0-89791-343-0, S2CID 3005134.
• Parigot, Michel (1992), "Lambda-mu-calculus: An algorithmic interpretation of classical natural deduction", International Conference on Logic Programming and Automated Reasoning: LPAR '92 Proceedings, St. Petersburg, Russia, Lecture Notes in Computer Science, vol. 624, Springer-Verlag, pp. 190–201, ISBN 978-3-540-55727-2.
• Herbelin, Hugo (1995), "A Lambda-Calculus Structure Isomorphic to Gentzen-Style Sequent Calculus Structure", in Pacholski, Leszek; Tiuryn, Jerzy (eds.), Computer Science Logic, 8th International Workshop, CSL '94, Kazimierz, Poland, September 25–30, 1994, Selected Papers, Lecture Notes in Computer Science, vol. 933, Springer-Verlag, pp. 61–75, ISBN 978-3-540-60017-6.
• Gabbay, Dov; de Queiroz, Ruy (1992). "Extending the Curry–Howard interpretation to linear, relevant and other resource logics". Journal of Symbolic Logic. Vol. 57. Association for Symbolic Logic. pp. 1319–1365. doi:10.2307/2275370. JSTOR 2275370. S2CID 7159005.. (Full version of the paper presented at Logic Colloquium '90, Helsinki. Abstract in JSL 56(3):1139–1140, 1991.)
• de Queiroz, Ruy; Gabbay, Dov (1994), "Equality in Labelled Deductive Systems and the Functional Interpretation of Propositional Equality", in Dekker, Paul; Stokhof, Martin (eds.), Proceedings of the Ninth Amsterdam Colloquium, ILLC/Department of Philosophy, University of Amsterdam, pp. 547–565, ISBN 978-90-74795-07-4.
• de Queiroz, Ruy; Gabbay, Dov (1995), "The Functional Interpretation of the Existential Quantifier", Bulletin of the Interest Group in Pure and Applied Logics, vol. 3, pp. 243–290. (Full version of a paper presented at Logic Colloquium '91, Uppsala. Abstract in JSL 58(2):753–754, 1993.)
• de Queiroz, Ruy; Gabbay, Dov (1997), "The Functional Interpretation of Modal Necessity", in de Rijke, Maarten (ed.), Advances in Intensional Logic, Applied Logic Series, vol. 7, Springer-Verlag, pp. 61–91, ISBN 978-0-7923-4711-8.
• de Queiroz, Ruy; Gabbay, Dov (1999), "Labelled Natural Deduction", in Ohlbach, Hans-Juergen; Reyle, Uwe (eds.), Logic, Language and Reasoning. Essays in Honor of Dov Gabbay, Trends in Logic, vol. 7, Kluwer, pp. 173–250, ISBN 978-0-7923-5687-5.
• de Oliveira, Anjolina; de Queiroz, Ruy (1999), "A Normalization Procedure for the Equational Fragment of Labelled Natural Deduction", Logic Journal of the Interest Group in Pure and Applied Logics, vol. 7, Oxford University Press, pp. 173–215. (Full version of a paper presented at 2nd WoLLIC'95, Recife. Abstract in Journal of the Interest Group in Pure and Applied Logics 4(2):330–332, 1996.)
• Poernomo, Iman; Crossley, John; Wirsing; Martin (2005), Adapting Proofs-as-Programs: The Curry–Howard Protocol, Monographs in Computer Science, Springer, ISBN 978-0-387-23759-6, concerns the adaptation of proofs-as-programs program synthesis to coarse-grain and imperative program development problems, via a method the authors call the Curry–Howard protocol. Includes a discussion of the Curry–Howard correspondence from a Computer Science perspective.
• de Queiroz, Ruy J.G.B.; de Oliveira, Anjolina (2011), "The Functional Interpretation of Direct Computations", Electronic Notes in Theoretical Computer Science, Elsevier, 269: 19–40, doi:10.1016/j.entcs.2011.03.003. (Full version of a paper presented at LSFA 2010, Natal, Brazil.)

### Books

• De Groote, Philippe, ed. (1995), The Curry–Howard Isomorphism, Cahiers du Centre de Logique (Université catholique de Louvain), vol. 8, Academia-Bruylant, ISBN 978-2-87209-363-2, reproduces the seminal papers of Curry-Feys and Howard, a paper by de Bruijn and a few other papers.
• Sørensen, Morten Heine; Urzyczyn, Paweł (2006) [1998], Lectures on the Curry–Howard isomorphism, Studies in Logic and the Foundations of Mathematics, vol. 149, Elsevier Science, CiteSeerX 10.1.1.17.7385, ISBN 978-0-444-52077-7, notes on proof theory and type theory, that includes a presentation of the Curry–Howard correspondence, with a focus on the formulae-as-types correspondence
• Girard, Jean-Yves (1987–1990), Proof and Types, Cambridge Tracts in Theoretical Computer Science, vol. 7, Translated by and with appendices by Lafont, Yves and Taylor, Paul, Cambridge University Press, ISBN 0-521-37181-3, archived from the original on 2008-04-18, notes on proof theory with a presentation of the Curry–Howard correspondence.
• Thompson, Simon (1991), Type Theory and Functional Programming, Addison–Wesley, ISBN 0-201-41667-0.
• Poernomo, Iman; Crossley, John; Wirsing; Martin (2005), Adapting Proofs-as-Programs: The Curry–Howard Protocol, Monographs in Computer Science, Springer, ISBN 978-0-387-23759-6, concerns the adaptation of proofs-as-programs program synthesis to coarse-grain and imperative program development problems, via a method the authors call the Curry–Howard protocol. Includes a discussion of the Curry–Howard correspondence from a Computer Science perspective.
• Binard, F.; Felty, A. (2008), "Genetic programming with polymorphic types and higher-order functions" (PDF), Proceedings of the 10th annual conference on Genetic and evolutionary computation, Association for Computing Machinery, pp. 1187–94, doi:10.1145/1389095.1389330, ISBN 9781605581309, S2CID 3669630
• de Queiroz, Ruy J.G.B.; de Oliveira, Anjolina G.; Gabbay, Dov M. (2011), The Functional Interpretation of Logical Deduction, Advances in Logic, vol. 5, Imperial College Press/World Scientific, ISBN 978-981-4360-95-1.