Equality in Coq

Coq views truth through the lens of provability. The hypotheses it manipulates are not mere assertions of truth, but formal proofs of the corresponding statements ─ data structures that can be inspected to build other proofs. It is not a coincidence that function types and logical implication use the same notation, A -> B , because proofs of implication in Coq are functions: they take proofs of the precondition as inputs and return a proof of the consequent as the output. Such proofs are written with the same language we use for programming in Coq; tactics are but scripts that build such programs for us. A proof that implication is transitive, for example, amounts to function composition.



Definition implication_is_transitive ( A B C : Prop ) :

( A -> B ) -> ( B -> C ) -> ( A -> C ) :=

fun HAB HBC HA => HBC ( HAB HA ).





Similarly, inductive propositions in Coq behave just like algebraic data types in typical functional programming languages. With pattern matching, we can check which constructor was used in a proof and act accordingly.



Definition or_false_r ( A : Prop ) : A \/ False -> A :=

fun ( H : A \/ False ) =>

match H with

| or_introl HA => HA

| or_intror contra => match contra with end

end .





\/ is an inductive proposition with two constructors, or_introl and or_intror , whose arguments are proofs of its left and right sides. In other words, a proof of A \/ B is either a proof of A or a proof of B . Falsehood, on the other hand, is an inductive proposition with no constructors. Matching on a proof of False does not require us to consider any cases, thus allowing the expression to have any type we please. This corresponds to the so-called The idea of viewing proofs as programs is known as the Defining equality The Coq standard library defines equality as an indexed inductive proposition. (The familiar x = y syntax is provided by the standard library using Coq's notation mechanism.) Disjunctionis an inductive proposition with two constructors,and, whose arguments are proofs of its left and right sides. In other words, a proof ofis either a proof ofor a proof of. Falsehood, on the other hand, is an inductive proposition with no constructors. Matching on a proof ofdoes not require us to consider any cases, thus allowing the expression to have any type we please. This corresponds to the so-called principle of explosion , which asserts that from a contradiction, anything follows.The idea of viewing proofs as programs is known as the Curry-Howard correspondence . It has been a fruitful source of inspiration for the design of many other logics and programming languages beyond Coq, other noteworthy examples including Agda and Nuprl . I will discuss a particular facet of this correspondence in Coq: the meaning of a proof of equality.The Coq standard library defines equality as an indexed inductive proposition. (The familiarsyntax is provided by the standard library using Coq's notation mechanism.)



Inductive eq ( T : Type ) ( x : T ) : T -> Prop :=

| eq_refl : eq T x x .





x = y is when x and y are the "same" term ─ not in the strict sense of syntactic equality, but in the more lenient sense of equality "up to computation" used in Coq's theory. For instance, we can use eq_refl to show that 1 + 1 = 2 , because Coq can simplify the left-hand side using the definition of + and arrive at the right-hand side. To prove interesting facts about equality, we generally use the rewrite tactic, which in turn is implemented by pattern matching. Matching on proofs of equality is more complicated than for typical data types because it is a non-uniform indexed proposition ─ that is, the value of its last argument is not fixed for the whole declaration, but depends on the constructor used. (This non-uniformity is what allows us to put two occurrences of x on the type of eq_refl .) Concretely, suppose that we have elements x and y of a type T , and a predicate P : T -> Prop . We want to prove that P y holds assuming that x = y and P x hold. This can be done with the following program: This declaration says that the most basic way of showingis whenandare the "same" term ─ not in the strict sense of syntactic equality, but in the more lenient sense of equality "up to computation" used in Coq's theory. For instance, we can useto show that, because Coq can simplify the left-hand side using the definition ofand arrive at the right-hand side.To prove interesting facts about equality, we generally use thetactic, which in turn is implemented by pattern matching. Matching on proofs of equality is more complicated than for typical data types because it is aindexed proposition ─ that is, the value of its last argument is not fixed for the whole declaration, but depends on the constructor used. (This non-uniformity is what allows us to put two occurrences ofon the type of.)Concretely, suppose that we have elementsandof a type, and a predicate. We want to prove thatholds assuming thatandhold. This can be done with the following program:



Definition rewriting

( T : Type ) ( P : T -> Prop ) ( x y : T ) ( p : x = y ) ( H : P x ) : P y :=

match p in _ = z return P z with

| eq_refl => H

end .





in _ = z , allows us to provide a name to the non-uniform argument of the type of p . The second, return P z , allows us to declare what the return type of the match expression is as a function of z . At the top level, z corresponds to y , which means that the whole match has type P y . When checking each individual branch, however, Coq requires proofs of P z using values of z that correspond to the constructors of that branch. Inside the eq_refl branch, z corresponds to x , which means that we have to provide a proof of P x . This is why the use of H there is valid. To illustrate, here are proofs of two basic facts about equality: transitivity and symmetry. Compared to common match expressions, this one has two extra clauses. The first,, allows us to provide a name to the non-uniform argument of the type of. The second,, allows us to declare what the return type of the match expression is as a function of. At the top level,corresponds to, which means that the wholehas type. When checking each individual branch, however, Coq requires proofs ofusing values ofthat correspond to the constructors of that branch. Inside thebranch,corresponds to, which means that we have to provide a proof of. This is why the use ofthere is valid.To illustrate, here are proofs of two basic facts about equality: transitivity and symmetry.



Definition etrans { T } { x y z : T } ( p : x = y ) ( q : y = z ) : x = z :=

match p in _ = w return w = z -> x = z with

| eq_refl => fun q' : x = z => q'

end q .



Definition esym { T } { x y : T } ( p : x = y ) : y = x :=

match p in _ = z return z = x with

| eq_refl => eq_refl

end .





w = z alone, as the final type of the expression would be y = z . The other reasonable guess, x = z , wouldn't work either, since we would have nothing of that type to return in the branch ─ q has type y = z , and Coq does not automatically change it to x = z just because we know that x and y ought to be equal inside the branch. The practice of returning a function is so common when matching on dependent types that it even has its own name: the convoy pattern, a term coined by Adam Chlipala in his In addition to functions, pretty much any type expression can go in the return clause of a match . This flexibility allows us to derive many basic reasoning principles ─ for instance, the fact that constructors are disjoint and injective. Notice the return clause in the first proof, which uses a function type. We cannot usealone, as the final type of the expression would be. The other reasonable guess,, wouldn't work either, since we would have nothing of that type to return in the branch ─has type, and Coq does not automatically change it tojust because we know thatandought to be equal inside the branch. The practice of returning a function is so common when matching on dependent types that it even has its own name: the, a term coined by Adam Chlipala in his CDPT book In addition to functions, pretty much any type expression can go in the return clause of a. This flexibility allows us to derive many basic reasoning principles ─ for instance, the fact that constructors are disjoint and injective.



Definition eq_Sn_m ( n m : nat ) ( p : S n = m ) :=

match p in _ = k return match k with

| 0 => False

| S m' => n = m'

end with

| eq_refl => eq_refl

end .



Definition succ_not_zero n : S n <> 0 :=

eq_Sn_m n 0.



Definition succ_inj n m : S n = S m -> n = m :=

eq_Sn_m n ( S m ).





eq_refl branch, we know that k is of the form S n . By substituting this value in the return type, we find that the result of the branch must have type n = n , which is why eq_refl is accepted. Since this is only value of k we have to handle, it doesn't matter that False appears in the return type of the match : that branch will never be used. The more familiar lemmas succ_not_zero and succ_inj simply correspond to special cases of eq_Sn_m . A similar trick can be used for many other inductive types, such as booleans, lists, and so on. Mixing proofs and computation Proofs can be used not only to build other proofs, but also in more conventional programs. If we know that a list is not empty, for example, we can write a function that extracts its first element. In thebranch, we know thatis of the form. By substituting this value in the return type, we find that the result of the branch must have type, which is whyis accepted. Since this is only value ofwe have to handle, it doesn't matter thatappears in the return type of the: that branch will never be used. The more familiar lemmasandsimply correspond to special cases of. A similar trick can be used for many other inductive types, such as booleans, lists, and so on.Proofs can be used not only to build other proofs, but also in more conventional programs. If we know that a list is not empty, for example, we can write a function that extracts its first element.



From mathcomp Require Import seq .



Definition first { T } ( s : seq T ) ( Hs : s <> [::]) : T :=

match s return s <> [::] -> T with

| [::] => fun Hs : [::] <> [::] => match Hs eq_refl with end

| x :: _ => fun _ => x

end Hs .





s , not just on the indices of its type. The rules for checking that this expression is valid are the same: we substitute the pattern of each branch for s in the return type, and ensure that it is compatible with the result it produces. On the first branch, this gives a contradictory hypothesis [::] <> [::] , which we can discard by pattern matching as we did earlier. On the second branch, we can just return the first element of the list. Proofs can also be stored in regular data structures. Consider for instance the subset type { x : T | P x } , which restricts the elements of the type T to those that satisfy the predicate P . Elements of this type are of the form exist x H , where x is an element of T and H is a proof of P x . Here is an alternative version of first , which expects the arguments s and Hs packed as an element of a subset type. Here we see a slightly different use of dependent pattern matching: the return type depends on the analyzed value, not just on the indices of its type. The rules for checking that this expression is valid are the same: we substitute the pattern of each branch forin the return type, and ensure that it is compatible with the result it produces. On the first branch, this gives a contradictory hypothesis, which we can discard by pattern matching as we did earlier. On the second branch, we can just return the first element of the list.Proofs can also be stored in regular data structures. Consider for instance the subset type, which restricts the elements of the typeto those that satisfy the predicate. Elements of this type are of the form, whereis an element ofandis a proof of. Here is an alternative version of, which expects the argumentsandpacked as an element of a subset type.



Definition first' { T } ( s : { s : seq T | s <> [::]}) : T :=

match s with

| exist s Hs => first s Hs

end .





exist x H1 and exist x H2 are equal just by reflexivity; we must explicitly argue that the proofs H1 and H2 are equal. Unfortunately, there are many cases in which this is not possible ─ for example, two proofs of a disjunction A \/ B need to use the same constructor to be considered equal. The situation is not as bad as it might sound, because Coq was designed to allow a proof irrelevance axiom without compromising its soundness. This axiom says that any two proofs of the same proposition are equal. While powerful, this idiom comes with a price: when reasoning about a term that mentions proofs, the proofs must be explicitly taken into account. For instance, we cannot show that two elementsandare equal just by reflexivity; we must explicitly argue that the proofsandare equal. Unfortunately, there are many cases in which this is not possible ─ for example, two proofs of a disjunctionneed to use the same constructor to be considered equal.The situation is not as bad as it might sound, because Coq was designed to allow aaxiom without compromising its soundness. This axiom says that any two proofs of the same proposition are equal.



Axiom proof_irrelevance : forall ( P : Prop ) ( p q : P ), p = q .





Proof irrelevance and equality A classical result of type theory says that equalities between elements of a type T are proof irrelevant provided that T has decidable equality. Many useful properties can be expressed in this way; in particular, any boolean function f : S -> bool can be seen as a predicate S -> Prop defined as fun x : S => f x = true . Thus, if we restrict subset types to computable predicates, we do not need to worry about the proofs that appear in its elements. You might wonder why any assumptions are needed in this result ─ after all, the definition of equality only had a single constructor; how could two proofs be different? Let us begin by trying to show the result without relying on any extra assumptions. We can show that general proof irrelevance can be reduced to irrelevance of "reflexive equality": all proofs of x = x are equal to eq_refl x . If we are willing to extend the theory with this axiom, much of the pain of mixing proofs and computation goes away; nevertheless, it is a bit upsetting that we need an extra axiom to make the use of proofs in computation practical. Fortunately, much of this convenience is already built into Coq's theory, thanks to the structure of proofs of equality.A classical result of type theory says that equalities between elements of a typeare proof irrelevanthas decidable equality. Many useful properties can be expressed in this way; in particular, any boolean functioncan be seen as a predicatedefined as. Thus, if we restrict subset types topredicates, we do not need to worry about the proofs that appear in its elements.You might wonder why any assumptions are needed in this result ─ after all, the definition of equality only had a single constructor; how could two proofs be different? Let us begin by trying to show the result without relying on any extra assumptions. We can show that general proof irrelevance can be reduced to irrelevance of "reflexive equality": all proofs ofare equal to



Section Irrelevance .



Variable T : Type .

Implicit Types x y : T .



Definition almost_irrelevance :

( forall x ( p : x = x ), p = eq_refl x ) ->

( forall x y ( p q : x = y ), p = q ) :=

fun H x y p q =>

match q in _ = z return forall p' : x = z , p' = q with

| eq_refl => fun p' => H x p'

end p .





first : the return type mentions q , the very element we are matching on. It also uses the convoy pattern to "update" the type of p with the extra information gained by matching on q . The almost_irrelevance lemma may look like progress, but it does not actually get us anywhere, because there is no way of proving its premise without assumptions. Here is a failed attempt. This proof uses the extended form of dependent pattern matching we have seen in the definition of: the return type mentions, the very element we are matching on. It also uses the convoy pattern to "update" the type ofwith the extra information gained by matching onThelemma may look like progress, but it does not actually get us anywhere, because there is no way of proving its premise without assumptions. Here is a failed attempt.



Fail Definition irrelevance x ( p : x = x ) : p = eq_refl x :=

match p in _ = y return p = eq_refl x with

| eq_refl => eq_refl

end .





x = x , but its left-hand side has type x = y . That is because when checking the return type, Coq does not use the original type of p , but the one obtained by generalizing the index of its type according to the in clause. It took many years to understand that, even though the inductive definition of equality only mentions one constructor, it is possible to extend the type theory to allow for provably different proofs of equality between two elements. univalence principle that says that proofs of equality between two types correspond to isomorphisms between them. Since there are often many different isomorphisms between two types, irrelevance cannot hold in full generality. To obtain an irrelevance result, we must assume that T has decidable equality. Coq complains that the return clause is ill-typed: its right-hand side has type, but its left-hand side has type. That is because when checking the return type, Coq does not use the original type of, but the one obtained by generalizing the index of its type according to theclause.It took many years to understand that, even though the inductive definition of equality only mentions one constructor, it is possible to extend the type theory to allow for provably different proofs of equality between two elements. Homotopy type theory , for example, introduced athat says that proofs of equality between two types correspond to isomorphisms between them. Since there are often many different isomorphisms between two types,cannot hold in full generality.To obtain an irrelevance result, we must assume thathas decidable equality.



Hypothesis eq_dec : forall x y , x = y \/ x <> y .





Here is the normalization procedure. The argument roughly proceeds as follows. We use decidable equality to define a normalization procedure that takes a proof of equality as input and produces a canonical proof of equality of the same type as output. Crucially, the output of the procedure does not depend on its input. We then show that the normalization procedure has an inverse, allowing us to conclude ─ all proofs must be equal to the canonical one.Here is the normalization procedure.



Definition normalize { x y } ( p : x = y ) : x = y :=

match eq_dec x y with

| or_introl e => e

| or_intror _ => p

end .





If x = y holds, eq_dec x y must return something of the form or_introl e , the other branch being contradictory. This implies that normalize is constant.



Lemma normalize_const { x y } ( p q : x = y ) : normalize p = normalize q .

Proof . by rewrite / normalize ; case : ( eq_dec x y ). Qed .





The inverse of normalize is defined by combining transitivity and symmetry of equality.



Notation "p * q" := ( etrans p q ).



Notation "p ^-1" := ( esym p )

( at level 3, left associativity , format "p ^-1").



Definition denormalize { x y } ( p : x = y ) := p * ( normalize ( eq_refl y ))^-1.





As the above notation suggests, we can show that esym is the inverse of etrans , in the following sense.



Definition etransK x y ( p : x = y ) : p * p ^-1 = eq_refl x :=

match p in _ = y return p * p ^-1 = eq_refl x with

| eq_refl => eq_refl ( eq_refl x )

end .





irrelevance , resulting from generalizing the right-hand side of p . In this return type, p * p ^-1 has type x = x , which matches the one of eq_refl x . Notice why the result of the eq_refl branch is valid: we must produce something of type eq_refl x * ( eq_refl x )^-1 = eq_refl x , but by the definitions of etrans and esym , the left-hand side computes precisely to eq_refl x . Armed with etransK , we can now relate normalize to its inverse, and conclude the proof of irrelevance. This proof avoids the problem that we encountered in our failed proof of, resulting from generalizing the right-hand side of. In this return type,has type, which matches the one of. Notice why the result of thebranch is valid: we must produce something of type, but by the definitions ofand, the left-hand side computes precisely toArmed with, we can now relateto its inverse, and conclude the proof of irrelevance.



Definition normalizeK x y ( p : x = y ) :

normalize p * ( normalize ( eq_refl y ))^-1 = p :=

match p in _ = y return normalize p * ( normalize ( eq_refl y ))^-1 = p with

| eq_refl => etransK x x ( normalize ( eq_refl x ))

end .



Lemma irrelevance x y ( p q : x = y ) : p = q .

Proof .

by rewrite -[ LHS ] normalizeK -[ RHS ] normalizeK ( normalize_const p q ).

Qed .



End Irrelevance .





Irrelevance of equality in practice The eqtype eq_irrelevance ). The class structure used by eqtype makes it easy for Coq to infer proofs of decidable equality, which considerably simplifies the use of this and other lemmas. The eq_proofs_unicity_on ), though it is a bit harder to use, since it does not make use of any mechanisms for inferring results of decidable equality. The Mathematical Components library provides excellent support for types with decidable equality in its module , including a generic result of proof irrelevance like the one I gave above (). The class structure used bymakes it easy for Coq to infer proofs of decidable equality, which considerably simplifies the use of this and other lemmas. The Coq standard library also provides a proof of this lemma (), though it is a bit harder to use, since it does not make use of any mechanisms for inferring results of decidable equality.

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