module plfa.part3.Soundness where

Introduction

In this chapter we prove that the reduction semantics is sound with respect to the denotational semantics, i.e., for any term L

L —↠ ƛ N  implies  ℰ L ≃ ℰ (ƛ N)

The proof is by induction on the reduction sequence, so the main lemma concerns a single reduction step. We prove that if any term M steps to a term N, then M and N are denotationally equal. We shall prove each direction of this if-and-only-if separately. One direction will look just like a type preservation proof. The other direction is like proving type preservation for reduction going in reverse. Recall that type preservation is sometimes called subject reduction. Preservation in reverse is a well-known property and is called subject expansion. It is also well-known that subject expansion is false for most typed lambda calculi!

Imports

open import Relation.Binary.PropositionalEquality
  using (_≡_; _≢_; refl; sym; cong; cong₂; cong-app)
open import Data.Product using (_×_; Σ; Σ-syntax; ; ∃-syntax; proj₁; proj₂)
  renaming (_,_ to ⟨_,_⟩)
open import Agda.Primitive using (lzero)
open import Relation.Nullary using (¬_)
open import Relation.Nullary.Negation using (contradiction)
open import Data.Empty using (⊥-elim)
open import Relation.Nullary using (Dec; yes; no)
open import Function using (_∘_)
open import plfa.part2.Untyped
     using (Context; _,_; _∋_; _⊢_; ; Z; S_; `_; ƛ_; _·_;
            subst; _[_]; subst-zero; ext; rename; exts;
            _—→_; ξ₁; ξ₂; β; ζ; _—↠_; _—→⟨_⟩_; _∎)
open import plfa.part2.Substitution using (Rename; Subst; ids)
open import plfa.part3.Denotational
     using (Value; ; Env; _⊢_↓_; _`,_; _⊑_; _`⊑_; `⊥; _`⊔_; init; last; init-last;
            ⊑-refl; ⊑-trans; `⊑-refl; ⊑-env; ⊑-env-conj-R1; ⊑-env-conj-R2; up-env;
            var; ↦-elim; ↦-intro; ⊥-intro; ⊔-intro; sub;
            rename-pres; ; _≃_; ≃-trans)
open import plfa.part3.Compositional using (lambda-inversion; var-inv)

Forward reduction preserves denotations

The proof of preservation in this section mixes techniques from previous chapters. Like the proof of preservation for the STLC, we are preserving a relation defined separately from the syntax, in contrast to the intrinsically-typed terms. On the other hand, we are using de Bruijn indices for variables.

The outline of the proof remains the same in that we must prove lemmas concerning all of the auxiliary functions used in the reduction relation: substitution, renaming, and extension.

Simultaneous substitution preserves denotations

Our next goal is to prove that simultaneous substitution preserves meaning. That is, if M results in v in environment γ, then applying a substitution σ to M gives us a program that also results in v, but in an environment δ in which, for every variable x, σ x results in the same value as the one for x in the original environment γ. We write δ ⊢ σ ↓ γ for this condition.

infix 3 _`⊢_↓_
_`⊢_↓_ : ∀{Δ Γ}  Env Δ  Subst Γ Δ  Env Γ  Set
_`⊢_↓_ {Δ}{Γ} δ σ γ = (∀ (x : Γ  )  δ  σ x  γ x)

As usual, to prepare for lambda abstraction, we prove an extension lemma. It says that applying the exts function to a substitution produces a new substitution that maps variables to terms that when evaluated in δ , v produce the values in γ , v.

subst-ext :  {Γ Δ v} {γ : Env Γ} {δ : Env Δ}
   (σ : Subst Γ Δ)
   δ `⊢ σ  γ
   --------------------------
   δ `, v `⊢ exts σ  γ `, v
subst-ext σ d Z = var
subst-ext σ d (S x′) = rename-pres S_  _  ⊑-refl) (d x′)

The proof is by cases on the de Bruijn index x.

  • If it is Z, then we need to show that δ , v ⊢ # 0 ↓ v, which we have by rule var.

  • If it is S x′,then we need to show that δ , v ⊢ rename S_ (σ x′) ↓ γ x′, which we obtain by the rename-pres lemma.

With the extension lemma in hand, the proof that simultaneous substitution preserves meaning is straightforward. Let’s dive in!

subst-pres :  {Γ Δ v} {γ : Env Γ} {δ : Env Δ} {M : Γ  }
   (σ : Subst Γ Δ)
   δ `⊢ σ  γ
   γ  M  v
    ------------------
   δ  subst σ M  v
subst-pres σ s (var {x = x}) = (s x)
subst-pres σ s (↦-elim d₁ d₂) =
  ↦-elim (subst-pres σ s d₁) (subst-pres σ s d₂)
subst-pres σ s (↦-intro d) =
  ↦-intro (subst-pres  {A}  exts σ) (subst-ext σ s) d)
subst-pres σ s ⊥-intro = ⊥-intro
subst-pres σ s (⊔-intro d₁ d₂) =
  ⊔-intro (subst-pres σ s d₁) (subst-pres σ s d₂)
subst-pres σ s (sub d lt) = sub (subst-pres σ s d) lt

The proof is by induction on the semantics of M. The two interesting cases are for variables and lambda abstractions.

  • For a variable x, we have that v ⊑ γ x and we need to show that δ ⊢ σ x ↓ v. From the premise applied to x, we have that δ ⊢ σ x ↓ γ x, so we conclude by the sub rule.

  • For a lambda abstraction, we must extend the substitution for the induction hypothesis. We apply the subst-ext lemma to show that the extended substitution maps variables to terms that result in the appropriate values.

Single substitution preserves denotations

For β reduction, (ƛ N) · M —→ N [ M ], we need to show that the semantics is preserved when substituting M for de Bruijn index 0 in term N. By inversion on the rules ↦-elim and ↦-intro, we have that γ , v ⊢ M ↓ w and γ ⊢ N ↓ v. So we need to show that γ ⊢ M [ N ] ↓ w, or equivalently, that γ ⊢ subst (subst-zero N) M ↓ w.

substitution :  {Γ} {γ : Env Γ} {N M v w}
    γ `, v  N  w
    γ  M  v
     ---------------
    γ  N [ M ]  w
substitution{Γ}{γ}{N}{M}{v}{w} dn dm =
  subst-pres (subst-zero M) sub-z-ok dn
  where
  sub-z-ok : γ `⊢ subst-zero M  (γ `, v)
  sub-z-ok Z = dm
  sub-z-ok (S x) = var

This result is a corollary of the lemma for simultaneous substitution. To use the lemma, we just need to show that subst-zero M maps variables to terms that produces the same values as those in γ , v. Let y be an arbitrary variable (de Bruijn index).

  • If it is Z, then (subst-zero M) y = M and (γ , v) y = v. By the premise we conclude that γ ⊢ M ↓ v.

  • If it is S x, then (subst-zero M) (S x) = x and (γ , v) (S x) = γ x. So we conclude that γ ⊢ x ↓ γ x by rule var.

Reduction preserves denotations

With the substitution lemma in hand, it is straightforward to prove that reduction preserves denotations.

preserve :  {Γ} {γ : Env Γ} {M N v}
   γ  M  v
   M —→ N
    ----------
   γ  N  v
preserve (var) ()
preserve (↦-elim d₁ d₂) (ξ₁ r) = ↦-elim (preserve d₁ r) d₂
preserve (↦-elim d₁ d₂) (ξ₂ r) = ↦-elim d₁ (preserve d₂ r)
preserve (↦-elim d₁ d₂) β = substitution (lambda-inversion d₁) d₂
preserve (↦-intro d) (ζ r) = ↦-intro (preserve d r)
preserve ⊥-intro r = ⊥-intro
preserve (⊔-intro d d₁) r = ⊔-intro (preserve d r) (preserve d₁ r)
preserve (sub d lt) r = sub (preserve d r) lt

We proceed by induction on the semantics of M with case analysis on the reduction.

  • If M is a variable, then there is no such reduction.

  • If M is an application, then the reduction is either a congruence (ξ₁ or ξ₂) or β. For each congruence, we use the induction hypothesis. For β reduction we use the substitution lemma and the sub rule.

  • The rest of the cases are straightforward.

Reduction reflects denotations

This section proves that reduction reflects the denotation of a term. That is, if N results in v, and if M reduces to N, then M also results in v. While there are some broad similarities between this proof and the above proof of semantic preservation, we shall require a few more technical lemmas to obtain this result.

The main challenge is dealing with the substitution in β reduction:

(ƛ N) · M —→ N [ M ]

We have that γ ⊢ N [ M ] ↓ v and need to show that γ ⊢ (ƛ N) · M ↓ v. Now consider the derivation of γ ⊢ N [ M ] ↓ v. The term M may occur 0, 1, or many times inside N [ M ]. At each of those occurrences, M may result in a different value. But to build a derivation for (ƛ N) · M, we need a single value for M. If M occurred more than 1 time, then we can join all of the different values using . If M occurred 0 times, then we do not need any information about M and can therefore use for the value of M.

Renaming reflects meaning

Previously we showed that renaming variables preserves meaning. Now we prove the opposite, that it reflects meaning. That is, if δ ⊢ rename ρ M ↓ v, then γ ⊢ M ↓ v, where (δ ∘ ρ)⊑ γ`.

First, we need a variant of a lemma given earlier.
ext-⊑′ :  {Γ Δ v} {γ : Env Γ} {δ : Env Δ}
   (ρ : Rename Γ Δ)
   (δ  ρ) `⊑ γ
    ------------------------------
   ((δ `, v)  ext ρ) `⊑ (γ `, v)
ext-⊑′ ρ lt Z = ⊑-refl
ext-⊑′ ρ lt (S x) = lt x
The proof is then as follows.
rename-reflect :  {Γ Δ v} {γ : Env Γ} {δ : Env Δ} { M : Γ  }
   {ρ : Rename Γ Δ}
   (δ  ρ) `⊑ γ
   δ  rename ρ M  v
    ------------------------------------
   γ  M  v
rename-reflect {M = ` x} all-n d with var-inv d
... | lt =  sub var (⊑-trans lt (all-n x))
rename-reflect {M = ƛ N}{ρ = ρ} all-n (↦-intro d) =
   ↦-intro (rename-reflect (ext-⊑′ ρ all-n) d)
rename-reflect {M = ƛ N} all-n ⊥-intro = ⊥-intro
rename-reflect {M = ƛ N} all-n (⊔-intro d₁ d₂) =
   ⊔-intro (rename-reflect all-n d₁) (rename-reflect all-n d₂)
rename-reflect {M = ƛ N} all-n (sub d₁ lt) =
   sub (rename-reflect all-n d₁) lt
rename-reflect {M = L · M} all-n (↦-elim d₁ d₂) =
   ↦-elim (rename-reflect all-n d₁) (rename-reflect all-n d₂)
rename-reflect {M = L · M} all-n ⊥-intro = ⊥-intro
rename-reflect {M = L · M} all-n (⊔-intro d₁ d₂) =
   ⊔-intro (rename-reflect all-n d₁) (rename-reflect all-n d₂)
rename-reflect {M = L · M} all-n (sub d₁ lt) =
   sub (rename-reflect all-n d₁) lt

We cannot prove this lemma by induction on the derivation of δ ⊢ rename ρ M ↓ v, so instead we proceed by induction on M.

  • If it is a variable, we apply the inversion lemma to obtain that v ⊑ δ (ρ x). Instantiating the premise to x we have δ (ρ x) = γ x, so we conclude by the var rule.

  • If it is a lambda abstraction ƛ N, we have rename ρ (ƛ N) = ƛ (rename (ext ρ) N). We proceed by cases on δ ⊢ ƛ (rename (ext ρ) N) ↓ v.

    • Rule ↦-intro: To satisfy the premise of the induction hypothesis, we prove that the renaming can be extended to be a mapping from γ , v₁ to δ , v₁.

    • Rule ⊥-intro: We simply apply ⊥-intro.

    • Rule ⊔-intro: We apply the induction hypotheses and ⊔-intro.

    • Rule sub: We apply the induction hypothesis and sub.

  • If it is an application L · M, we have rename ρ (L · M) = (rename ρ L) · (rename ρ M). We proceed by cases on δ ⊢ (rename ρ L) · (rename ρ M) ↓ v and all the cases are straightforward.

In the upcoming uses of rename-reflect, the renaming will always be the increment function. So we prove a corollary for that special case.

rename-inc-reflect :  {Γ v′ v} {γ : Env Γ} { M : Γ  }
   (γ `, v′)  rename S_ M  v
    ----------------------------
   γ  M  v
rename-inc-reflect d = rename-reflect `⊑-refl d

Substitution reflects denotations, the variable case

We are almost ready to begin proving that simultaneous substitution reflects denotations. That is, if γ ⊢ (subst σ M) ↓ v, then γ ⊢ σ k ↓ δ k and δ ⊢ M ↓ v for any k and some δ. We shall start with the case in which M is a variable x. So instead the premise is γ ⊢ σ x ↓ v and we need to show that δ ⊢ x ↓ v for some δ. The δ that we choose shall be the environment that maps x to v and every other variable to .

Next we define the environment that maps x to v and every other variable to , that is const-env x v. To tell variables apart, we define the following function for deciding equality of variables.

_var≟_ :  {Γ}  (x y : Γ  )  Dec (x  y)
Z var≟ Z  =  yes refl
Z var≟ (S _)  =  no λ()
(S _) var≟ Z  =  no λ()
(S x) var≟ (S y) with  x var≟ y
...                 |  yes refl =  yes refl
...                 |  no neq   =  no λ{refl  neq refl}

var≟-refl :  {Γ} (x : Γ  )  (x var≟ x)  yes refl
var≟-refl Z = refl
var≟-refl (S x) rewrite var≟-refl x = refl

Now we use var≟ to define const-env.

const-env : ∀{Γ}  (x : Γ  )  Value  Env Γ
const-env x v y with x var≟ y
...             | yes _       = v
...             | no _        = 

Of course, const-env x v maps x to value v

same-const-env : ∀{Γ} {x : Γ  } {v}  (const-env x v) x  v
same-const-env {x = x} rewrite var≟-refl x = refl

and const-env x v maps y to ⊥, so long asx ≢ y`.

diff-const-env : ∀{Γ} {x y : Γ  } {v}
   x  y
    -------------------
   const-env x v y  
diff-const-env {Γ} {x} {y} neq with x var≟ y
...  | yes eq  =  ⊥-elim (neq eq)
...  | no _    =  refl

So we choose const-env x v for δ and obtain δ ⊢ x ↓ v with the var rule.

It remains to prove that γ ⊢ σ ↓ δ and δ ⊢ M ↓ v for any k, given that we have chosen const-env x v for δ. We shall have two cases to consider, x ≡ y or x ≢ y.

Now to finish the two cases of the proof.

  • In the case where x ≡ y, we need to show that γ ⊢ σ y ↓ v, but that’s just our premise.
  • In the case where x ≢ y, we need to show that γ ⊢ σ y ↓ ⊥, which we do via rule ⊥-intro.

Thus, we have completed the variable case of the proof that simultaneous substitution reflects denotations. Here is the proof again, formally.

subst-reflect-var :  {Γ Δ} {γ : Env Δ} {x : Γ  } {v} {σ : Subst Γ Δ}
   γ  σ x  v
    -----------------------------------------
   Σ[ δ  Env Γ ] γ `⊢ σ  δ  ×  δ  ` x  v
subst-reflect-var {Γ}{Δ}{γ}{x}{v}{σ} xv
  rewrite sym (same-const-env {Γ}{x}{v}) =
     const-env x v ,  const-env-ok , var  
  where
  const-env-ok : γ `⊢ σ  const-env x v
  const-env-ok y with x var≟ y
  ... | yes x≡y rewrite sym x≡y | same-const-env {Γ}{x}{v} = xv
  ... | no x≢y rewrite diff-const-env {Γ}{x}{y}{v} x≢y = ⊥-intro

Substitutions and environment construction

Every substitution produces terms that can evaluate to .

subst-⊥ : ∀{Γ Δ}{γ : Env Δ}{σ : Subst Γ Δ}
    -----------------
   γ `⊢ σ  `⊥
subst-⊥ x = ⊥-intro

If a substitution produces terms that evaluate to the values in both γ₁ and γ₂, then those terms also evaluate to the values in γ₁ ⊔ γ₂.

subst-⊔ : ∀{Γ Δ}{γ : Env Δ}{γ₁ γ₂ : Env Γ}{σ : Subst Γ Δ}
            γ `⊢ σ  γ₁
            γ `⊢ σ  γ₂
             -------------------------
            γ `⊢ σ  (γ₁ `⊔ γ₂)
subst-⊔ γ₁-ok γ₂-ok x = ⊔-intro (γ₁-ok x) (γ₂-ok x)

The Lambda constructor is injective

lambda-inj :  {Γ} {M N : Γ ,    }
   _≡_ {A = Γ  } (ƛ M) (ƛ N)
    ---------------------------
   M  N
lambda-inj refl = refl

Simultaneous substitution reflects denotations

In this section we prove a central lemma, that substitution reflects denotations. That is, if γ ⊢ subst σ M ↓ v, then δ ⊢ M ↓ v and γ ⊢ σ ↓ δ for some δ. We shall proceed by induction on the derivation of γ ⊢ subst σ M ↓ v. This requires a minor restatement of the lemma, changing the premise to γ ⊢ L ↓ v and L ≡ subst σ M.

split :  {Γ} {M : Γ ,   } {δ : Env (Γ , )} {v}
   δ  M  v
    --------------------------
   (init δ `, last δ)  M  v
split {δ = δ} δMv rewrite init-last δ = δMv

subst-reflect :  {Γ Δ} {δ : Env Δ} {M : Γ  } {v} {L : Δ  } {σ : Subst Γ Δ}
   δ  L  v
   subst σ M  L
    ---------------------------------------
   Σ[ γ  Env Γ ] δ `⊢ σ  γ  ×  γ  M  v

subst-reflect {M = M}{σ = σ} (var {x = y}) eqL with M
... | ` x  with var {x = y}
...           | yv  rewrite sym eqL = subst-reflect-var {σ = σ} yv
subst-reflect {M = M} (var {x = y}) () | M₁ · M₂
subst-reflect {M = M} (var {x = y}) () | ƛ M′

subst-reflect {M = M}{σ = σ} (↦-elim d₁ d₂) eqL
         with M
...    | ` x with ↦-elim d₁ d₂
...    | d′ rewrite sym eqL = subst-reflect-var {σ = σ} d′
subst-reflect (↦-elim d₁ d₂) () | ƛ M′
subst-reflect{Γ}{Δ}{γ}{σ = σ} (↦-elim d₁ d₂)
   refl | M₁ · M₂
     with subst-reflect {M = M₁} d₁ refl | subst-reflect {M = M₂} d₂ refl
...     |  δ₁ ,  subst-δ₁ , m1   |  δ₂ ,  subst-δ₂ , m2   =
      δ₁ `⊔ δ₂ ,  subst-⊔ {γ₁ = δ₁}{γ₂ = δ₂}{σ = σ} subst-δ₁ subst-δ₂ ,
                    ↦-elim (⊑-env m1 (⊑-env-conj-R1 δ₁ δ₂))
                           (⊑-env m2 (⊑-env-conj-R2 δ₁ δ₂))  

subst-reflect {M = M}{σ = σ} (↦-intro d) eqL with M
...    | ` x with (↦-intro d)
...             | d′ rewrite sym eqL = subst-reflect-var {σ = σ} d′
subst-reflect {σ = σ} (↦-intro d) eq | ƛ M′
      with subst-reflect {σ = exts σ} d (lambda-inj eq)
... |  δ′ ,  exts-σ-δ′ , m′   =
     init δ′ ,  ((λ x  rename-inc-reflect (exts-σ-δ′ (S x)))) ,
             ↦-intro (up-env (split m′) (var-inv (exts-σ-δ′ Z)))  
subst-reflect (↦-intro d) () | M₁ · M₂

subst-reflect {σ = σ} ⊥-intro eq =
     `⊥ ,  subst-⊥ {σ = σ} , ⊥-intro  

subst-reflect {σ = σ} (⊔-intro d₁ d₂) eq
  with subst-reflect {σ = σ} d₁ eq | subst-reflect {σ = σ} d₂ eq
... |  δ₁ ,  subst-δ₁ , m1   |  δ₂ ,  subst-δ₂ , m2   =
      δ₁ `⊔ δ₂ ,  subst-⊔ {γ₁ = δ₁}{γ₂ = δ₂}{σ = σ} subst-δ₁ subst-δ₂ ,
                    ⊔-intro (⊑-env m1 (⊑-env-conj-R1 δ₁ δ₂))
                            (⊑-env m2 (⊑-env-conj-R2 δ₁ δ₂))  
subst-reflect (sub d lt) eq
    with subst-reflect d eq
... |  δ ,  subst-δ , m   =  δ ,  subst-δ , sub m lt  
  • Case var: We have subst σ M ≡ y, so M must also be a variable, say x. We apply the lemma subst-reflect-var to conclude.

  • Case ↦-elim: We have subst σ M ≡ L₁ · L₂. We proceed by cases on M.

    • Case M ≡ x: We apply the subst-reflect-var lemma again to conclude.

    • Case M ≡ M₁ · M₂: By the induction hypothesis, we have some δ₁ and δ₂ such that δ₁ ⊢ M₁ ↓ v₁ ↦ v₃ and γ ⊢ σ ↓ δ₁, as well as δ₂ ⊢ M₂ ↓ v₁ and γ ⊢ σ ↓ δ₂. By ⊑-env we have δ₁ ⊔ δ₂ ⊢ M₁ ↓ v₁ ↦ v₃ and δ₁ ⊔ δ₂ ⊢ M₂ ↓ v₁ (using ⊑-env-conj-R1 and ⊑-env-conj-R2), and therefore δ₁ ⊔ δ₂ ⊢ M₁ · M₂ ↓ v₃. We conclude this case by obtaining γ ⊢ σ ↓ δ₁ ⊔ δ₂ by the subst-⊔ lemma.

  • Case ↦-intro: We have subst σ M ≡ ƛ L′. We proceed by cases on M.

    • Case M ≡ x: We apply the subst-reflect-var lemma.

    • Case M ≡ ƛ M′: By the induction hypothesis, we have (δ′ , v′) ⊢ M′ ↓ v₂ and (δ , v₁) ⊢ exts σ ↓ (δ′ , v′). From the later we have (δ , v₁) ⊢ # 0 ↓ v′. By the lemma var-inv we have v′ ⊑ v₁, so by the up-env lemma we have (δ′ , v₁) ⊢ M′ ↓ v₂ and therefore δ′ ⊢ ƛ M′ ↓ v₁ → v₂. We also need to show that δ⊢ σ ↓ δ′. Fixk. We have(δ , v₁) ⊢ rename S_ σ k ↓ δ k′. We then apply the lemmarename-inc-reflectto obtainδ ⊢ σ k ↓ δ k′`, so this case is complete.

  • Case ⊥-intro: We choose for δ. We have ⊥ ⊢ M ↓ ⊥ by ⊥-intro. We have δ ⊢ σ ↓ ⊥ by the lemma subst-empty.

  • Case ⊔-intro: By the induction hypothesis we have δ₁ ⊢ M ↓ v₁, δ₂ ⊢ M ↓ v₂, δ ⊢ σ ↓ δ₁, and δ ⊢ σ ↓ δ₂. We have δ₁ ⊔ δ₂ ⊢ M ↓ v₁ and δ₁ ⊔ δ₂ ⊢ M ↓ v₂ by ⊑-env with ⊑-env-conj-R1 and ⊑-env-conj-R2. So by ⊔-intro we have δ₁ ⊔ δ₂ ⊢ M ↓ v₁ ⊔ v₂. By subst-⊔ we conclude that δ ⊢ σ ↓ δ₁ ⊔ δ₂.

Single substitution reflects denotations

Most of the work is now behind us. We have proved that simultaneous substitution reflects denotations. Of course, β reduction uses single substitution, so we need a corollary that proves that single substitution reflects denotations. That is, given terms N : (Γ , ★ ⊢ ★) and M : (Γ ⊢ ★), if γ ⊢ N [ M ] ↓ w, then γ ⊢ M ↓ v and (γ , v) ⊢ N ↓ w for some value v. We have N [ M ] = subst (subst-zero M) N.

We first prove a lemma about subst-zero, that if δ ⊢ subst-zero M ↓ γ, then γ ⊑ (δ , w) × δ ⊢ M ↓ w for some w.

subst-zero-reflect :  {Δ} {δ : Env Δ} {γ : Env (Δ , )} {M : Δ  }
   δ `⊢ subst-zero M  γ
    ----------------------------------------
   Σ[ w  Value ] γ `⊑ (δ `, w) × δ  M  w
subst-zero-reflect {δ = δ} {γ = γ} δσγ =  last γ ,  lemma , δσγ Z  
  where
  lemma : γ `⊑ (δ `, last γ)
  lemma Z  =  ⊑-refl
  lemma (S x) = var-inv (δσγ (S x))

We choose w to be the last value in γ and we obtain δ ⊢ M ↓ w by applying the premise to variable Z. Finally, to prove γ ⊑ (δ , w), we prove a lemma by induction in the input variable. The base case is trivial because of our choice of w. In the induction case, S x, the premise δ ⊢ subst-zero M ↓ γ gives us δ ⊢ x ↓ γ (S x) and then using var-inv we conclude that γ (S x) ⊑ (δ, w) (S x)`.

Now to prove that substitution reflects denotations.

substitution-reflect :  {Δ} {δ : Env Δ} {N : Δ ,   } {M : Δ  } {v}
   δ  N [ M ]  v
    ------------------------------------------------
   Σ[ w  Value ] δ  M  w  ×  (δ `, w)  N  v
substitution-reflect d with subst-reflect d refl
...  |  γ ,  δσγ , γNv   with subst-zero-reflect δσγ
...    |  w ,  ineq , δMw   =  w ,  δMw , ⊑-env γNv ineq  

We apply the subst-reflect lemma to obtain δ ⊢ subst-zero M ↓ γ and γ ⊢ N ↓ v for some γ. Using the former, the subst-zero-reflect lemma gives us γ ⊑ (δ , w) and δ ⊢ M ↓ w. We conclude that δ , w ⊢ N ↓ v by applying the ⊑-env lemma, using γ ⊢ N ↓ v and γ ⊑ (δ , w).

Reduction reflects denotations

Now that we have proved that substitution reflects denotations, we can easily prove that reduction does too.

reflect-beta : ∀{Γ}{γ : Env Γ}{M N}{v}
     γ  (N [ M ])  v
     γ  (ƛ N) · M  v
reflect-beta d
    with substitution-reflect d
... |  v₂′ ,  d₁′ , d₂′   = ↦-elim (↦-intro d₂′) d₁′


reflect :  {Γ} {γ : Env Γ} {M M′ N v}
   γ  N  v    M —→ M′     M′  N
    ---------------------------------
   γ  M  v
reflect var (ξ₁ r) ()
reflect var (ξ₂ r) ()
reflect{γ = γ} (var{x = x}) β mn
    with var{γ = γ}{x = x}
... | d′ rewrite sym mn = reflect-beta d′
reflect var (ζ r) ()
reflect (↦-elim d₁ d₂) (ξ₁ r) refl = ↦-elim (reflect d₁ r refl) d₂
reflect (↦-elim d₁ d₂) (ξ₂ r) refl = ↦-elim d₁ (reflect d₂ r refl)
reflect (↦-elim d₁ d₂) β mn
    with ↦-elim d₁ d₂
... | d′ rewrite sym mn = reflect-beta d′
reflect (↦-elim d₁ d₂) (ζ r) ()
reflect (↦-intro d) (ξ₁ r) ()
reflect (↦-intro d) (ξ₂ r) ()
reflect (↦-intro d) β mn
    with ↦-intro d
... | d′ rewrite sym mn = reflect-beta d′
reflect (↦-intro d) (ζ r) refl = ↦-intro (reflect d r refl)
reflect ⊥-intro r mn = ⊥-intro
reflect (⊔-intro d₁ d₂) r mn rewrite sym mn =
   ⊔-intro (reflect d₁ r refl) (reflect d₂ r refl)
reflect (sub d lt) r mn = sub (reflect d r mn) lt

Reduction implies denotational equality

We have proved that reduction both preserves and reflects denotations. Thus, reduction implies denotational equality.

reduce-equal :  {Γ} {M : Γ  } {N : Γ  }
   M —→ N
    ---------
    M   N
reduce-equal {Γ}{M}{N} r γ v =
      m  preserve m r) ,  n  reflect n r refl) 

We conclude with the soundness property, that multi-step reduction to a lambda abstraction implies denotational equivalence with a lambda abstraction.

soundness : ∀{Γ} {M : Γ  } {N : Γ ,   }
   M —↠ ƛ N
    -----------------
    M   (ƛ N)
soundness (.(ƛ _) ) γ v =   x  x) ,  x  x) 
soundness {Γ} (L —→⟨ r  M—↠N) γ v =
   let ih = soundness M—↠N in
   let e = reduce-equal r in
   ≃-trans {Γ} e ih γ v

Unicode

This chapter uses the following unicode:

≟  U+225F  QUESTIONED EQUAL TO (\?=)