DeBruijn: Inherently typed de Bruijn representation
module plfa.DeBruijn where
The previous two chapters introduced lambda calculus, with a formalisation based on named variables, and terms defined separately from types. We began with that approach because it is traditional, but it is not the one we recommend. This chapter presents an alternative approach, where named variables are replaced by de Bruijn indices and terms are inherently typed. Our new presentation is more compact, using less than half the lines of code required previously to cover the same ground.
The inherently typed representation was first proposed by Thorsten Altenkirch and Bernhard Reus. The formalisation of renaming and substitution we use is due to Conor McBride. Related work has been carried out by James Chapman, James McKinna, and many others.
Imports
import Relation.Binary.PropositionalEquality as Eq open Eq using (_≡_; refl) open import Data.Empty using (⊥; ⊥elim) open import Data.Nat using (ℕ; zero; suc) open import Relation.Nullary using (¬_)
Introduction
There is a close correspondence between the structure of a term and the structure of the derivation showing that it is well typed. For example, here is the term for the Church numeral two:
twoᶜ : Term
twoᶜ = ƛ "s" ⇒ ƛ "z" ⇒ ` "s" · (` "s" · ` "z")
And here is its corresponding type derivation:
⊢twoᶜ : ∀ {A} → ∅ ⊢ twoᶜ ⦂ Ch A
⊢twoᶜ = ⊢ƛ (⊢ƛ (⊢` ∋s · (⊢` ∋s · ⊢` ∋z)))
where
∋s = S ("s" ≠ "z") Z
∋z = Z
(These are both taken from Chapter Lambda and you can see the corresponding derivation tree written out in full here.) The two definitions are in close correspondence, where:
`_
corresponds to⊢`
ƛ_⇒_
corresponds to⊢ƛ
_·_
corresponds to_·_
Further, if we think of Z
as zero and S
as successor, then
the lookup derivation for each variable corresponds to a
number which tells us how many enclosing binding terms to
count to find the binding of that variable. Here "z"
corresponds to Z
or zero and "s"
corresponds to S Z
or
one. And, indeed, "z"
is bound by the inner abstraction
(count outward past zero abstractions) and "s"
is bound by the
outer abstraction (count outward past one abstraction).
In this chapter, we are going to exploit this correspondence, and introduce a new notation for terms that simultaneously represents the term and its type derivation. Now we will write the following:
twoᶜ : ∅ ⊢ Ch `ℕ
twoᶜ = ƛ ƛ (# 1 · (# 1 · # 0))
A variable is represented by a natural number (written with
Z
and S
, and abbreviated in the usual way), and tells us
how many enclosing binding terms to count to find the binding
of that variable. Thus, # 0
is bound at the inner ƛ
, and
# 1
at the outer ƛ
.
Replacing variables by numbers in this way is called de Bruijn representation, and the numbers themselves are called de Bruijn indices, after the Dutch mathematician Nicolaas Govert (Dick) de Bruijn (1918—2012), a pioneer in the creation of proof assistants. One advantage of replacing named variables with de Bruijn indices is that each term now has a unique representation, rather than being represented by the equivalence class of terms under alpha renaming.
The other important feature of our chosen representation is
that it is inherently typed. In the previous two chapters,
the definition of terms and the definition of types are
completely separate. All terms have type Term
, and nothing
in Agda prevents one from writing a nonsense term such as
`zero · `suc `zero
which has no type. Such terms that
exist independent of types are sometimes called preterms or
raw terms. Here we are going to replace the type Term
of
raw terms by the more sophisticated type Γ ⊢ A
of inherently
typed terms, which in context Γ
have type A
.
While these two choices fit well, they are independent. One can use de Bruijn indices in raw terms, or (with more difficulty) have inherently typed terms with names. In Chapter Untyped, we will introduce terms with de Bruijn indices that are inherently scoped but not typed.
A second example
De Bruijn indices can be tricky to get the hang of, so before proceeding further let’s consider a second example. Here is the term that adds two naturals:
plus : Term
plus = μ "+" ⇒ ƛ "m" ⇒ ƛ "n" ⇒
case ` "m"
[zero⇒ ` "n"
suc "m" ⇒ `suc (` "+" · ` "m" · ` "n") ]
Note variable "m"
is bound twice, once in a lambda abstraction
and once in the successor branch of the case. Any appearance
of "m"
in the successor branch must refer to the latter
binding, due to shadowing.
Here is its corresponding type derivation:
⊢plus : ∅ ⊢ plus ⦂ `ℕ ⇒ `ℕ ⇒ `ℕ
⊢plus = ⊢μ (⊢ƛ (⊢ƛ (⊢case (⊢` ∋m) (⊢` ∋n)
(⊢suc (⊢` ∋+ · ⊢` ∋m′ · ⊢` ∋n′)))))
where
∋+ = (S ("+" ≠ "m") (S ("+" ≠ "n") (S ("+" ≠ "m") Z)))
∋m = (S ("m" ≠ "n") Z)
∋n = Z
∋m′ = Z
∋n′ = (S ("n" ≠ "m") Z)
The two definitions are in close correspondence, where in addition to the previous correspondences we have:
`zero
corresponds to⊢zero
`suc_
corresponds to⊢suc
case_[zero⇒_suc_⇒_]
corresponds to⊢case
μ_⇒_
corresponds to⊢μ
Note the two lookup judgments ∋m
and ∋m′
refer to two
different bindings of variables named "m"
. In contrast, the
two judgments ∋n
and ∋n′
both refer to the same binding
of "n"
but accessed in different contexts, the first where
"n"
is the last binding in the context, and the second after
"m"
is bound in the successor branch of the case.
Here is the term and its type derivation in the notation of this chapter:
plus : ∀ {Γ} → Γ ⊢ `ℕ ⇒ `ℕ ⇒ `ℕ
plus = μ ƛ ƛ case (# 1) (# 0) (`suc (# 3 · # 0 · # 1))
Reading from left to right, each de Bruijn index corresponds to a lookup derivation:
# 1
corresponds to∋m
# 0
corresponds to∋n
# 3
corresponds to∋+
# 0
corresponds to∋m′
# 1
corresponds to∋n′
The de Bruijn index counts the number of S
constructs in the
corresponding lookup derivation. Variable "n"
bound in the
inner abstraction is referred to as # 0
in the zero branch
of the case but as # 1
in the successor branch of the case,
because of the intervening binding. Variable "m"
bound in the
lambda abstraction is referred to by the first # 1
in the
code, while variable "m"
bound in the successor branch of the
case is referred to by the second # 0
. There is no
shadowing: with variable names, there is no way to refer to
the former binding in the scope of the latter, but with de
Bruijn indices it could be referred to as # 2
.
Order of presentation
In the current chapter, the use of inherentlytyped terms necessitates that we cannot introduce operations such as substitution or reduction without also showing that they preserve types. Hence, the order of presentation must change.
The syntax of terms now incorporates their typing rules, and the definition of values now incorporates the Canonical Forms lemma. The definition of substitution is somewhat more involved, but incorporates the trickiest part of the previous proof, the lemma establishing that substitution preserves types. The definition of reduction incorporates preservation, which no longer requires a separate proof.
Syntax
We now begin our formal development.
First, we get all our infix declarations out of the way. We list separately operators for judgments, types, and terms:
infix 4 _⊢_ infix 4 _∋_ infixl 5 _,_ infixr 7 _⇒_ infix 5 ƛ_ infix 5 μ_ infixl 7 _·_ infix 8 `suc_ infix 9 `_ infix 9 S_ infix 9 #_
Since terms are inherently typed, we must define types and contexts before terms.
Types
As before, we have just two types, functions and naturals. The formal definition is unchanged:
data Type : Set where _⇒_ : Type → Type → Type `ℕ : Type
Contexts
Contexts are as before, but we drop the names. Contexts are formalised as follows:
data Context : Set where ∅ : Context _,_ : Context → Type → Context
A context is just a list of types, with the type of the most
recently bound variable on the right. As before, we let Γ
and Δ
range over contexts. We write ∅
for the empty
context, and Γ , A
for the context Γ
extended by type A
.
For example
_ : Context _ = ∅ , `ℕ ⇒ `ℕ , `ℕ
is a context with two variables in scope, where the outer
bound one has type `ℕ ⇒ `ℕ
, and the inner bound one has
type `ℕ
.
Variables and the lookup judgment
Inherently typed variables correspond to the lookup judgment. They are represented by de Bruijn indices, and hence also correspond to natural numbers. We write
Γ ∋ A
for variables which in context Γ
have type A
. Their
formalisation looks exactly like the old lookup judgment, but
with all variable names dropped:
data _∋_ : Context → Type → Set where Z : ∀ {Γ A}  → Γ , A ∋ A S_ : ∀ {Γ A B} → Γ ∋ A  → Γ , B ∋ A
Constructor S
no longer requires an additional parameter,
since without names shadowing is no longer an issue. Now
constructors Z
and S
correspond even more closely to the
constructors here
and there
for the elementof relation
_∈_
on lists, as well as to constructors zero
and suc
for natural numbers.
For example, consider the following oldstyle lookup judgments:
∅ , "s" ⦂ `ℕ ⇒ `ℕ , "z" ⦂ `ℕ ∋ "z" ⦂ `ℕ
∅ , "s" ⦂ `ℕ ⇒ `ℕ , "z" ⦂ `ℕ ∋ "s" ⦂ `ℕ ⇒ `ℕ
They correspond to the following inherentlytyped variables:
_ : ∅ , `ℕ ⇒ `ℕ , `ℕ ∋ `ℕ _ = Z _ : ∅ , `ℕ ⇒ `ℕ , `ℕ ∋ `ℕ ⇒ `ℕ _ = S Z
In the given context, "z"
is represented by Z
(as the most recently bound variable),
and "s"
by S Z
(as the next most recently bound variable).
Terms and the typing judgment
Inherently typed terms correspond to the typing judgment. We write
Γ ⊢ A
for terms which in context Γ
has type A
. Their
formalisation looks exactly like the old typing judgment, but
with all terms and variable names dropped:
data _⊢_ : Context → Type → Set where `_ : ∀ {Γ A} → Γ ∋ A  → Γ ⊢ A ƛ_ : ∀ {Γ A B} → Γ , A ⊢ B  → Γ ⊢ A ⇒ B _·_ : ∀ {Γ A B} → Γ ⊢ A ⇒ B → Γ ⊢ A  → Γ ⊢ B `zero : ∀ {Γ}  → Γ ⊢ `ℕ `suc_ : ∀ {Γ} → Γ ⊢ `ℕ  → Γ ⊢ `ℕ case : ∀ {Γ A} → Γ ⊢ `ℕ → Γ ⊢ A → Γ , `ℕ ⊢ A  → Γ ⊢ A μ_ : ∀ {Γ A} → Γ , A ⊢ A  → Γ ⊢ A
The definition exploits the close correspondence between the structure of terms and the structure of a derivation showing that it is well typed: now we use the derivation as the term.
For example, consider the following oldstyle typing judgments:
∅ , "s" ⦂ `ℕ ⇒ `ℕ , "z" ⦂ `ℕ ⊢ ` "z" ⦂ `ℕ
∅ , "s" ⦂ `ℕ ⇒ `ℕ , "z" ⦂ `ℕ ⊢ ` "s" ⦂ `ℕ ⇒ `ℕ
∅ , "s" ⦂ `ℕ ⇒ `ℕ , "z" ⦂ `ℕ ⊢ ` "s" · ` "z" ⦂ `ℕ
∅ , "s" ⦂ `ℕ ⇒ `ℕ , "z" ⦂ `ℕ ⊢ ` "s" · (` "s" · ` "z") ⦂ `ℕ
∅ , "s" ⦂ `ℕ ⇒ `ℕ ⊢ (ƛ "z" ⇒ ` "s" · (` "s" · ` "z")) ⦂ `ℕ ⇒ `ℕ
∅ ⊢ ƛ "s" ⇒ ƛ "z" ⇒ ` "s" · (` "s" · ` "z")) ⦂ (`ℕ ⇒ `ℕ) ⇒ `ℕ ⇒ `ℕ
They correspond to the following inherentlytyped terms:
_ : ∅ , `ℕ ⇒ `ℕ , `ℕ ⊢ `ℕ _ = ` Z _ : ∅ , `ℕ ⇒ `ℕ , `ℕ ⊢ `ℕ ⇒ `ℕ _ = ` S Z _ : ∅ , `ℕ ⇒ `ℕ , `ℕ ⊢ `ℕ _ = ` S Z · ` Z _ : ∅ , `ℕ ⇒ `ℕ , `ℕ ⊢ `ℕ _ = ` S Z · (` S Z · ` Z) _ : ∅ , `ℕ ⇒ `ℕ ⊢ `ℕ ⇒ `ℕ _ = ƛ (` S Z · (` S Z · ` Z)) _ : ∅ ⊢ (`ℕ ⇒ `ℕ) ⇒ `ℕ ⇒ `ℕ _ = ƛ ƛ (` S Z · (` S Z · ` Z))
The final inherentlytyped term represents the Church numeral two.
Abbreviating de Bruijn indices
We can use a natural number to select a type from a context:
lookup : Context → ℕ → Type lookup (Γ , A) zero = A lookup (Γ , _) (suc n) = lookup Γ n lookup ∅ _ = ⊥elim impossible where postulate impossible : ⊥
We intend to apply the function only when the natural is
shorter than the length of the context, which we indicate by
postulating an impossible
term, just as we did
here.
Given the above, we can convert a natural to a corresponding de Bruijn index, looking up its type in the context:
count : ∀ {Γ} → (n : ℕ) → Γ ∋ lookup Γ n count {Γ , _} zero = Z count {Γ , _} (suc n) = S (count n) count {∅} _ = ⊥elim impossible where postulate impossible : ⊥
This requires the same trick as before.
We can then introduce a convenient abbreviation for variables:
#_ : ∀ {Γ} → (n : ℕ) → Γ ⊢ lookup Γ n # n = ` count n
With this abbreviation, we can rewrite the Church numeral two more compactly:
_ : ∅ ⊢ (`ℕ ⇒ `ℕ) ⇒ `ℕ ⇒ `ℕ _ = ƛ ƛ (# 1 · (# 1 · # 0))
Test examples
We repeat the test examples from Chapter Lambda. You can find them here for comparison.
First, computing two plus two on naturals:
two : ∀ {Γ} → Γ ⊢ `ℕ two = `suc `suc `zero plus : ∀ {Γ} → Γ ⊢ `ℕ ⇒ `ℕ ⇒ `ℕ plus = μ ƛ ƛ (case (# 1) (# 0) (`suc (# 3 · # 0 · # 1))) 2+2 : ∀ {Γ} → Γ ⊢ `ℕ 2+2 = plus · two · two
We generalise to arbitrary contexts because later we will give examples
where two
appears nested inside binders.
Next, computing two plus two on Church numerals:
Ch : Type → Type Ch A = (A ⇒ A) ⇒ A ⇒ A twoᶜ : ∀ {Γ A} → Γ ⊢ Ch A twoᶜ = ƛ ƛ (# 1 · (# 1 · # 0)) plusᶜ : ∀ {Γ A} → Γ ⊢ Ch A ⇒ Ch A ⇒ Ch A plusᶜ = ƛ ƛ ƛ ƛ (# 3 · # 1 · (# 2 · # 1 · # 0)) sucᶜ : ∀ {Γ} → Γ ⊢ `ℕ ⇒ `ℕ sucᶜ = ƛ `suc (# 0) 2+2ᶜ : ∀ {Γ} → Γ ⊢ `ℕ 2+2ᶜ = plusᶜ · twoᶜ · twoᶜ · sucᶜ · `zero
As before we generalise everything to arbitrary
contexts. While we are at it, we also generalise twoᶜ
and
plusᶜ
to Church numerals over arbitrary types.
Exercise (mul
) (recommended)
Write out the definition of a lambda term that multiplies two natural numbers, now adapted to the inherently typed DeBruijn representation.
 Your code goes here
Renaming
Renaming is a necessary prelude to substitution, enabling us to “rebase” a term from one context to another. It corresponds directly to the renaming result from the previous chapter, but here the theorem that ensures renaming preserves typing also acts as code that performs renaming.
As before, we first need an extension lemma that allows us to extend the context when we encounter a binder. Given a map from variables in one context to variables in another, extension yields a map from the first context extended to the second context similarly extended. It looks exactly like the old extension lemma, but with all names and terms dropped:
ext : ∀ {Γ Δ} → (∀ {A} → Γ ∋ A → Δ ∋ A)  → (∀ {A B} → Γ , B ∋ A → Δ , B ∋ A) ext ρ Z = Z ext ρ (S x) = S (ρ x)
Let ρ
be the name of the map that takes variables in Γ
to variables in Δ
. Consider the de Bruijn index of the
variable in Γ , B
:

If it is
Z
, which has typeB
inΓ , B
, then we returnZ
, which also has typeB
inΔ , B
. 
If it is
S x
, for some variablex
inΓ
, thenρ x
is a variable inΔ
, and henceS (ρ x)
is a variable inΔ , B
.
With extension under our belts, it is straightforward to define renaming. If variables in one context map to variables in another, then terms in the first context map to terms in the second:
rename : ∀ {Γ Δ} → (∀ {A} → Γ ∋ A → Δ ∋ A)  → (∀ {A} → Γ ⊢ A → Δ ⊢ A) rename ρ (` x) = ` (ρ x) rename ρ (ƛ N) = ƛ (rename (ext ρ) N) rename ρ (L · M) = (rename ρ L) · (rename ρ M) rename ρ (`zero) = `zero rename ρ (`suc M) = `suc (rename ρ M) rename ρ (case L M N) = case (rename ρ L) (rename ρ M) (rename (ext ρ) N) rename ρ (μ N) = μ (rename (ext ρ) N)
Let ρ
be the name of the map that takes variables in Γ
to variables in Δ
. Let’s unpack the first three cases:

If the term is a variable, simply apply
ρ
. 
If the term is an abstraction, use the previous result to extend the map
ρ
suitably and recursively rename the body of the abstraction. 
If the term is an application, recursively rename both the function and the argument.
The remaining cases are similar, recursing on each subterm, and extending the map whenever the construct introduces a bound variable.
Whereas before renaming was a result that carried evidence that a term is well typed in one context to evidence that it is well typed in another context, now it actually transforms the term, suitably altering the bound variables. Typechecking the code in Agda ensures that it is only passed and returns terms that are well typed by the rules of simplytyped lambda calculus.
Here is an example of renaming a term with one free and one bound variable:
M₀ : ∅ , `ℕ ⇒ `ℕ ⊢ `ℕ ⇒ `ℕ M₀ = ƛ (# 1 · (# 1 · # 0)) M₁ : ∅ , `ℕ ⇒ `ℕ , `ℕ ⊢ `ℕ ⇒ `ℕ M₁ = ƛ (# 2 · (# 2 · # 0)) _ : rename S_ M₀ ≡ M₁ _ = refl
In general, rename S_
will increment the de Bruijn index for
each free variable by one, while leaving the index for each
bound variable unchanged. The code achieves this naturally:
the map originally increments each variable by one, and is
extended for each bound variable by a map that leaves it
unchanged.
We will see below that renaming by S_
plays a key role in
substitution. For traditional uses of de Bruijn indices
without inherent typing, this is a little tricky. The code
keeps count of a number where all greater indexes are free and
all smaller indexes bound, and increment only indexes greater
than the number. It’s easy to have offbyone errors. But
it’s hard to imagine an offbyone error that preserves
typing, and hence the Agda code for inherentlytyped de Bruijn
terms is inherently reliable.
Simultaneous Substitution
Because de Bruijn indices free us of concerns with renaming, it becomes easy to provide a definition of substitution that is more general than the one considered previously. Instead of substituting a closed term for a single variable, it provides a map that takes each free variable of the original term to another term. Further, the substituted terms are over an arbitrary context, and need not be closed.
The structure of the definition and the proof is remarkably close to that for renaming. Again, we first need an extension lemma that allows us to extend the context when we encounter a binder. Whereas renaming concerned a map from variables in one context to variables in another, substitution takes a map from variables in one context to terms in another. Given a map from variables in one context to terms over another, extension yields a map from the first context extended to the second context similarly extended:
exts : ∀ {Γ Δ} → (∀ {A} → Γ ∋ A → Δ ⊢ A)  → (∀ {A B} → Γ , B ∋ A → Δ , B ⊢ A) exts σ Z = ` Z exts σ (S x) = rename S_ (σ x)
Let σ
be the name of the map that takes variables in Γ
to terms over Δ
. Consider the de Bruijn index of the
variable in Γ , B
:

If it is
Z
, which has typeB
inΓ , B
, then we return the term` Z
, which also has typeB
inΔ , B
. 
If it is
S x
, for some variablex
inΓ
, thenσ x
is a term inΔ
, and hencerename S_ (σ x)
is a term inΔ , B
.
This is why we had to define renaming first, since
we require it to convert a term over context Δ
to a term over the extended context Δ , B
.
With extension under our belts, it is straightforward to define substitution. If variables in one context map to terms over another, then terms in the first context map to terms in the second:
subst : ∀ {Γ Δ} → (∀ {A} → Γ ∋ A → Δ ⊢ A)  → (∀ {A} → Γ ⊢ A → Δ ⊢ A) subst σ (` k) = σ k subst σ (ƛ N) = ƛ (subst (exts σ) N) subst σ (L · M) = (subst σ L) · (subst σ M) subst σ (`zero) = `zero subst σ (`suc M) = `suc (subst σ M) subst σ (case L M N) = case (subst σ L) (subst σ M) (subst (exts σ) N) subst σ (μ N) = μ (subst (exts σ) N)
Let σ
be the name of the map that takes variables in Γ
to terms over Δ
. Let’s unpack the first three cases:

If the term is a variable, simply apply
σ
. 
If the term is an abstraction, use the previous result to extend the map
σ
suitably and recursively substitute over the body of the abstraction. 
If the term is an application, recursively substitute over both the function and the argument.
The remaining cases are similar, recursing on each subterm, and extending the map whenever the construct introduces a bound variable.
Single substitution
From the general case of substitution for multiple free variables it is easy to define the special case of substitution for one free variable:
_[_] : ∀ {Γ A B} → Γ , B ⊢ A → Γ ⊢ B  → Γ ⊢ A _[_] {Γ} {A} {B} N M = subst {Γ , B} {Γ} σ {A} N where σ : ∀ {A} → Γ , B ∋ A → Γ ⊢ A σ Z = M σ (S x) = ` x
In a term of type A
over context Γ , B
, we replace the
variable of type B
by a term of type B
over context Γ
.
To do so, we use a map from the context Γ , B
to the context
Γ
, that maps the last variable in the context to the term of
type B
and every other free variable to itself.
Consider the previous example:
(ƛ "z" ⇒ ` "s" · (` "s" · ` "z")) [ "s" := sucᶜ ]
yieldsƛ "z" ⇒ sucᶜ · (sucᶜ · ` "z")
Here is the example formalised:
M₂ : ∅ , `ℕ ⇒ `ℕ ⊢ `ℕ ⇒ `ℕ M₂ = ƛ # 1 · (# 1 · # 0) M₃ : ∅ ⊢ `ℕ ⇒ `ℕ M₃ = ƛ `suc # 0 M₄ : ∅ ⊢ `ℕ ⇒ `ℕ M₄ = ƛ (ƛ `suc # 0) · ((ƛ `suc # 0) · # 0) _ : M₂ [ M₃ ] ≡ M₄ _ = refl
Previously, we presented an example of substitution that we did not implement, since it needed to rename the bound variable to avoid capture:
(ƛ "x" ⇒ ` "x" · ` "y") [ "y" := ` "x" · `zero ]
should yieldƛ "z" ⇒ ` "z" · (` "x" · `zero)
Say the bound "x"
has type `ℕ ⇒ `ℕ
, the substituted
"y"
has type `ℕ
, and the free "x"
also has type
`ℕ ⇒ `ℕ
. Here is the example formalised:
M₅ : ∅ , `ℕ ⇒ `ℕ , `ℕ ⊢ (`ℕ ⇒ `ℕ) ⇒ `ℕ M₅ = ƛ # 0 · # 1 M₆ : ∅ , `ℕ ⇒ `ℕ ⊢ `ℕ M₆ = # 0 · `zero M₇ : ∅ , `ℕ ⇒ `ℕ ⊢ (`ℕ ⇒ `ℕ) ⇒ `ℕ M₇ = ƛ (# 0 · (# 1 · `zero)) _ : M₅ [ M₆ ] ≡ M₇ _ = refl
The logician Haskell Curry observed that getting the definition of substitution right can be a tricky business. It can be even trickier when using de Bruijn indices, which can often be hard to decipher. Under the current approach, any definition of substitution must, of necessity, preserves types. While this makes the definition more involved, it means that once it is done the hardest work is out of the way. And combining definition with proof makes it harder for errors to sneak in.
Values
The definition of value is much as before, save that the added types incorporate the same information found in the Canonical Forms lemma:
data Value : ∀ {Γ A} → Γ ⊢ A → Set where Vƛ : ∀ {Γ A B} {N : Γ , A ⊢ B}  → Value (ƛ N) Vzero : ∀ {Γ}  → Value (`zero {Γ}) Vsuc : ∀ {Γ} {V : Γ ⊢ `ℕ} → Value V  → Value (`suc V)
Here zero
requires an implicit parameter to aid inference,
much in the same way that []
did in
Lists.
Reduction
The reduction rules are the same as those given earlier, save
that for each term we must specify its types. As before, we
have compatibility rules that reduce a part of a term,
labelled with ξ
, and rules that simplify a constructor
combined with a destructor, labelled with β
:
infix 2 _—→_ data _—→_ : ∀ {Γ A} → (Γ ⊢ A) → (Γ ⊢ A) → Set where ξ·₁ : ∀ {Γ A B} {L L′ : Γ ⊢ A ⇒ B} {M : Γ ⊢ A} → L —→ L′  → L · M —→ L′ · M ξ·₂ : ∀ {Γ A B} {V : Γ ⊢ A ⇒ B} {M M′ : Γ ⊢ A} → Value V → M —→ M′  → V · M —→ V · M′ βƛ : ∀ {Γ A B} {N : Γ , A ⊢ B} {W : Γ ⊢ A} → Value W  → (ƛ N) · W —→ N [ W ] ξsuc : ∀ {Γ} {M M′ : Γ ⊢ `ℕ} → M —→ M′  → `suc M —→ `suc M′ ξcase : ∀ {Γ A} {L L′ : Γ ⊢ `ℕ} {M : Γ ⊢ A} {N : Γ , `ℕ ⊢ A} → L —→ L′  → case L M N —→ case L′ M N βzero : ∀ {Γ A} {M : Γ ⊢ A} {N : Γ , `ℕ ⊢ A}  → case `zero M N —→ M βsuc : ∀ {Γ A} {V : Γ ⊢ `ℕ} {M : Γ ⊢ A} {N : Γ , `ℕ ⊢ A} → Value V  → case (`suc V) M N —→ N [ V ] βμ : ∀ {Γ A} {N : Γ , A ⊢ A}  → μ N —→ N [ μ N ]
The definition states that M —→ N
can only hold of terms M
and N
which both have type Γ ⊢ A
for some context Γ
and type A
. In other words, it is builtin to our
definition that reduction preserves types. There is no
separate Preservation theorem to prove. The Agda typechecker
validates that each term preserves types. In the case of β
rules, preservation depends on the fact that substitution
preserves types, which is builtin to our
definition of substitution.
Reflexive and transitive closure
The reflexive and transitive closure is exactly as before. We simply cutandpaste the previous definition:
infix 2 _—↠_ infix 1 begin_ infixr 2 _—→⟨_⟩_ infix 3 _∎ data _—↠_ : ∀ {Γ A} → (Γ ⊢ A) → (Γ ⊢ A) → Set where _∎ : ∀ {Γ A} (M : Γ ⊢ A)  → M —↠ M _—→⟨_⟩_ : ∀ {Γ A} (L : Γ ⊢ A) {M N : Γ ⊢ A} → L —→ M → M —↠ N  → L —↠ N begin_ : ∀ {Γ A} {M N : Γ ⊢ A} → M —↠ N  → M —↠ N begin M—↠N = M—↠N
Examples
We reiterate each of our previous examples. First, the Church numeral two applied to the successor function and zero yields the natural number two:
_ : twoᶜ · sucᶜ · `zero {∅} —↠ `suc `suc `zero _ = begin twoᶜ · sucᶜ · `zero —→⟨ ξ·₁ (βƛ Vƛ) ⟩ (ƛ (sucᶜ · (sucᶜ · # 0))) · `zero —→⟨ βƛ Vzero ⟩ sucᶜ · (sucᶜ · `zero) —→⟨ ξ·₂ Vƛ (βƛ Vzero) ⟩ sucᶜ · `suc `zero —→⟨ βƛ (Vsuc Vzero) ⟩ `suc (`suc `zero) ∎
As before, we need to supply an explicit context to `zero
.
Next, a sample reduction demonstrating that two plus two is four:
_ : plus {∅} · two · two —↠ `suc `suc `suc `suc `zero _ = plus · two · two —→⟨ ξ·₁ (ξ·₁ βμ) ⟩ (ƛ ƛ case (` S Z) (` Z) (`suc (plus · ` Z · ` S Z))) · two · two —→⟨ ξ·₁ (βƛ (Vsuc (Vsuc Vzero))) ⟩ (ƛ case two (` Z) (`suc (plus · ` Z · ` S Z))) · two —→⟨ βƛ (Vsuc (Vsuc Vzero)) ⟩ case two two (`suc (plus · ` Z · two)) —→⟨ βsuc (Vsuc Vzero) ⟩ `suc (plus · `suc `zero · two) —→⟨ ξsuc (ξ·₁ (ξ·₁ βμ)) ⟩ `suc ((ƛ ƛ case (` S Z) (` Z) (`suc (plus · ` Z · ` S Z))) · `suc `zero · two) —→⟨ ξsuc (ξ·₁ (βƛ (Vsuc Vzero))) ⟩ `suc ((ƛ case (`suc `zero) (` Z) (`suc (plus · ` Z · ` S Z))) · two) —→⟨ ξsuc (βƛ (Vsuc (Vsuc Vzero))) ⟩ `suc (case (`suc `zero) (two) (`suc (plus · ` Z · two))) —→⟨ ξsuc (βsuc Vzero) ⟩ `suc (`suc (plus · `zero · two)) —→⟨ ξsuc (ξsuc (ξ·₁ (ξ·₁ βμ))) ⟩ `suc (`suc ((ƛ ƛ case (` S Z) (` Z) (`suc (plus · ` Z · ` S Z))) · `zero · two)) —→⟨ ξsuc (ξsuc (ξ·₁ (βƛ Vzero))) ⟩ `suc (`suc ((ƛ case `zero (` Z) (`suc (plus · ` Z · ` S Z))) · two)) —→⟨ ξsuc (ξsuc (βƛ (Vsuc (Vsuc Vzero)))) ⟩ `suc (`suc (case `zero (two) (`suc (plus · ` Z · two)))) —→⟨ ξsuc (ξsuc βzero) ⟩ `suc (`suc (`suc (`suc `zero))) ∎
And finally, a similar sample reduction for Church numerals:
_ : plusᶜ · twoᶜ · twoᶜ · sucᶜ · `zero —↠ `suc `suc `suc `suc `zero {∅} _ = begin plusᶜ · twoᶜ · twoᶜ · sucᶜ · `zero —→⟨ ξ·₁ (ξ·₁ (ξ·₁ (βƛ Vƛ))) ⟩ (ƛ ƛ ƛ twoᶜ · ` S Z · (` S S Z · ` S Z · ` Z)) · twoᶜ · sucᶜ · `zero —→⟨ ξ·₁ (ξ·₁ (βƛ Vƛ)) ⟩ (ƛ ƛ twoᶜ · ` S Z · (twoᶜ · ` S Z · ` Z)) · sucᶜ · `zero —→⟨ ξ·₁ (βƛ Vƛ) ⟩ (ƛ twoᶜ · sucᶜ · (twoᶜ · sucᶜ · ` Z)) · `zero —→⟨ βƛ Vzero ⟩ twoᶜ · sucᶜ · (twoᶜ · sucᶜ · `zero) —→⟨ ξ·₁ (βƛ Vƛ) ⟩ (ƛ sucᶜ · (sucᶜ · ` Z)) · (twoᶜ · sucᶜ · `zero) —→⟨ ξ·₂ Vƛ (ξ·₁ (βƛ Vƛ)) ⟩ (ƛ sucᶜ · (sucᶜ · ` Z)) · ((ƛ sucᶜ · (sucᶜ · ` Z)) · `zero) —→⟨ ξ·₂ Vƛ (βƛ Vzero) ⟩ (ƛ sucᶜ · (sucᶜ · ` Z)) · (sucᶜ · (sucᶜ · `zero)) —→⟨ ξ·₂ Vƛ (ξ·₂ Vƛ (βƛ Vzero)) ⟩ (ƛ sucᶜ · (sucᶜ · ` Z)) · (sucᶜ · `suc `zero) —→⟨ ξ·₂ Vƛ (βƛ (Vsuc Vzero)) ⟩ (ƛ sucᶜ · (sucᶜ · ` Z)) · `suc (`suc `zero) —→⟨ βƛ (Vsuc (Vsuc Vzero)) ⟩ sucᶜ · (sucᶜ · `suc (`suc `zero)) —→⟨ ξ·₂ Vƛ (βƛ (Vsuc (Vsuc Vzero))) ⟩ sucᶜ · `suc (`suc (`suc `zero)) —→⟨ βƛ (Vsuc (Vsuc (Vsuc Vzero))) ⟩ `suc (`suc (`suc (`suc `zero))) ∎
Values do not reduce
We have now completed all the definitions, which of necessity subsumed some of the propositions from the earlier development: Canonical Forms, Substitution preserves types, and Preservation. We now turn to proving the remaining results from the previous development.
Exercise V¬—→
Following the previous development, show values do not reduce, and its corollary, terms that reduce are not values.
 Your code goes here
Progress
As before, every term that is well typed and closed is either a value or takes a reduction step. The formulation of progress is just as before, but annotated with types:
data Progress {A} (M : ∅ ⊢ A) : Set where step : ∀ {N : ∅ ⊢ A} → M —→ N  → Progress M done : Value M  → Progress M
The statement and proof of progress is much as before, appropriately annotated. We no longer need to explicitly refer to the Canonical Forms lemma, since it is builtin to the definition of value:
progress : ∀ {A} → (M : ∅ ⊢ A) → Progress M progress (` ()) progress (ƛ N) = done Vƛ progress (L · M) with progress L ...  step L—→L′ = step (ξ·₁ L—→L′) ...  done Vƛ with progress M ...  step M—→M′ = step (ξ·₂ Vƛ M—→M′) ...  done VM = step (βƛ VM) progress (`zero) = done Vzero progress (`suc M) with progress M ...  step M—→M′ = step (ξsuc M—→M′) ...  done VM = done (Vsuc VM) progress (case L M N) with progress L ...  step L—→L′ = step (ξcase L—→L′) ...  done Vzero = step (βzero) ...  done (Vsuc VL) = step (βsuc VL) progress (μ N) = step (βμ)
Evaluation
Before, we combined progress and preservation to evaluate a term. We can do much the same here, but we no longer need to explicitly refer to preservation, since it is builtin to the definition of reduction.
As previously, gas is specified by a natural number:
data Gas : Set where gas : ℕ → Gas
When our evaluator returns a term N
, it will either give evidence that
N
is a value or indicate that it ran out of gas:
data Finished {Γ A} (N : Γ ⊢ A) : Set where done : Value N  → Finished N outofgas :  Finished N
Given a term L
of type A
, the evaluator will, for some N
, return
a reduction sequence from L
to N
and an indication of whether
reduction finished:
data Steps : ∀ {A} → ∅ ⊢ A → Set where steps : ∀ {A} {L N : ∅ ⊢ A} → L —↠ N → Finished N  → Steps L
The evaluator takes gas and a term and returns the corresponding steps:
eval : ∀ {A} → Gas → (L : ∅ ⊢ A)  → Steps L eval (gas zero) L = steps (L ∎) outofgas eval (gas (suc m)) L with progress L ...  done VL = steps (L ∎) (done VL) ...  step {M} L—→M with eval (gas m) M ...  steps M—↠N fin = steps (L —→⟨ L—→M ⟩ M—↠N) fin
The definition is a little simpler than previously, as we no longer need to invoke preservation.
Examples
We reiterate each of our previous examples. We redefine the term
sucμ
that loops forever:
sucμ : ∅ ⊢ `ℕ sucμ = μ (`suc (# 0))
To compute the first three steps of the infinite reduction sequence, we evaluate with three steps worth of gas:
_ : eval (gas 3) sucμ ≡ steps (μ `suc ` Z —→⟨ βμ ⟩ `suc (μ `suc ` Z) —→⟨ ξsuc βμ ⟩ `suc (`suc (μ `suc ` Z)) —→⟨ ξsuc (ξsuc βμ) ⟩ `suc (`suc (`suc (μ `suc ` Z))) ∎) outofgas _ = refl
The Church numeral two applied to successor and zero:
_ : eval (gas 100) (twoᶜ · sucᶜ · `zero) ≡ steps ((ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · `zero —→⟨ ξ·₁ (βƛ Vƛ) ⟩ (ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · `zero —→⟨ βƛ Vzero ⟩ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · `zero) —→⟨ ξ·₂ Vƛ (βƛ Vzero) ⟩ (ƛ `suc ` Z) · `suc `zero —→⟨ βƛ (Vsuc Vzero) ⟩ `suc (`suc `zero) ∎) (done (Vsuc (Vsuc Vzero))) _ = refl
Two plus two is four:
_ : eval (gas 100) (plus · two · two) ≡ steps ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · `suc (`suc `zero) · `suc (`suc `zero) —→⟨ ξ·₁ (ξ·₁ βμ) ⟩ (ƛ (ƛ case (` (S Z)) (` Z) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · ` (S Z))))) · `suc (`suc `zero) · `suc (`suc `zero) —→⟨ ξ·₁ (βƛ (Vsuc (Vsuc Vzero))) ⟩ (ƛ case (`suc (`suc `zero)) (` Z) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · ` (S Z)))) · `suc (`suc `zero) —→⟨ βƛ (Vsuc (Vsuc Vzero)) ⟩ case (`suc (`suc `zero)) (`suc (`suc `zero)) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · `suc (`suc `zero))) —→⟨ βsuc (Vsuc Vzero) ⟩ `suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · `suc `zero · `suc (`suc `zero)) —→⟨ ξsuc (ξ·₁ (ξ·₁ βμ)) ⟩ `suc ((ƛ (ƛ case (` (S Z)) (` Z) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · ` (S Z))))) · `suc `zero · `suc (`suc `zero)) —→⟨ ξsuc (ξ·₁ (βƛ (Vsuc Vzero))) ⟩ `suc ((ƛ case (`suc `zero) (` Z) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · ` (S Z)))) · `suc (`suc `zero)) —→⟨ ξsuc (βƛ (Vsuc (Vsuc Vzero))) ⟩ `suc case (`suc `zero) (`suc (`suc `zero)) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · `suc (`suc `zero))) —→⟨ ξsuc (βsuc Vzero) ⟩ `suc (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · `zero · `suc (`suc `zero))) —→⟨ ξsuc (ξsuc (ξ·₁ (ξ·₁ βμ))) ⟩ `suc (`suc ((ƛ (ƛ case (` (S Z)) (` Z) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · ` (S Z))))) · `zero · `suc (`suc `zero))) —→⟨ ξsuc (ξsuc (ξ·₁ (βƛ Vzero))) ⟩ `suc (`suc ((ƛ case `zero (` Z) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · ` (S Z)))) · `suc (`suc `zero))) —→⟨ ξsuc (ξsuc (βƛ (Vsuc (Vsuc Vzero)))) ⟩ `suc (`suc case `zero (`suc (`suc `zero)) (`suc ((μ (ƛ (ƛ case (` (S Z)) (` Z) (`suc (` (S (S (S Z))) · ` Z · ` (S Z)))))) · ` Z · `suc (`suc `zero)))) —→⟨ ξsuc (ξsuc βzero) ⟩ `suc (`suc (`suc (`suc `zero))) ∎) (done (Vsuc (Vsuc (Vsuc (Vsuc Vzero))))) _ = refl
And the corresponding term for Church numerals:
_ : eval (gas 100) (plusᶜ · twoᶜ · twoᶜ · sucᶜ · `zero) ≡ steps ((ƛ (ƛ (ƛ (ƛ ` (S (S (S Z))) · ` (S Z) · (` (S (S Z)) · ` (S Z) · ` Z))))) · (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · `zero —→⟨ ξ·₁ (ξ·₁ (ξ·₁ (βƛ Vƛ))) ⟩ (ƛ (ƛ (ƛ (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · ` (S Z) · (` (S (S Z)) · ` (S Z) · ` Z)))) · (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · `zero —→⟨ ξ·₁ (ξ·₁ (βƛ Vƛ)) ⟩ (ƛ (ƛ (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · ` (S Z) · ((ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · ` (S Z) · ` Z))) · (ƛ `suc ` Z) · `zero —→⟨ ξ·₁ (βƛ Vƛ) ⟩ (ƛ (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · ((ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · ` Z)) · `zero —→⟨ βƛ Vzero ⟩ (ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · ((ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · `zero) —→⟨ ξ·₁ (βƛ Vƛ) ⟩ (ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · ((ƛ (ƛ ` (S Z) · (` (S Z) · ` Z))) · (ƛ `suc ` Z) · `zero) —→⟨ ξ·₂ Vƛ (ξ·₁ (βƛ Vƛ)) ⟩ (ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · ((ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · `zero) —→⟨ ξ·₂ Vƛ (βƛ Vzero) ⟩ (ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · ((ƛ `suc ` Z) · ((ƛ `suc ` Z) · `zero)) —→⟨ ξ·₂ Vƛ (ξ·₂ Vƛ (βƛ Vzero)) ⟩ (ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · ((ƛ `suc ` Z) · `suc `zero) —→⟨ ξ·₂ Vƛ (βƛ (Vsuc Vzero)) ⟩ (ƛ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · ` Z)) · `suc (`suc `zero) —→⟨ βƛ (Vsuc (Vsuc Vzero)) ⟩ (ƛ `suc ` Z) · ((ƛ `suc ` Z) · `suc (`suc `zero)) —→⟨ ξ·₂ Vƛ (βƛ (Vsuc (Vsuc Vzero))) ⟩ (ƛ `suc ` Z) · `suc (`suc (`suc `zero)) —→⟨ βƛ (Vsuc (Vsuc (Vsuc Vzero))) ⟩ `suc (`suc (`suc (`suc `zero))) ∎) (done (Vsuc (Vsuc (Vsuc (Vsuc Vzero))))) _ = refl
We omit the proof that reduction is deterministic, since it is tedious and almost identical to the previous proof.
Exercise mulexample
(recommended)
Using the evaluator, confirm that two times two is four.
 Your code goes here
Inherentlytyped is golden
Counting the lines of code is instructive. While this chapter covers the same formal development as the previous two chapters, it has much less code. Omitting all the examples, and all proofs that appear in Properties but not DeBruijn (such as the proof that reduction is deterministic), the number of lines of code is as follows:
Lambda 216
Properties 235
DeBruijn 275
The relation between the two approaches approximates the golden ratio: raw terms with a separate typing relation require about 1.6 times as much code as inherentlytyped terms with de Bruijn indices.
Unicode
This chapter uses the following unicode:
σ U+03C3 GREEK SMALL LETTER SIGMA (\Gs or \sigma)
₀ U+2080 SUBSCRIPT ZERO (\_0)
₃ U+20B3 SUBSCRIPT THREE (\_3)
₄ U+2084 SUBSCRIPT FOUR (\_4)
₅ U+2085 SUBSCRIPT FIVE (\_5)
₆ U+2086 SUBSCRIPT SIX (\_6)
₇ U+2087 SUBSCRIPT SEVEN (\_7)
≠ U+2260 NOT EQUAL TO (\=n)