For my thesis I've needed to use strictly positive types to avoid the problems which can occur with types of the form: νD. D → D, which were discussed briefly in my post System F is Amazing.

In order to ensure types are positive they should be restricted in some way. It is therefore convenient to have formation rules for types which ensure that we can't make bad types. My starting point was a wonderful paper by Andreas Abel and Thorsten Altenkirch [1]

Specifically, this paper describes interleaved inductive types, which is close to what I want. However, I also want to be able to have polymorphism using ∀. Since I'm interested in becoming more familiar with Agda (a language in which I'm still a novice) I've written the following as a prototype of my type formation rules, as an experiment.

First, we have to set up the appropriate libraries.

The first data type we define is a context of type variables. Since we are going to use de Bruijn indices, we can use a trick and say that the type variable context is really just a bound on the number of free variables. Since our types are not themselves typed (as our theory is impredicative) we are justified in doing this as the only information needed formation is that the variable is in a context.

We then define an ∈ relation which is true whenever a variable is below the bound.

It turns out to be convenient to define lemmas which perform inversions on data-types in Agda. In Coq, I would generally just use a tactic to perform the inversion. However, there are advantages to Agda's way of doing things, one of which is that it obviates a lot of playing with equational reasoning. So far I've yet to be forced to learn how equational reasoning works in Agda since I haven't had to use it. I'm a little frightened that things work so well!

Next we prove that ∈ is decidable, which is pretty straight forward since it is defined using ≤ on natural numbers.

Next we define the syntax for types. We have ① for unit, ι which creates a type variable, η which creates a fixed-point type variable, ⇒ which is the function type, × which is the type of a pair, ⊕ which is the type of an injection into the left or right, μ which makes a fixed point, and Π which is the type of a type abstraction.

Next we demonstrate the formation rules. Here we use two contexts. One is for ι variables, the other for η. The reason for segregating them is that we can allow ι variables to be placed more freely, while η variables need to be tightly controlled to ensure strictly positive types. The restriction manifests itself in the rule for ⇒, which allows formation with a type to the left of the arrow, only if the context for η vars is empty. The reader will notice that every time we go under a Π binder, the first context is increased in size, and every time we go under a μ, we increase the second. When we check that variables are in context, we refer to these segregated contexts with our two different type variables.

Again, we find it convenient to define a number of inversion rules, this type on type formation.

Now we would like to show that we can demonstrate whether or not a type is valid or not. This program will construct a proof of the validity of a given type, or reject it with a proof that one can not be constructed. Intermediate computations which attempt to construct subproofs or proofs that no subproof exists are done using with. This allows us to further pattern much, just as we would do with parameters of functions in Haskell. The notation (λ { (VarValid q) → p q }) is especially cool. It is a pattern matching lambda, which allows us to destruct the argument with each possible case. In this case, because of the type, there is exactly one case which allows us to prove that this type can not be valid.

Finally we give an example of some terms which can be typed, and some which can not. When there is no inhabitant of a pattern in Agda, we can use (). Agda is smart enough to know that there is no sense in a right-hand-side of a pattern which has no inhabitants because the type is indexed in such a way that inversion would yield no cases. It's very clever!

[1] A Predicative Strong Normalisation Proof for a λ-calculus with Interleaving Inductive Types

In order to ensure types are positive they should be restricted in some way. It is therefore convenient to have formation rules for types which ensure that we can't make bad types. My starting point was a wonderful paper by Andreas Abel and Thorsten Altenkirch [1]

Specifically, this paper describes interleaved inductive types, which is close to what I want. However, I also want to be able to have polymorphism using ∀. Since I'm interested in becoming more familiar with Agda (a language in which I'm still a novice) I've written the following as a prototype of my type formation rules, as an experiment.

First, we have to set up the appropriate libraries.

module StrictlyPositive where open import Data.Nat open import Data.Bool open import Data.List open import Data.Unit using (⊤) open import Relation.Nullary.Core open import Data.Sum

The first data type we define is a context of type variables. Since we are going to use de Bruijn indices, we can use a trick and say that the type variable context is really just a bound on the number of free variables. Since our types are not themselves typed (as our theory is impredicative) we are justified in doing this as the only information needed formation is that the variable is in a context.

Ctx : Set Ctx = ℕ

We then define an ∈ relation which is true whenever a variable is below the bound.

data _∈_ : ℕ → Ctx → Set where bounded : ∀ {n Δ} → suc n ≤ Δ → n ∈ Δ

It turns out to be convenient to define lemmas which perform inversions on data-types in Agda. In Coq, I would generally just use a tactic to perform the inversion. However, there are advantages to Agda's way of doing things, one of which is that it obviates a lot of playing with equational reasoning. So far I've yet to be forced to learn how equational reasoning works in Agda since I haven't had to use it. I'm a little frightened that things work so well!

inInversion : ∀ {n Δ} → n ∈ Δ → suc n ≤ Δ inInversion (bounded p) = p

Next we prove that ∈ is decidable, which is pretty straight forward since it is defined using ≤ on natural numbers.

inctx : ∀ n Δ → Dec (n ∈ Δ) inctx n Δ with (suc n) ≤? Δ inctx n Δ | no p = no (λ x → p (inInversion x)) inctx n Δ | yes q = yes (bounded q)

Next we define the syntax for types. We have ① for unit, ι which creates a type variable, η which creates a fixed-point type variable, ⇒ which is the function type, × which is the type of a pair, ⊕ which is the type of an injection into the left or right, μ which makes a fixed point, and Π which is the type of a type abstraction.

data Ty : Set where ① : Ty ι : ℕ → Ty η : ℕ → Ty _⇒_ : Ty → Ty → Ty _×_ : Ty → Ty → Ty _⊕_ : Ty → Ty → Ty μ : Ty → Ty Π : Ty → Ty

Next we demonstrate the formation rules. Here we use two contexts. One is for ι variables, the other for η. The reason for segregating them is that we can allow ι variables to be placed more freely, while η variables need to be tightly controlled to ensure strictly positive types. The restriction manifests itself in the rule for ⇒, which allows formation with a type to the left of the arrow, only if the context for η vars is empty. The reader will notice that every time we go under a Π binder, the first context is increased in size, and every time we go under a μ, we increase the second. When we check that variables are in context, we refer to these segregated contexts with our two different type variables.

data _·_⊢_ : Ctx → Ctx → Ty → Set where UnitValid : ∀ {Δ Δ+} → Δ · Δ+ ⊢ ① VarValid : ∀ {n Δ Δ+} → (n ∈ Δ) → Δ · Δ+ ⊢ (ι n) FixValid : ∀ {n Δ Δ+} → (n ∈ Δ+) → Δ · Δ+ ⊢ (η n) ImpValid : ∀ {Δ Δ+ α β} → Δ · 0 ⊢ α → Δ · Δ+ ⊢ β → Δ · Δ+ ⊢ (α ⇒ β) AndValid : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ α → Δ · Δ+ ⊢ β → Δ · Δ+ ⊢ (α × β) OrValid : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ α → Δ · Δ+ ⊢ β → Δ · Δ+ ⊢ (α ⊕ β) MuValid : ∀ {Δ Δ+ α} → Δ · (suc Δ+) ⊢ α → Δ · Δ+ ⊢ μ α AllValid : ∀ {Δ Δ+ α} → (suc Δ) · Δ+ ⊢ α → Δ · Δ+ ⊢ Π α

Again, we find it convenient to define a number of inversion rules, this type on type formation.

varValidInversion : ∀ {Δ Δ+ n} → Δ · Δ+ ⊢ ι n → n ∈ Δ varValidInversion {Δ} {Δ+} {n} (VarValid p) = p fixValidInversion : ∀ {Δ Δ+ n} → Δ · Δ+ ⊢ η n → n ∈ Δ+ fixValidInversion {Δ} {Δ+} {n} (FixValid p) = p orValidInversionL : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ (α ⊕ β) → (Δ · Δ+ ⊢ α) orValidInversionL {Δ} {Δ+} {α} {β} (OrValid p q) = p orValidInversionR : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ (α ⊕ β) → (Δ · Δ+ ⊢ β) orValidInversionR {Δ} {Δ+} {α} {β} (OrValid p q) = q impValidInversionL : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ (α ⇒ β) → (Δ · 0 ⊢ α) impValidInversionL {Δ} {Δ+} {α} {β} (ImpValid p q) = p impValidInversionR : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ (α ⇒ β) → (Δ · Δ+ ⊢ β) impValidInversionR {Δ} {Δ+} {α} {β} (ImpValid p q) = q andValidInversionL : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ (α × β) → (Δ · Δ+ ⊢ α) andValidInversionL {Δ} {Δ+} {α} {β} (AndValid p q) = p andValidInversionR : ∀ {Δ Δ+ α β} → Δ · Δ+ ⊢ (α × β) → (Δ · Δ+ ⊢ β) andValidInversionR {Δ} {Δ+} {α} {β} (AndValid p q) = q allValidInversion : ∀ {Δ Δ+ α} → Δ · Δ+ ⊢ Π α → ((suc Δ) · Δ+ ⊢ α) allValidInversion {Δ} {Δ+} {α} (AllValid p) = p muValidInversion : ∀ {Δ Δ+ α} → Δ · Δ+ ⊢ μ α → (Δ · (suc Δ+) ⊢ α) muValidInversion {Δ} {Δ+} {α} (MuValid p) = p

Now we would like to show that we can demonstrate whether or not a type is valid or not. This program will construct a proof of the validity of a given type, or reject it with a proof that one can not be constructed. Intermediate computations which attempt to construct subproofs or proofs that no subproof exists are done using with. This allows us to further pattern much, just as we would do with parameters of functions in Haskell. The notation (λ { (VarValid q) → p q }) is especially cool. It is a pattern matching lambda, which allows us to destruct the argument with each possible case. In this case, because of the type, there is exactly one case which allows us to prove that this type can not be valid.

isValid : ∀ Δ Δ+ α → Dec (Δ · Δ+ ⊢ α) isValid Δ Δ+ ① = yes UnitValid isValid Δ Δ+ (ι n) with (inctx n Δ) isValid Δ Δ+ (ι n) | yes p = yes (VarValid p) isValid Δ Δ+ (ι n) | no p = no (λ { (VarValid q) → p q }) isValid Δ Δ+ (η n) with (inctx n Δ+) isValid Δ Δ+ (η n) | yes p = yes (FixValid p) isValid Δ Δ+ (η n) | no p = no (λ { (FixValid q) → p q }) isValid Δ Δ+ (α ⇒ β) with isValid Δ 0 α isValid Δ Δ+ (α ⇒ β) | yes p with isValid Δ Δ+ β isValid Δ Δ+ (α ⇒ β) | yes p | yes q = yes (ImpValid p q) isValid Δ Δ+ (α ⇒ β) | yes p | no q = no (λ x → q (impValidInversionR x)) isValid Δ Δ+ (α ⇒ β) | no p = no (λ x → p (impValidInversionL x)) isValid Δ Δ+ (α × β) with isValid Δ Δ+ α isValid Δ Δ+ (α × β) | yes p with isValid Δ Δ+ β isValid Δ Δ+ (α × β) | yes p | yes q = yes (AndValid p q) isValid Δ Δ+ (α × β) | yes p | no q = no (λ x → q (andValidInversionR x)) isValid Δ Δ+ (α × β) | no p = no (λ x → p (andValidInversionL x)) isValid Δ Δ+ (α ⊕ β) with isValid Δ Δ+ α isValid Δ Δ+ (α ⊕ β) | yes p with isValid Δ Δ+ β isValid Δ Δ+ (α ⊕ β) | yes p | yes q = yes (OrValid p q) isValid Δ Δ+ (α ⊕ β) | yes p | no q = no (λ x → q (orValidInversionR x)) isValid Δ Δ+ (α ⊕ β) | no p = no (λ x → p (orValidInversionL x)) isValid Δ Δ+ (Π α) with isValid (suc Δ) Δ+ α isValid Δ Δ+ (Π α) | yes p = yes (AllValid p) isValid Δ Δ+ (Π α) | no p = no (λ x → p (allValidInversion x)) isValid Δ Δ+ (μ α) with isValid Δ (suc Δ+) α isValid Δ Δ+ (μ α) | yes p = yes (MuValid p) isValid Δ Δ+ (μ α) | no p = no (λ x → p (muValidInversion x))

Finally we give an example of some terms which can be typed, and some which can not. When there is no inhabitant of a pattern in Agda, we can use (). Agda is smart enough to know that there is no sense in a right-hand-side of a pattern which has no inhabitants because the type is indexed in such a way that inversion would yield no cases. It's very clever!

idTy : Ty idTy = Π ((ι zero) ⇒ (ι zero)) idValid : 0 · 0 ⊢ idTy idValid = AllValid (ImpValid (VarValid (bounded (s≤s z≤n))) (VarValid (bounded (s≤s z≤n)))) listTy : Ty listTy = Π (μ (① ⊕ ((ι 0) × (η 0)))) listValid : 0 · 0 ⊢ listTy listValid = AllValid (MuValid (OrValid UnitValid (AndValid (VarValid (bounded (s≤s z≤n))) (FixValid (bounded (s≤s z≤n)))))) notPos : Ty notPos = μ ((ι 0) ⇒ (ι 0)) notPosInvalid : ¬ (0 · 0 ⊢ notPos) notPosInvalid x with muValidInversion x notPosInvalid x | p with impValidInversionL p notPosInvalid x | p | q with varValidInversion q notPosInvalid x | p | q | bounded () dodgy : Ty dodgy = μ (μ ((ι 1) ⇒ (ι 0))) dodgyInvalid : ¬ (0 · 0 ⊢ dodgy) dodgyInvalid x with muValidInversion x dodgyInvalid x | p with muValidInversion p dodgyInvalid x | p | q with impValidInversionL q dodgyInvalid x | p | q | r with varValidInversion r dodgyInvalid x | p | q | r | bounded () isThisDodgy : Ty isThisDodgy = Π (μ (μ ((((ι 0) ⇒ (η 1)) ⇒ (η 0))))) isThisDodgyValid : ¬ (0 · 0 ⊢ isThisDodgy) isThisDodgyValid x with allValidInversion x isThisDodgyValid x | p with muValidInversion p isThisDodgyValid x | p | q with muValidInversion q isThisDodgyValid x | p | q | r with impValidInversionL r isThisDodgyValid x | p | q | r | s with impValidInversionR s isThisDodgyValid x | p | q | r | s | t with fixValidInversion t isThisDodgyValid x | p | q | r | s | t | bounded ()

[1] A Predicative Strong Normalisation Proof for a λ-calculus with Interleaving Inductive Types

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