NbE for lambda calculus with sums and for modal logic

Maxim Koltsov

What is NbE

Terms

Syntax

Normalization

What is NbE

Terms

Denotations

Syntax

Semantics

Evaluation

Reification

Normalization by Evaluation

Question: how to build NbE procedure?

Pick semantics model (= meta language) that makes reification easy

Danvy's Type Directed Partial Evaluation

Terms

a, b, c, x, y, ... : variables
\x -> t: lambda-abstractions
a b: application
<t1, t2>: pairs
p1 t, p2 t: projections

Danvy's Type Directed Partial Evaluation

reify_t : D -> Normal terms

Let D be some model

reify_o t = t, for o — atomic type
reify_(a -> b) t = \x -> reify_b (t (reflect_a x))
reify_(<a, b>) t = <p1 (reify_a t), p2 (reify_b t)>

reflect_o t = t, for o — atomic type
reflect_(a -> b) t = fun(s) { reflect_b (t (reify_a s)) }
reflect_(<a, b>) t = (reflect_a (p1 t), reflect_b (p2 t))

reflect_t : Neutral terms -> D

So, model (D) must support variables, functions and pairs

In Haskell it may be:

data Ty
  = Int
  | Fun Ty Ty
  deriving (Eq, Show)

data Term
  = Var String
  | VInt Int
  | Lam String Term
  | App Term Term
  | Pair Term Term
  | P1 Term
  | P2 Term
  deriving (Eq, Show)

-- | Semantics of terms
data Sem
  = SNeu Term
  -- ^ neutral: only atomic types and applications
  | SFun (Sem -> Sem)
  | SPair Sem Sem

Can we add sums?

Terms:

Left t, Right t
case t of (Left x -> t, Right x -> t)

reflect_(Either a b) t =
  case t of ...

NO — we need more complicated model and meta-language

Danvy: shift/reset

NbE and calculus properties

Soundness: if $ \Gamma \vdash p : A $, then in a model where $ \Gamma $ holds, $ A $ holds

$ \Gamma \vdash p : A $ — proof of proposition A under hypotheses Г (Curry-Howard)

Completeness: if in any model Г implies A, then there exists term $ p $ such that $ Г \vdash p : A $

$ Г \vdash p : A \to (w \Vdash Г => w \Vdash A) $

$ (\forall w.\ w \Vdash Г \to w \Vdash A) \to Г \vdash p : A $

Evaluation

Reification

Ilik's CPS model

Key ideas:

shift/reset can be written in full CPS
Use Kripke models in soundness and completeness formulations

Kripke model for lambda calculus

$ (K, \leq) $ — preorder of possible worlds

Relation of forcing: for world w in K and n-ary predicate X, $ w \Vdash X $ — n-ary relation on w, such that

For $ w' \geq w $, $ (w \Vdash X)(d_1, \ldots, d_n) \to (w', \Vdash X)(d_1, \ldots, d_n) $

For compound propositions (= types):

$ w \Vdash A \to B $ means for all $ w' \ge w $ $ w' \Vdash A $ implies $ w' \Vdash B $

$ w \Vdash <A, B> $ means $ w \Vdash A $ AND $w \Vdash B $

$ w \Vdash \text{Either}(A, B) $ means $ w \Vdash A $ OR $w \Vdash B $

Again, meta language must have functions, pairs and sums

Soundness: If $ \Gamma \vdash p:A $, then in any model for any world, if $ w \Vdash \Gamma $, then $ w \Vdash A $

Completeness: if in any model, in any world $ w \Vdash \Gamma $ implies $ w \Vdash A $, then there exists a term $ p $ such that $ \Gamma \vdash p:A $

We can construct an Universal Model for proving completeness

Universal model

$ K $ (possible worlds) will be a set of different contexts $ \Gamma $

$ \Gamma \leq \Gamma' $ if $ \Gamma \subseteq \Gamma' $

Forcing: $ \Gamma \Vdash A $ — a set of normal form derivations $ \Gamma \vdash^{nf} A $, where $ A $ — closed and atomic

Then, for example, by definition

$ \Gamma \Vdash A \to B $ when for all $ \Gamma' \geq \Gamma $ $\Gamma' \vdash^{nf} A $ implies $ \Gamma' \vdash^{nf} B $ — a function in meta-language

Kripke model for logic with sums

New binary relation: $ w \Vdash_\bot^C $ — world is exploding at formula $ C $

Strong forcing $ w \Vdash_s A $ defined as before for atomic formulas

Non-strong forcing: $ w \Vdash A $ if

$$ \forall C\, \forall w' \geq w \left( \forall w'' \geq w'\, w'' \Vdash_s A \to w'' \Vdash_\bot^C \right) \to w' \Vdash_\bot^C $$

Extend strong forcing to all formulas:

$ w \Vdash_s A \to B $ if for all $ w' \geq w $ $ w' \Vdash A $ implies $ w' \Vdash B $
$ w \Vdash_s <A, B> $ if $ w \Vdash A $ AND $ w \Vdash B $
...

Universal model for Kripke model with sums

K is the set of contexts again
Strong forcing: $ \Gamma \Vdash_s A $ if there is a derivation in normal form $ \Gamma \vdash^{nf} A $, for atomic formula $ A $
Exploding: $ \Gamma \Vdash_\bot^C $ if there is a derivation in normal form $ \Gamma \vdash^{nf} A $

Forcing:

$$ \forall C\, \forall \Gamma' \geq \Gamma \left( \forall \Gamma'' \geq \Gamma'\, \Gamma'' \Vdash_s A \to \Gamma'' \vdash^{nf} C \right) \to \Gamma' \vdash^{nf} C $$

In Haskell:

data Sem
  = Sem ((StrongForcing -> Term) -> Term)

data StrongForcing
  = FInt Term
  | FPair Sem Sem
  | FSum (Either Sem Sem)
  | FFun (Sem -> Sem)

Reification

"unit": $ \eta: w \Vdash_s A \to w \Vdash A $, $ \eta(a) = \lambda k.\, k a $

"bind": $ * : (\forall w' \geq w\, w' \Vdash_s A \to w' \Vdash B) \to w \Vdash A \to w \Vdash B $,

$\phi*a = \lambda k.\, a\ (\lambda b.\, \phi\ b\ k) $

Forcing ($ w \Vdash A $) is kinda "Cont" monad storing strong forcing ($ w \Vdash_s A $)

"run": $ \mu : w \Vdash X \to w \Vdash_s X $ for atomic $ X $, by definition of $ \Vdash $

Reification

reify_X (a) = $ \mu(a) $, $X$ — atomic

reify_(A->B) (f) = $ \eta\left(\lambda s.\, \lambda x.\, \text{reify}_B (s (\text{reflect}_A x))\right) $

reify_(Either A B) (a) =

$$ \eta\left( \lambda s.\, (\text{Left}\ l \to \text{reify}_A l; \text{Right}\ r \to \text{reify}_B r)\right) $$

reflect_X (a) = $ \eta(a) $, $ X $ — atomic

reflect_(A->B) (f) = $ \eta\left( \lambda x.\, \text{reflect}_B (f (\text{reify}_A x)) \right) $

reflect_(Either A B) (t) =

$$ \lambda k.\, \text{case}\ t\ \text{of}\ \Bigl[\text{Left}\ l \to k (\text{Left}\ (\text{reflect}_A l));$$

$$ \text{Right}\ r \to k (\text{Right}\ (\text{reflect}_B r)) \Bigr]$$

Examples

t :: Either Church () -> Church -> Church
t = \c -> \n0 ->
  case c of
    Left n -> add n n0
    Right _ -> \f -> \x -> f (n0 f x)

nbe(t (Left 3)):

t' :: Church -> Church
t' = \n0 -> \f -> \x -> f $ f $ f (n0 $ \k -> f k) x

nbe(t (Right ())):

t' :: Church -> Church
t' = \n0 -> \f -> \x -> f (n0 $ \k -> f k) x

Modal logic

New type: $ \Box A $

New terms:

quot t : $\Box A$ if $ t : A $

let box u = t in e: A if $ t : \Box A $ and $ \Gamma; \Delta, u:A \vdash e : B $

Two contexts: $ \Gamma; \Delta \vdash t : A $

exp :: Int -> [] (Int -> Int)
exp b =
  if zero b
    then box (\_ -> 1)
    else
      let box u = exp (b - 1)
      in box (\x -> x * (u x))
      
\x -> let box u = exp 3 in u x :: Int -> Int

Kripke CPS model

Strong forcing:

$ \Gamma; \Delta \Vdash_s \text{box}\ A $ is a pair of $ \Gamma; \emptyset \vdash t $ and $ \Gamma; \Delta \Vdash A $

eval(box t) = $ \eta( <\text{graft}\ t, \text{eval}\ t>) $

eval(let box u = t in e) =

$$ \lambda k.\, (\text{eval}\ t) (<t, s> \to (\text{eval}\ b) k) $$

reify([] A)(t) = $ t (<t, s> \to \text{box}\ t) $

reflect([] A)(t) =

$$ \lambda k.\, \text{let box}\ u = t\ \text{in}\ k <u, \text{reflect}_A u> $$

Example

nbe (exp 3) :: [] (Int -> Int)
nbe (exp 3) = box \x -> x * ( (\y -> y * ((\z -> ...)) y) x)

nbe (exp 3)

Full normalization: nbe(let box u = exp 3 in u)

nbe (let box u = exp 3 in u) :: Int -> Int
\x -> x * x * x * 1

Next steps

Convert NbE procedure to abstract machine through functional correspondence
Compare machine's behavior to known operational semantics of S4

NbE for sums and S4

By Maxim Koltsov

NbE for lambda calculus with sums and for modal logic

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