Relational Algebra with Fancy Types

Last time, I tried to give a primer of relations and relational algebra using the Haskell type type Rel a b = [(a,b)]. In this post we’re going to look at these ideas from a slightly different angle. Instead of encoding relations using value level sets, we’ll encode relations in the type system. The Algebra of Programming Agda repo and the papers quoted therein are very relevant, so if you’re comfortable wading into those waters, give them a look. You can find my repo for fiddling here

At this point, depending on what you’ve seen before, you’re either thinking “Yeah, sure. That’s a thing.” or you’re thinking “How and why the hell would you do such a ridiculous thing.”

Most of this post will be about how, so let’s address why first:

  1. Examining relations in this style illuminates some constructions that appear around the Haskell ecosystem, particularly some peculiar fellows in the profunctor package.
  2. More syntactic approach to relations allows discussion of larger/infinite domains. The finite enumerations of the previous post is nice for simplicity, but it seems you can’t get far that way.
  3. Mostly because we can – It’s a fun game. Maybe a useful one? TBD.

With that out of the way, let’s go on to how.

Translating Functions to Relation GADTs

We will be using some Haskell extensions in this post, at the very least GADTs and DataKinds. For an introduction to GADTs and DataKinds, check out this blog post. DataKinds is an extension that reflects every data constructor of data types to a type constructor. Because there are values True and False there are corresponding types created'True and 'False. GADTs is an extension of the type definition mechanism of standard Haskell. They allow you to declare refined types for the constructors of your data and they infer those refined type when you pattern match out of the data as well, such that the whole process is kind of information preserving.

We will use the GADT extension to define relational datatypes with the kind

a -> b -> *
. That way it has a slot a for the “input” and b for the “output” of the relation. What will goes in these type slots will be DataKind lifted types like 'True, not ordinary Haskell types like Int. This is a divergence from from the uses of similar kinds you see in Category, Profunctor, or Arrow. We’re doing a more typelevel flavored thing than you’ll see in those libraries. What we’re doing is clearly a close brother of the singleton approach to dependently typed programming in Haskell.

Some examples are in order for what I mean. Here are two simple boolean functions, not and and defined in ordinary Haskell functions, and their equivalent GADT relation data type.

not True = False
not False = True

data Not a b where
    NotTF :: Not 'True 'False
    NotFT :: Not 'False 'True

and True True = True
and False _ = False
and _ False = False

data And a b where
    AndTT :: And '( 'True, 'True) 'True
    AndFU :: And '( 'False, a) 'False
    AndUF :: And '( a, 'False) 'False

You can already start to see how mechanical the correspondence between the ordinary function definition and our new fancy relation type. A function is often defined via cases. Each case corresponds to a new constructor of the relation and each pattern that occurs in that case is the pattern that appears in the GADT. Multiple arguments to the relations are encoded by uncurrying everything by default.

Any function calls that occur on the right hand side of a function definition becomes fields in the constructor of our relation. This includes recursive calls and external function calls. Here are some examples with a Peano style natural number data type.

data Nat = S Nat | Z

plus Z x = x
plus (S x) y = S (plus x y)

data Plus a b where
    PZ :: Plus '( 'Z, a) a
    PS :: Plus '( a,b) c -> Plus '( 'S a, b) c 

We can also define things that aren’t functions. Relations are a larger class of things than functions are, which is part of their utility. Here is a “less than equal” relation LTE.

data LTE a b where
    LTERefl :: LTE n n
    LTESucc :: LTE n m -> LTE n ('S m)

You can show that elements are in a particular relation by finding a value of that relational type. Is ([4,7], 11) in the relation Plus? Yes, and I can show it with with the value PS (PS (PS (PS PZ))) :: Plus (4,7) 11 . This is very much the Curry-Howard correspondence. The type R a b corresponds to the proposition/question is (a,b) \in R .

The Fun Stuff : Relational Combinators

While you need to build some primitive relations using new data type definitions, others can be built using relational combinators. If you avoid defining too many primitive relations like the above and build them out of combinators, you expose a rich high level manipulation algebra. Otherwise you are stuck in the pattern matching dreck. We are traveling down the same road we did in the previous post, so look there for less confusing explanations of the relational underpinnings of these constructions, or better yet some of the references below.

Higher order relational operators take in a type parameters of kind

a -> b -> *
and produce new types of a similar kind. The types appearing in these combinators is the AST of our relational algebra language.

The first two combinators of interest is the composition operator and the identity relation. An element (a,c) is in R \cdot Q if there exists a b such that (a,b) \in R and (b,c) \in Q. The fairly direct translation of this into a type is

{- rcompose :: Rel b c -> Rel a b -> Rel a c  -}

data RCompose k1 k2 a c where
   RCompose :: k1 b c -> k2 a b -> RCompose k1 k2 a c

type k <<< k' = RCompose k k' 
type k >>> k' = RCompose k' k

The type of the composition is the same as that of Profunctor composition found in the profunctors package.

type RCompose = Procompose

Alongside a composition operator, it is a knee jerk to look for an identity relation and we do have one

data Id a b where
   Refl :: Id a a

-- monomorphic identity. Leave this out?
data IdBool a b where
  ReflTrue :: IdBool 'True 'True
  ReflFalse :: IdBool 'False 'False

This is also a familiar friend. The identity relation in this language is the Equality type.

-- identity function is the same as Equality
type Id a b = Id (a :~: b)

We can build an algebra for handling product and sum types by defining the appropriate relational combinators. These are very similar to the combinators in the Control.Arrow package.

-- Product types

data Fan k k' a b where
    Fan :: k a b -> k' a c -> Fan k k' a '(b,c)

type k &&& k' = Fan k k'

data Fst a b where
    Prj1 :: Fst '(a, b) a

data Snd a b where
    Prj2 :: Snd '(a, b) b

-- Sum type

data Split k k' a b where
    CaseLeft :: k a c -> Split k k' ('Left a) c
    CaseRight :: k' b c -> Split k k' ('Right b) c

type k ||| k' = Split k k'

data Inj1 a b where
    Inj1 :: Inj1 a ('Left a)
data Inj2 a b where
    Inj2 :: Inj2 a ('Right a)

-- some derived combinators
type Par f g = Fan (f <<< Fst) (g <<< Snd)
type Dup  = Fan Id Id
type Swap = Fan Snd Fst

The converse of relations is very interesting operation and is the point where relations really differ from functions. Inverting a function is tough. Conversing a relation always works. This data type has no analog in profunctor to my knowledge and probably shouldn't.

data RConverse k a b where
    RConverse :: k a b -> RConverse k b a
-- Shorter synonym
type RCon = RConverse

Relations do not have a notion of currying. The closest thing they have is

data Trans k a b where
    Trans :: k '(a,b) c -> Trans k a '(b,c)

Lattice Operators

For my purposes, lattices are descriptions of sets that trade away descriptive power for efficiency. So most operations you'd perform on sets have an analog in the lattice structure, but it isn't a perfect matching and you're forced into approximation. It is nice to have the way you perform these approximation be principled, so that you can know at the end of your analysis whether you've actually really shown anything or not about the actual sets in question.

? No. No... Yes? Oh. OH! IT IS!

The top relation holds all values. This is represented by making no conditions on the type parameters. They are completely phantom.

newtype Top a b = Top ()

Bottom is a relation with no inhabitants.

newtype Bottom a b = Bottom Void

The meet is basically the intersection of the relations, the join is basically the union.

newtype RMeet k k' a b = RMeet (k a b, k' a b)
type k /\ k' = RMeet k k'  

newtype RJoin k k' a b = RJoin (Either (k a b) (k' a b))
type k \/ k' = RJoin k k' 

A Lattice has an order on it. This order is given by relational inclusion. This is the same as the :-> combinator can be found in the profunctors package.

type (:->) p q = forall a b. p a b -> q a b
type RSub p q = p :-> q

Relational equality can be written as back and forth inclusion, a natural isomorphism between the relations. There is also an interesting indirect form.

data REq k k' = REq {to' :: k :-> k', from' :: k' :-> k }

Relational Division

If we consider the equation (r <<< p) :-> q with p and q given, in what sense is there a solution for r? By analogy, this looks rather like r*p = q, so we're asking a kind of division question. Well, unfortunately, this equation may not necessarily have a solution (neither do linear algebraic equations for that matter), but we can ask for the best under approximation instead. This is the operation of relational division. It also appears in the profunctor package as the right Kan Extension. You'll also find the universal property of the right division under the name curryRan and uncurryRan in that module.

newtype Ran p q a b = Ran { runRan :: forall x. p x a -> q x b }
type RDiv = Ran

One formulation of Galois connections can be found in the adjunctions file. Galois Connections are very slick, but I'm running out of steam, so let's leave that one for another day.

Properties and Proofs

We can prove many properties about these relational operations. Here a a random smattering that we showed using quickcheck last time.

prop_ridleft ::  (k <<< Id) :-> k
prop_ridleft (RCompose k IdRefl) = k

prop_ridright ::  (Id <<< k) :-> k
prop_ridright (RCompose IdRefl k) = k

prop_meet :: p /\ q :-> p
prop_meet (RMeet (p, q)) = p

prop_join :: p :-> p \/ q
prop_join p = RJoin (Left p)

meet_assoc :: RMeet k (RMeet k' k'') a b -> RMeet (RMeet k k') k'' a b
meet_assoc (RMeet (k, (RMeet (k',k'')))) = RMeet (RMeet (k,k'), k'')

prop_top :: k :-> Top
prop_top _ = top

prop_bottom :: Bottom :-> k
prop_bottom (Bottom x) = absurd x

bottom_compose :: REq (k <<< Bottom) Bottom
bottom_compose = REq (\(RCompose k (Bottom b)) -> absurd b) prop_bottom

data Iso a b = Iso {to :: a -> b, from :: b -> a}
type a <-> b = Iso a b

meet_universal :: (p ::-> RMeet k k') <-> (p ::-> k, p ::-> k')
meet_universal = Iso to from where
    to (RSub f) = (RSub $ \p -> case f p of RMeet (k,k') -> k  , RSub $ \p -> case f p of RMeet (k,k') -> k')
    from (RSub f,RSub g) = RSub $ \p -> RMeet (f p, g p) 

prop_con :: RCon (RCon k) :-> k
prop_con (RConverse (RConverse k)) = k

Odds and Ends

  • Recursion Schemes - Recursion schemes are a methodology to talk about recursion in a point free style and where the rubber meets the road in the algebra of programming. Here is an excellent series of articles about them. Here is a sample of how I think they go:
data MapMaybe k a b where
    MapJust :: k a b -> MapMaybe k ('Just a) ('Just b)
    MapNothing :: MapMaybe k 'Nothing 'Nothing

data Cata map k a b where
    Cata :: k fa a -> map (Cata map k) x fa  -> Cata map k ('Fix x) 
  • Higher Order Relations?
  • Examples of use. Check out the examples folder in the AoP Agda repo. These are probably translatable into Haskell.
  • Interfacing with Singletons. Singletonized functions are a specialized case or relations. Something like?
  •  newtype SFun a b = SFun (Sing a -> Sing b) 
  • A comment to help avoid confusion. What we've done here feels confusingly similar to profunctor, but it is in fact distinct I think. Profunctors are described as a categorical generalization of relations , but to be honest, I kind of don't get it. Despite many of our constructions appearing in the profunctor package, the profunctor typeclass itself appears to not play a role in our formulation. There just isn't a good way to dimap under our relations as written, unless you construct free profunctors. Converse at the least is a wrench in the works.
  • Star and graphs. Transition relations are a powerful methodology. A transition relation is in some respects the analog of a square matrix. We can iteratively compose it with itself.
-- Check out "term rewriting and all that"
-- This is also the reflection without remorse data type
-- TSequence
-- this is also a free instance of Category
data Star k a b where
    Done :: Star k a a
    Roll :: k b c -> Star k a b -> Star k a c

data KPlus k a b where
    PDone :: k a b -> KPlus k a b
    PRoll :: k b c -> KPlus k a b -> KPlus k a c

type SymClos k a b = RJoin k (RCon k) a b
type RefClos k a b = RJoin k Id a b
{- n-fold composition -}
-- similar to Fin.
-- This is also the Vec n is to list and this is to reflection without remorse. Kind of interesting
data NFold n k a b where
    One :: k a b -> NFold ('S n) k a b
    More :: k b c -> NFold n k a b -> NFold ('S n) k a b


CAV 2019 Notes: Probably Nothin Interestin’ for You. A bit of noodling with Liquid Haskell

I went to the opening workshops of CAV 2019 in New York this year (on my own dime mind you!) after getting back from joining Ben on the long trail for a bit. The tutorials on Rosette and Liquid Haskell were really fun. Very interesting technology. Also got some good ramen and mochi pudding, so it’s all good. Big Gay Ice Cream was dece.

Day 1 workshops

Calin Belta a new book. Control of Temporal logic systems. Automata. Optimized. Partition space into abstraction. Bisimulation

Control Lyapunov Function (CLF) – guarantees you are going where you want to go

Control Barrier Function – Somehow controls regions you don’t want to go to.

Lyapunov function based trajectory optimization. You somehow have (Ames 2014) Is this it?

Differential flatness , input output linearization

Sadradiini worked there.

Temproal logic with

Rise of Temporal Logic

Linear Temporal Logic vs CTL

Fixpoint logic,

Buchi automata – visit accepting state infinite times

equivalency to first order logic

monadic logic, propositions only take 1 agrument. Decidable. Lowenheim. Quantifier elimination. Bounded Mondel property

Languages: ForSpec, SVA, LDL, PSL, Sugar

Monadic second order logic (MSO).

method of tableau


Polytopic regions. Can push forward the dynmaics around a trajectory and the polytope that you lie in. RRT/LQR polytopic tree. pick random poitn. Run.

Evauating branching heuristics

branch and prune icp. dreal.

branch and prune. Take set. Propagate constraints until none fire.

branching heuristics on variables

largest first, smearing, lookahead. Try different options see who has the most pruning. Non clear that helped that muhc

QF_NRA. dreal benchmarks. flyspeck, control, robotics, SMT-lib


commands: saolver adied programming

verify – find an input on which the assertions fail. exists x. not safe

debug – Minimal unsat core if you give an unsat query. x=42/\ safe(s,P(x))$ we know thia is unsat because of previous step

solve – exists v si.t safe(v)

synthesis – exists e forall x safe(x,P(x))

define-symbolic, assert, verify, debug, solve, sythesis

Rosette. Alloy is also connected to her. Z Method. Is related to relational logic?

Building solver aided programming tool.

symbolic compiler. reduce program all possible paths to a constraint

Cling – symbolic execution engine for llvm

implement intepreter in rosette

Symbolic virtual machine

layering of languages. DSL. library (shallow) embedding. interpreter (deep) embedding.

deep embedding for sythesis.

I can extract coq to rosette?

how does it work?

reverse and filter keeping only positive queries.

symbolic execution vs bounded model checking

symbolic checks every possible branch of the program. Cost is expoentntial


type driven state merging. Merge instances of primitiv types. (like BMC), value types structurally ()

instance Merge Int, Bool, Real — collect up SMT context

vs. Traversable f => Merge (f c) – do using Traversable

symbolic union a set of guarded values with diskoint guard.

merging union. at most one of any shape. bounded by number of possible shapes.

puts some branching in rosette and some branch (on primitives) in SMT.

symbolic propfiling. Repair the encdoing.

tools people have built.

veify radiation

strategy generation. That’s interesting. builds good rewrite rules.


certikso komodo keystone. fintie programss

IS rosette going to be useful for my work? coooooould be

Liquid Haskell

{-# LANGUAGE GADTs, DataKinds, PolyKinds #-}
{-@ LIQUID "--reflection" @-}
{-@ LIQUID "--short-names" @-}
{-@ LIQUID "--ple" @-}

{-@ type TRUE = {v: Bool | v = True }@-}
{-@  type NonEmpty a = {v : [a] | len v > 0}  @-}
{-@ head :: NonEmpty a -> a @-}
head (x : _) = x

{-@ measure f :: Int -> Int @-}
f x = 2 * x

{-@ true :: TRUE @-}
true = True

-- impl x y = x ==> y
-- {-@ congruence :: Int -> Int -> TRUE @-}
-- congruence x y = (x == y) ==> (f x == f y)
-- {-@ (==>) :: {x : Bool |} -> {y : Bool |} -> {z : Bool | z = x ==> y} @-}
-- x ==> y = (not x) || y

-- aws automated reaosning group
nullaway uber. 
sadowski google static anaylis
infer faecobook

give programmer early 

refinement types

why and how to use
how to implement refinement types

mostly like floyd hoare logic

types + predicates = refinement type

t := {x:b | p} erefinemed b
      x : t -> t -- refined function type
linear arithemtic

congruence axioms

emit a bunch of verification conditions (VC)
p1 => p2 => p3 ... 

SMT can tell if VC is alwasy true


{-@ type Zero = {v : Int | v == 0} @-}
{-@ zero :: Zero @-}
zero = (0 :: Int) -- why?

{-@ type NAT = {v: Int | v >= 0} @-}
{-@ nats :: [Nat]  @-}
nats = [0,1,1,1] :: [Int]



in an environemnt Gamma, t1 is a subtype of t2
x are vairables are in scope.
/\x |- {v | q} <= {v | r}

True  => 


{-@ plus :: x : Int -> y : Int -> {v : Int | v = x + y} @-}
plus :: Int -> Int -> Int
plus x y = x + y

-- measure. uninterpeteyd function called measure.
-- {-@ measure vlen :: Vector a -> Int @-}

-- {-@ at :: vec :Vector a -> {i : Nat | i < vlen vec} ->  @-}

-- {-@  size :: vec : Vector a -> {n : Nat | n == vlen vec} @-}

horn contrints infer 

Collections and Higher and order fucntions

reduce :: (a -> b -> a) -> a -> [b] -> a
type a is an invaraint that holds on initial acc and indictively on f
Huh. That is true.

-- Huh. I could prove sum formulas this way.

sum'' = vec = let is = range 0 (size vec)
              add = \tot i -> ot + at vec i

              properties of data structures

size of of a list
allow uninterpetyed functions inside refinements

{ measre length}

LISTNE a = {v : [a] | 0 < length v}

measure yields refined constructions

[] :: {v : [a] | legnth v = 0}


              Q: measure?
              Q : environment?
              Q : where is standard libraru?
              no foralls anywhere? All in decidable fragment

              p : ([a],[a]) | (fst p) + (snd p) == lenght xs
              measures fst and snd


              impossible :: {v : String | false} -> a

imperative language


/inlcude folder has prelude

basic values
{v : Int | lo <= v && v < hi }

invaraint properties of structures

encoe invraints in constructor

data OrdPair = OP {opX :: Int, opY :: Int}

{-@ data OrdPair = OP {opX :: Int, opY :: {v : Int | opX < v}@-}

class Liquid Int
class Liquid Bool
class Liquid ... Real?

{-@   @-}

Liquid Relations?
{  }

data OList 
{-@data OList a LB = Empt | :< {oHd :: {v : a | LB = v}   oTl :: OList a oHd }   @-}

{-@   {oHd :: a, oTl :: OList {v : a | oHd < v}   }@-}
data MyList a where
    Nil :: MyList a
    {-@ Cons :: v : a -> MyList {x : a | x < v} ->  MyList a @-}
    Cons :: a -> MyList a -> MyList a

test :: MyList Int
test = Cons 2 (Cons 1 Nil)


abstracting the invaraint from the data structure
parametrzie by relations

data [a]<rel :: a -> a -> Bool> where
    = []
    | (:) {hd :: }

    rel != is unique list

{\x y -> x >= y}
type level lambdas!? .... uh.... maybe.

reflecting singletons into liquid?

termination metrics

/ [length xs + len ys] -- merge sort

{-@ Half a s = }@-}

Oncey ou have temrination proofs you have proofs of correctness

Propositions as Types

Plus commutes is trivial
{n = n + n}


{-@ easyProof :: {True} @-}
easyProof = ()

-- hot damn. I mean this is in it's legerdomain. But prettttty sweet.
{-@ commute :: x : Int -> y : Int -> {x + y = y + x} @-}
commute :: Int -> Int -> ()
commute x y = ()

{-@ reflect mysum @-}
{-@ mysum :: Nat -> Nat @-}
mysum :: Int -> Int
mysum 0 = 0 -- if n <= 0 then 0 else  2 * n + (mysum (n - 1))
mysum n = 2 * n + (mysum (n - 1))

-- what is going on here? why do I need _?
{-@ mysumpf :: _ -> {mysum 0 = 0 } @-}
-- mysumpf :: Proof
mysumpf _ = let x = mysum 0 in x

{-@ mysumpf' :: {mysum 3 = 12 } @-}
-- mysumpf :: Proof
mysumpf' = ()

{-@ reflect fastsum @-}
{-@ fastsum :: Nat -> Nat @-}
fastsum :: Int -> Int
fastsum n = n * (n + 1)

type Proof = ()
{-@ pfsum :: x : Nat -> {fastsum x = mysum x} @-}
pfsum :: Int -> Proof
pfsum 0 = () -- let _ = fastsum 0 in let _ = mysum 0 in ()
pfsum n = pfsum (n-1)  

{-@ pfsum :: x : Nat -> {fastsum x = mysum x} @-}
pfsum :: Int -> Proof
pfsum 0 = () -- let _ = fastsum 0 in let _ = mysum 0 in ()
pfsum n = pfsum (n-1)  



reflect takes the prcondition of sum and dumps it as the poscondition

sum3 _ = let s0 =sum 0
             s1  = sum 1
             s2  = sum 3 -- all are going to be in scope.
z3 will connect the dots.

using proof combinatos from Proof Combinators
long chains of claculations

reflection of singletons

data SS s where
    {-@ SZero :: {v : Int | v = 0} -> SS 'Zero @-} 
    SZero :: Int -> SS 'Zero 
    {-@ SZero :: {v : Int | v = 0} -> SS 'S a @-} 
    SZero :: Int -> SS 'Zero 

    proof by induction
    sum n = n * (n + 1)/2

    2 * sum n  = n * (n + 1)

point free liquid types

(.) :: (a -> b) -> (a -> )
? Can I abstract over predicates like this?
({v:a | p} -> {s:}) -> 


cauchy schwartz


data V2 a = V2 a a 

{-@ reflect dot @-} 
dot (V2 x y) (V2 x' y') = x * x' + y * y'

{-@ reflect vplus @-}
vplus (V2 x y) (V2 x' y') = V2 (x + x') (y + y')

{-@ reflect smul @-}
smul s (V2 x' y') = V2 (s * x') (s * y')
{-@ cauchy :: x : V2 Int -> y : V2 Int -> {(dot x y) * (dot x y) <= (dot x x) * (dot y y) } @-}
cauchy :: V2 Int -> V2 Int -> Proof
cauchy x y = let q = dotpos (vplus x y) in let r = dotpos (vplus x (smul (-1 :: Int) y)) in (\_ _ -> ()) q r

-- {-@ square :: Int -> Nat @-} -- basiclly the same thing
{-@ reflect square @-}
square :: Int -> Int
square x = x * x

{-@ sqpos :: x: Int -> {square x >= 0} @-}
sqpos :: Int -> ()
sqpos x = ()

{-@ dotpos :: x: V2 Int -> {dot x x >= 0} @-}
dotpos :: V2 Int -> ()
dotpos x = ()

{-@ dotsym :: x: V2 Int -> y : V2 Int -> {dot x y = dot y x} @-}
dotsym :: V2 Int -> V2 Int ->  ()
dotsym x y = ()

{-@ vpluscomm :: x: V2 Int -> y : V2 Int -> {vplus x y = vplus y x} @-}
vpluscomm :: V2 Int -> V2 Int ->  ()
vpluscomm x y = ()

{-@ dotlin :: x: V2 Int -> y : V2 Int -> z : V2 Int -> {dot (vplus x y) z = dot x z + dot y z} @-}
dotlin :: V2 Int -> V2 Int ->  V2 Int ->  ()
dotlin x y z = ()


What else is interesting to prove?

verify stuff about ODEs?

fold [1 .. t] where t = 10

could give little spiel about how dynamical systems are like imperative programming

get some rationals.

p a b
a -> b are refined functions

I should learn how to abstract over typeclasses. Verified typeclasses?

SMT has built in rationals prob?

data Rat = Rat Int Int 

{-@ reflect rplus @-}
rplus :: Rat -> Rat -> Rat
rplus (Rat x y) (Rat x' y') = Rat (x*y' + x'*y) (y * y')

{-@ reflect rmul @-}
rmul :: Rat -> Rat -> Rat
rmul (Rat x y) (Rat x' y') = Rat (x*x') (y * y')

data Nat' = S Nat' | Z

{-@ measure nlen @-}
{-@ nlen :: Nat' -> Nat @-}
nlen :: Nat' -> Int
nlen Z = 0
nlen (S x) = 1 + (nlen x) 
-- failing?
-- crash: SMTLIB2 respSat = Error "line 31 column 169: unknown sort 'Main.SNat'"
data SNat a where
    SZ :: SNat 'Z
    SS :: SNat x -> SNat ('S x)

{-@ reflect conv @-}
{-@ conv :: x : Nat -> {v : Nat' | nlen v = x} @-}
conv :: Int -> Nat'
conv 0 = Z
conv x = S (conv (x-1))

-- It's an isomorphism

{-@ pfconv :: x : Nat -> {nlen (conv x) = x} @-}
pfconv :: Int -> Proof
pfconv 0 = ()
pfconv x = pfconv (x - 1)

{-@ pfconv' :: x : Nat' -> {conv (nlen x) = x} @-}
pfconv' :: Nat' -> Proof
pfconv' Z = ()
pfconv' (S x) = pfconv' x

{-@ reflect plus' @-}
plus' :: Nat' -> Nat' -> Nat'
plus' Z x = x
plus' (S x) y = S (plus' x y)

{-@ plusz' :: x : Nat' -> {plus' x Z = plus' Z x}  @-}
plusz' :: Nat' -> Proof
plusz' Z = ()
plusz' (S x) = plusz' x

{-@ pluscomm' :: x : Nat' -> y : Nat' -> {plus' x y = plus' y x} / [nlen x, nlen y] @-}
pluscomm' :: Nat' -> Nat' -> Proof
pluscomm' Z y = plusz' y
pluscomm' (S x) (S y) = const (pluscomm' (S x) y) $ const (pluscomm' x (S y)) $  pluscomm' x y
    -- const () $ const (plus' (S x) (S y)) $ const (plus' x (S y)) (plus' x y)
    -- const (pluscomm' (S x) y) $ const (pluscomm' x (S y)) $  pluscomm' x y
 -- flip const is proof combinator .==   
    {-let q = pluscomm' x (S y) in
                        let w = pluscomm' (S x) y in 
                        let r = pluscomm' x y in (\b n m -> ()) q w r -- ? Was this necessary? -}
pluscomm' x Z = plusz' x

-- {-@ data Iso = @-}
data Iso a b = Iso { to :: a -> b, from :: b -> a, p1 :: Proof, p2 :: Proof}


We also have type level lambdas.
refinement polymorphism

LH is somewhat like singletons in the sense there is a manual reflection step.
In singletons the manual reflection is in the Sing type
in LH it is kind of all over the place. (+) has a type. Where is it defined?

How does it know that the Haskell function + is the same as the SMT solver function?

Coq and Agda and Idris type checking is powered quite a bit by an internal unification engine
explicit annotation may lessen the burden somewhat

SMT solvers as a unification engine
structure unification vs uninterpeted functions.
f a ~ Int is not a valid Haskell constraint. Maybe with the unmatchable arrow it is?
In a funny sense, there is a difference between  Just and (+ 1). 
One being a constructor means we can match out of it
Just :: a ->> b
(+ 1) :: Int -> Int


-- test' :: (f a ~ Int) => ()
-- test' = ()

Liquid Haskell – What is?

another thing we could do is galois connections between refinements. Pos, Zero, Neg <-> Int

Liquid Haskell uses SMT solvers to resolve it’s type checking requirements.

Agda et al also work very much via unification. Unification is a broad term but it’s true.

It also has a horn clause solver for inference. Every language needs some kind of inference or you’d go insane. Also it is piggybacking on haskell

It’s not as magical as I thought? Like seeing the magicians trick. It really does understand haskell code. Like it isn’t interpretting it. When it knows facts about how (+) works, that is because the refined type was put in by hand in the prelude connecting it to SMT facts. What is imported by liquid haskell?

The typing environment is clutch. You need to realize what variables are in scope and what their types are, because that is all the SMT can use to push through type checking requirements.

Installing the stack build worked for me. It takes a while . I couldn’t get cabal install to work, because I am not l33t.

Uninterpeted functions. Unmatchability?

It wouldn’t be haskell without a bunch of compiler directives. It is somewhat difficult to find in a single cohesive place what all the syntax, and directives are from liquid haskell. Poking around it best.

  • ple
  • reflection
  • no-termination
  • higherorder – what is this? course notes some of software foundations niki vazou’s pubs. Check out refinement reflection draft work? Shows stuff about typeclasses. This is a haskell 2017 paper though intro to liquid haskell. Interesting to a see a different author’s take presentations. They are fairly similar to one another.

Liquid haskell gives us the ability to put types on stuff that wasn’t possible before.

Linearity :: f :: {a -> b | f (s ^* a) == s ^* (f a) }

Pullback. {(a,b) | f a == g b}


Many things in categoruy theory rely on the exists unique. Do we have functiona extensionality in Liquid haskell?

product : {(a,b) | f q = x, g q = y, => }

Pushing the boundaries on what liquid haskell can do sounds fun.

Equalizer. The eqaulizer seems prominents in sheaves. Pre-sheaves are basically functors. Sheaves require extra conditions. Restriction maps have to work? Open covers seem important

type Equalizer f g a b = {(e :: a , eq :: a -> b) | f (eq e) = g (eq e) }

I think both the type a and eq are special. e is like an explcit parametrization.

type Eq f g a = {e :: a | f e = g e} I think this is more in the spirit. Use f and g both as measures.

presheaf is functor. But then sheaf is functor that

(a, Eq (F a) (G a)). typelevel equalizer? All types a that F and G agree on.

Records are sheaves – Jon Sterling. Records have subtyping. This gives you a toplogy feeling thing.

What about purescript records?

{foo | a} {bar | a} -> intersection = {foo bar | b} can inhabit either

union is
or do you want closed records? union is union of fields. intersection is intersection of fields.

In this case a cover would be a set of records with possibly overlapping fields whose combined labels cover the whle space we want to talk about. consistency condition of sheaf/equalizer is that overlapping records fields have to match. I guess { = } ?There is a way to combine all the stuff up. This is exactly what Ghrist was getting at with tables. Tables with shared columns.

data R1 = R1 {foo :: Int, bar :: Int}

{ (r1 :: R1, r2 :: R2) | (foo r1) = (foo r2) } — we manitain duplicates across records.

{. }

if you have a “cover” {foo bar |} {bar fred} {gary larry} whose in

Sheaves. As a model of concurrency? Gaguen paper.

sheaves as constraint satisfcation? sheafifcation. Constraint solving as a way of fusing the local constraints to be globally consistent.

sheaves as records

sheaves as data fusion

Escardo. Compact data types are those finitely searchable

Continuous funcitons are ~computable? Productive?

typed recursion theory toplogy

typed computatabiltity theory

Topological notions in computation. Dictionary of terms realted decidable, searchable, semi decidablee

Through NDArray overloading, a significant fragment of numpy code is probably verifiable.

Start with functional arithmetic programs.

Need to inspect function annotations to know how to build input type.

@verify() tag

Use (Writer a) style monad.

If statements are branching. We are again approaching inspecting functions via probing. But what if we lazily probe. At every __bool__ point, we run a z3 program to determine if there is an avaiable bool left possible (we don’t need to inspect dead code regions. Also would be cool to mention it is a dead region). Curious. We’re reflecting via Z3.

Loops present a problem. Fixed loops are fine. but what about loops that depend on the execution? for i in range(n). I guess again we can hack it…? Maybe. range only takes an integer. we don’t have overload access.

Maybe we need to go into a full eval loop. utterly deconstructing the function and evaluating it statelemnt by statement.

(compare :: a -> a -> Comparison). We could select a choice based on if there is a new one avaialable. Requires access to external store. We lose the thread. How can we know a choice was made? How can we know what the choice was? Did it ask var1 or var2? We can probably do it in python via access to a global store. But in haskell?

while loops take invariant annotations.

It would be cool to have a program that takes

pre conditions. Post ocnditions, but then also a Parameter keyword to declare const variables as deriveable. exists parameter. forall x precondition x => post condition.

Parameter could be of a type to take a dsl of reasonable computations. Perhaps with complexity predicates. and then interpretting the parameter defines the computation.

Or simpler case is parameter is an integer. a magic number.

@pre(lambda x: None)
@post(lambda r: r >= 0)
def square(x):
	return x**2

@verify(pre, post) # Easier. because now we can also verify the individual function. Call Z3 at function definition time.
def pre(f,cond):
   return fnew
   def fnew(x):
	   		if(x == VerificationEnv):
	   			newenv = x.copy
	            newVar = Z3.variable()
	            newenv.add(newVar == f(x.var))

	        	return f(x)

def post(f, cond):
   def fnew(x):
    if x == VerifcationEnv:
        Z3.findmodel(not cond(x.var), x.env)
        if can't find one:
           we're good.
           return x
           assert(False, model, function name, function postcondition code.)

#overloading assignment. isn't a problem.
class VerifyArray():
   #numpy Z3 shim.

#termination requires measure decreasing at every recursive call.
#arbaitrary loops? How to deal with those?
#maybe a hierarchy of envs to make it easier on z3. Like it doesn't need to give the whole program history if the local use is good enough.

class VerificationEnv():
    self.var = []
    self.pre = [] = []

Proving some Inductive Facts about Lists using Z3 python

Z3 is an SMT solver which has a good rep. Here are some excellent tutorials.

SMT stands for satisfiability modulo theories. The exact nature of power of these kinds of solvers has been and is still hazy to me. I have known for a long time that they can slam sudoku or picross or other puzzles, but what about more infinite or logic looking things? I think I may always be hazy, as one can learn and discover more and more encoding tricks to get problems and proofs that you previously thought weren’t solvable into the system. It’s very similar to learning how to encode to linear programming solvers in that respect.

SMT solvers combine a general smart search procedure with a ton of specialized solvers for particular domains, like linear algebra, polynomials, linear inequalities and more.

The search procedure goes by the name DPLL(T). It is an adjustment of the procedure of SAT solvers, which are very powerful and fast. SAT solvers find an assignment of boolean variables in order to make a complicated boolean expression true, or to eventually find that it can never be made true. SAT solvers work on the principle of guessing and deducing. If a OR b needs to be true and we already know a is false, we can deduce b must be true. When the deduction process stalls, we just guess and then backtrack if it doesn’t work out. This is the same process you use manually in solving Sudoku.

The modern era of SAT solvers came when a couple new tricks vastly extended their power. In particular Conflict Driven Clause Learning (CDCL), where when the solver finds itself in a dead end, it figures out the choices it made that put it in the dead end and adds a clause to its formula so that it will never make those choices again.

SMT works by now having the boolean variables of the SAT solver contain inner structure, like the boolean p actually represents the fact x + y < 5. During the search process it can take the pile of booleans that have been set to true and ask a solver (maybe a linear programming solver in this case) whether those facts can all be made true in the underlying theory. This is an extra spice on top of the SAT search.

Something that wasn’t apparent to me at first is how important the theory of uninterpreted formulas is to SMT solvers. It really is their bread and butter. This theory is basically solved by unification, which is the fairly intuitive process of finding assignments to variables to make a set of equations true. If I ask how to make fred(x,4) = fred(7,y), obviously the answer is y=4, x=7. That is unification. Unification is a completely syntax driven way to deduce facts. It starts to give you something quite similar to first order logic.

I was also under the impression that quantifiers \forall, \exists were available but heavily frowned upon. I don’t think this is quite true. I think they are sort of a part of the entire point of the SMT solver now, although yes, they are rather flaky. There are a number of methods to deal with the quantifier, but one common one is to look for a pattern or parts of the subformula, and instantiate a new set of free variables for all of the quantified ones and add the theorem every time the patterns match. This is called E-matching.

Here are a couple tutorials on proving inductive facts in SMT solvers. You kind of have to hold their hand a bit.

SMT solvers queries usually have the flavor of finding something, in this case a counterexample. The idea is that you try to ask for the first counterexample where induction failed. Assuming that proposition P was true for (n-1), find n such that P is not true. If you can’t find it, then the induction has gone through.

And here is some code where I believe I’m showing that some simple list functions like reverse, append, and length have some simple properties like \forall t. rev (rev(t)) == t .

from z3 import *

# Very useful reference

f = Function('f', IntSort(), IntSort())
s = Solver()
#s.add(f(True) == False, f(False) == True)
x = Int('x')
s.add(ForAll([x], f(x) >= x)) #> and < do not seem to be returning

# Rolling my own list data type
# Z3 has a built in which will probably be better?
s = Solver()

L = Datatype("MyList")
L.declare("Cons", ("hd", IntSort()), ("tail", L))
L = L.create()

t = Const('t', L)
u = Const('u', L)

y = Int('y')

rev = Function('reverse', L, L)
app = Function('append', L, L, L)
leng = Function('length', L, IntSort())

#defining my functions. Micro Haskell, BABY
s.add( leng(L.Nil) == 0 )
s.add(  ForAll([u,y],  leng(L.Cons(y,u)) == 1 + leng(u))) #  patterns = [leng(L.Cons(y,u))] 

s.add(  ForAll([u], app(L.Nil, u) == u)) 
s.add(  ForAll([t, u, y] , app(L.Cons(y,t), u)  ==  L.Cons(y, app(t, u ))))

s.add( rev(L.Nil) == L.Nil)
s.add(  ForAll([y,t],  rev(L.Cons(y,t)) == app(rev(t), L.Cons(y, L.Nil))))

print("proving leng(t) >= 0")
#s.add( Or(And(t == L.Cons(y,u),  leng(u) >= 0 ), t == L.Nil))
s.add(  Not(leng(t) >= 0 ))
s.add( Implies(t == L.Cons(L.hd(t), L.tail(t)),  leng(L.tail(t)) >= 0 ))

#s.add( leng(app(L.Nil, u)) == leng(u) )

print("prove length is preserved under app.")
s.add( leng(app(t,u)) != leng(t) + leng(u))
s.add( Implies(t == L.Cons(L.hd(t), L.tail(t)),   leng(app(L.tail(t),u)) == leng(L.tail(t)) + leng(u)  ))


print("reverse preserves length")
#Lemma Could place in clause with the above proof of this theorem
s.add( ForAll([t,u], leng(app(t,u)) == leng(t) + leng(u)   )) #subgoal needed
s.add( leng(rev(t)) != leng(t))
s.add( Implies(t == L.Cons(L.hd(t), L.tail(t)),   leng(rev(L.tail(t))) == leng(L.tail(t)) ))


print("reverse reverse = id")
s.add( ForAll( [t,u], rev(app(t,u)) ==  app(rev(u), rev(t)) ) ) #subgoal needed
s.add( rev(rev(t)) != t )
s.add( Implies(t == L.Cons(L.hd(t), L.tail(t)),   rev(rev(L.tail(t))) == L.tail(t) ))


#s.add(t != L.Nil ) 
#s.add( ForAll([t], rev(L.Cons(y,t)) ==   ) , rev(L.Nil) == L.Nil)
#s.add( leng(L.Cons()) == 0 )
#s.add(  ForAll(y,  leng(L.Cons(y,L.Nil)) == 1 + leng(L.Nil)))
#s.add(  Not( leng(app(t,u))  == leng(t) + leng(u) )) 

# prove length of app + nil is same
#s.add( leng(app(t,L.Nil)) != leng(t))
#s.add( Implies(t == L.Cons(L.hd(t), L.tail(t)),   leng(app(L.tail(t),L.Nil)) == leng(L.tail(t))))

#s.add( app(L.Nil,L.Nil) == L.Nil)
#s.add( app(t,L.Nil) != t)
#s.add( Implies(t == L.Cons(L.hd(t), L.tail(t)),   app(L.tail(t),L.Nil) == L.tail(t)))

#s.add(Or(t == L.Nil , And( t == L.Cons(y, u),   app(u,L.Nil) == u)  ))
#s.add( Implies(t == L.Nil, leng(app(L.Nil, u)) == leng(u) ))
# all of these don't work
#s.add( Implies(u == L.tail(t),  leng(u) >= 0 ))
#s.add( ForAll(u, Implies(t == L.Cons(y,u),  leng(u) >= 0 )))

def induction(freevar, f, construtors?):
    x = Int('x')
    return And(Not(f(x)), 

# controller synthesis
pos = Array(Real, 10) 
for t in range(10):
    s.add(pos[t] == pos[t] +  v * dt)
    s.add(v[t] == v[t] + f(pos[t]) * dt)
    s.add(f(pos[t] <= 1)) # constraints on force
s.add(ForAll([init], Implies(And(init <= 1, pos[0] == init), pos[9] <= 0.1) ))
s.add(Forall[x], f(x) == a * x) # parametrizing. 
#s.set("produce-proofs", True
#s.set("mbqi", False)