ghc-9.12.1: GHC/Tc/Solver/InertSet.hs
{-# LANGUAGE DerivingStrategies #-}
{-# LANGUAGE RankNTypes #-}
{-# LANGUAGE TypeApplications #-}
{-# OPTIONS_GHC -Wno-incomplete-uni-patterns #-}
module GHC.Tc.Solver.InertSet (
-- * The work list
WorkList(..), isEmptyWorkList, emptyWorkList,
extendWorkListNonEq, extendWorkListCt,
extendWorkListCts, extendWorkListCtList,
extendWorkListEq, extendWorkListEqs,
appendWorkList, extendWorkListImplic,
workListSize,
selectWorkItem,
-- * The inert set
InertSet(..),
InertCans(..),
emptyInert,
noMatchableGivenDicts,
noGivenNewtypeReprEqs, updGivenEqs,
mightEqualLater,
prohibitedSuperClassSolve,
-- * Inert equalities
InertEqs,
foldTyEqs, delEq, findEq,
partitionInertEqs, partitionFunEqs,
foldFunEqs, addEqToCans,
-- * Inert Dicts
updDicts, delDict, addDict, filterDicts, partitionDicts,
addSolvedDict,
-- * Inert Irreds
InertIrreds, delIrred, addIrreds, addIrred, foldIrreds,
findMatchingIrreds, updIrreds, addIrredToCans,
-- * Kick-out
KickOutSpec(..), kickOutRewritableLHS,
-- * Cycle breaker vars
CycleBreakerVarStack,
pushCycleBreakerVarStack,
addCycleBreakerBindings,
forAllCycleBreakerBindings_,
-- * Solving one from another
InteractResult(..), solveOneFromTheOther
) where
import GHC.Prelude
import GHC.Tc.Types.Constraint
import GHC.Tc.Types.Origin
import GHC.Tc.Types.CtLoc( CtLoc, ctLocOrigin, ctLocSpan, ctLocLevel )
import GHC.Tc.Solver.Types
import GHC.Tc.Utils.TcType
import GHC.Types.Var
import GHC.Types.Var.Env
import GHC.Types.Var.Set
import GHC.Types.Basic( SwapFlag(..) )
import GHC.Core.Reduction
import GHC.Core.Predicate
import GHC.Core.TyCo.FVs
import qualified GHC.Core.TyCo.Rep as Rep
import GHC.Core.Class( Class )
import GHC.Core.TyCon
import GHC.Core.Class( classTyCon )
import GHC.Core.Unify
import GHC.Utils.Misc ( partitionWith )
import GHC.Utils.Outputable
import GHC.Utils.Panic
import GHC.Data.Maybe
import GHC.Data.Bag
import Data.List.NonEmpty ( NonEmpty(..), (<|) )
import qualified Data.List.NonEmpty as NE
import Data.Function ( on )
import Control.Monad ( forM_ )
{-
************************************************************************
* *
* Worklists *
* Canonical and non-canonical constraints that the simplifier has to *
* work on. Including their simplification depths. *
* *
* *
************************************************************************
Note [WorkList priorities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
A WorkList contains canonical and non-canonical items (of all flavours).
Notice that each Ct now has a simplification depth. We may
consider using this depth for prioritization as well in the future.
As a simple form of priority queue, our worklist separates out
* equalities (wl_eqs); see Note [Prioritise equalities]
* all the rest (wl_rest)
Note [Prioritise equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
It's very important to process equalities over class constraints:
* (Efficiency) The general reason to do so is that if we process a
class constraint first, we may end up putting it into the inert set
and then kicking it out later. That's extra work compared to just
doing the equality first.
* (Avoiding fundep iteration) As #14723 showed, it's possible to
get non-termination if we
- Emit the fundep equalities for a class constraint,
generating some fresh unification variables.
- That leads to some unification
- Which kicks out the class constraint
- Which isn't solved (because there are still some more
equalities in the work-list), but generates yet more fundeps
Solution: prioritise equalities over class constraints
* (Class equalities) We need to prioritise equalities even if they
are hidden inside a class constraint; see Note [Prioritise class equalities]
* (Kick-out) We want to apply this priority scheme to kicked-out
constraints too (see the call to extendWorkListCt in kick_out_rewritable)
E.g. a CIrredCan can be a hetero-kinded (t1 ~ t2), which may become
homo-kinded when kicked out, and hence we want to prioritise it.
Further refinements:
* Among the equalities we prioritise ones with an empty rewriter set;
see Note [Wanteds rewrite Wanteds] in GHC.Tc.Types.Constraint, wrinkle (W1).
* Among equalities with an empty rewriter set, we prioritise nominal equalities.
* They have more rewriting power, so doing them first is better.
* Prioritising them ameliorates the incompleteness of newtype
solving: see (Ex2) in Note [Decomposing newtype equalities] in
GHC.Tc.Solver.Equality.
Note [Prioritise class equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We prioritise equalities in the solver (see selectWorkItem). But class
constraints like (a ~ b) and (a ~~ b) are actually equalities too;
see Note [The equality types story] in GHC.Builtin.Types.Prim.
Failing to prioritise these is inefficient (more kick-outs etc).
But, worse, it can prevent us spotting a "recursive knot" among
Wanted constraints. See comment:10 of #12734 for a worked-out
example.
So we arrange to put these particular class constraints in the wl_eqs.
NB: since we do not currently apply the substitution to the
inert_solved_dicts, the knot-tying still seems a bit fragile.
But this makes it better.
Note [Residual implications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The wl_implics in the WorkList are the residual implication
constraints that are generated while solving or canonicalising the
current worklist. Specifically, when canonicalising
(forall a. t1 ~ forall a. t2)
from which we get the implication
(forall a. t1 ~ t2)
See GHC.Tc.Solver.Monad.deferTcSForAllEq
-}
-- See Note [WorkList priorities]
data WorkList
= WL { wl_eqs_N :: [Ct] -- /Nominal/ equalities (s ~#N t), (s ~ t), (s ~~ t)
-- with empty rewriter set
, wl_eqs_X :: [Ct] -- CEqCan, CDictCan, CIrredCan
-- with empty rewriter set
-- All other equalities: contains both equality constraints and
-- their class-level variants (a~b) and (a~~b);
-- See Note [Prioritise equalities]
-- See Note [Prioritise class equalities]
, wl_rw_eqs :: [Ct] -- Like wl_eqs, but ones that have a non-empty
-- rewriter set; or, more precisely, did when
-- added to the WorkList
-- We prioritise wl_eqs over wl_rw_eqs;
-- see Note [Prioritise Wanteds with empty RewriterSet]
-- in GHC.Tc.Types.Constraint for more details.
, wl_rest :: [Ct]
, wl_implics :: Bag Implication -- See Note [Residual implications]
}
appendWorkList :: WorkList -> WorkList -> WorkList
appendWorkList
(WL { wl_eqs_N = eqs1_N, wl_eqs_X = eqs1_X, wl_rw_eqs = rw_eqs1
, wl_rest = rest1, wl_implics = implics1 })
(WL { wl_eqs_N = eqs2_N, wl_eqs_X = eqs2_X, wl_rw_eqs = rw_eqs2
, wl_rest = rest2, wl_implics = implics2 })
= WL { wl_eqs_N = eqs1_N ++ eqs2_N
, wl_eqs_X = eqs1_X ++ eqs2_X
, wl_rw_eqs = rw_eqs1 ++ rw_eqs2
, wl_rest = rest1 ++ rest2
, wl_implics = implics1 `unionBags` implics2 }
workListSize :: WorkList -> Int
workListSize (WL { wl_eqs_N = eqs_N, wl_eqs_X = eqs_X, wl_rw_eqs = rw_eqs, wl_rest = rest })
= length eqs_N + length eqs_X + length rw_eqs + length rest
extendWorkListEq :: RewriterSet -> Ct -> WorkList -> WorkList
extendWorkListEq rewriters ct
wl@(WL { wl_eqs_N = eqs_N, wl_eqs_X = eqs_X, wl_rw_eqs = rw_eqs })
| isEmptyRewriterSet rewriters -- A wanted that has not been rewritten
-- isEmptyRewriterSet: see Note [Prioritise Wanteds with empty RewriterSet]
-- in GHC.Tc.Types.Constraint
= if isNomEqPred (ctPred ct)
then wl { wl_eqs_N = ct : eqs_N }
else wl { wl_eqs_X = ct : eqs_X }
| otherwise
= wl { wl_rw_eqs = ct : rw_eqs }
extendWorkListEqs :: RewriterSet -> Bag Ct -> WorkList -> WorkList
-- Add [eq1,...,eqn] to the work-list
-- They all have the same rewriter set
-- The constraints will be solved in left-to-right order:
-- see Note [Work-list ordering] in GHC.Tc.Solver.Equality
--
-- Precondition: new_eqs is non-empty
extendWorkListEqs rewriters new_eqs
wl@(WL { wl_eqs_N = eqs_N, wl_eqs_X = eqs_X, wl_rw_eqs = rw_eqs })
| isEmptyRewriterSet rewriters
-- isEmptyRewriterSet: see Note [Prioritise Wanteds with empty RewriterSet]
-- in GHC.Tc.Types.Constraint
= case partitionBag is_nominal new_eqs of
(new_eqs_N, new_eqs_X)
| isEmptyBag new_eqs_N -> wl { wl_eqs_X = new_eqs_X `push_on_front` eqs_X }
| isEmptyBag new_eqs_X -> wl { wl_eqs_N = new_eqs_N `push_on_front` eqs_N }
| otherwise -> wl { wl_eqs_N = new_eqs_N `push_on_front` eqs_N
, wl_eqs_X = new_eqs_X `push_on_front` eqs_X }
-- These isEmptyBag tests are just trying
-- to avoid creating unnecessary thunks
| otherwise
= wl { wl_rw_eqs = new_eqs `push_on_front` rw_eqs }
where
push_on_front :: Bag Ct -> [Ct] -> [Ct]
-- push_on_front puts the new equlities on the front of the queue
push_on_front new_eqs eqs = foldr (:) eqs new_eqs
is_nominal ct = isNomEqPred (ctPred ct)
extendWorkListNonEq :: Ct -> WorkList -> WorkList
-- Extension by non equality
extendWorkListNonEq ct wl = wl { wl_rest = ct : wl_rest wl }
extendWorkListImplic :: Implication -> WorkList -> WorkList
extendWorkListImplic implic wl = wl { wl_implics = implic `consBag` wl_implics wl }
extendWorkListCt :: Ct -> WorkList -> WorkList
-- Agnostic about what kind of constraint
extendWorkListCt ct wl
= case classifyPredType (ctEvPred ev) of
EqPred {}
-> extendWorkListEq rewriters ct wl
ClassPred cls _ -- See Note [Prioritise class equalities]
| isEqualityClass cls
-> extendWorkListEq rewriters ct wl
_ -> extendWorkListNonEq ct wl
where
ev = ctEvidence ct
rewriters = ctEvRewriters ev
extendWorkListCtList :: [Ct] -> WorkList -> WorkList
extendWorkListCtList cts wl = foldr extendWorkListCt wl cts
extendWorkListCts :: Cts -> WorkList -> WorkList
extendWorkListCts cts wl = foldr extendWorkListCt wl cts
isEmptyWorkList :: WorkList -> Bool
isEmptyWorkList (WL { wl_eqs_N = eqs_N, wl_eqs_X = eqs_X
, wl_rest = rest, wl_implics = implics })
= null eqs_N && null eqs_X && null rest && isEmptyBag implics
emptyWorkList :: WorkList
emptyWorkList = WL { wl_eqs_N = [], wl_eqs_X = []
, wl_rw_eqs = [], wl_rest = [], wl_implics = emptyBag }
selectWorkItem :: WorkList -> Maybe (Ct, WorkList)
-- See Note [Prioritise equalities]
selectWorkItem wl@(WL { wl_eqs_N = eqs_N, wl_eqs_X = eqs_X
, wl_rw_eqs = rw_eqs, wl_rest = rest })
| ct:cts <- eqs_N = Just (ct, wl { wl_eqs_N = cts })
| ct:cts <- eqs_X = Just (ct, wl { wl_eqs_X = cts })
| ct:cts <- rw_eqs = Just (ct, wl { wl_rw_eqs = cts })
| ct:cts <- rest = Just (ct, wl { wl_rest = cts })
| otherwise = Nothing
-- Pretty printing
instance Outputable WorkList where
ppr (WL { wl_eqs_N = eqs_N, wl_eqs_X = eqs_X, wl_rw_eqs = rw_eqs
, wl_rest = rest, wl_implics = implics })
= text "WL" <+> (braces $
vcat [ ppUnless (null eqs_N) $
text "Eqs_N =" <+> vcat (map ppr eqs_N)
, ppUnless (null eqs_X) $
text "Eqs_X =" <+> vcat (map ppr eqs_X)
, ppUnless (null rw_eqs) $
text "RwEqs =" <+> vcat (map ppr rw_eqs)
, ppUnless (null rest) $
text "Non-eqs =" <+> vcat (map ppr rest)
, ppUnless (isEmptyBag implics) $
ifPprDebug (text "Implics =" <+> vcat (map ppr (bagToList implics)))
(text "(Implics omitted)")
])
{- *********************************************************************
* *
InertSet: the inert set
* *
* *
********************************************************************* -}
type CycleBreakerVarStack = NonEmpty (Bag (TcTyVar, TcType))
-- ^ a stack of (CycleBreakerTv, original family applications) lists
-- first element in the stack corresponds to current implication;
-- later elements correspond to outer implications
-- used to undo the cycle-breaking needed to handle
-- Note [Type equality cycles] in GHC.Tc.Solver.Equality
-- Why store the outer implications? For the use in mightEqualLater (only)
--
-- Why NonEmpty? So there is always a top element to add to
data InertSet
= IS { inert_cans :: InertCans
-- Canonical Given, Wanted
-- Sometimes called "the inert set"
, inert_cycle_breakers :: CycleBreakerVarStack
, inert_famapp_cache :: FunEqMap Reduction
-- Just a hash-cons cache for use when reducing family applications
-- only
--
-- If F tys :-> (co, rhs, flav),
-- then co :: F tys ~N rhs
-- all evidence is from instances or Givens; no coercion holes here
-- (We have no way of "kicking out" from the cache, so putting
-- wanteds here means we can end up solving a Wanted with itself. Bad)
, inert_solved_dicts :: DictMap DictCt
-- All Wanteds, of form (C t1 .. tn)
-- Always a dictionary solved by an instance decl; never an implict parameter
-- See Note [Solved dictionaries]
-- and Note [Do not add superclasses of solved dictionaries]
}
instance Outputable InertSet where
ppr (IS { inert_cans = ics
, inert_solved_dicts = solved_dicts })
= vcat [ ppr ics
, ppUnless (null dicts) $
text "Solved dicts =" <+> vcat (map ppr dicts) ]
where
dicts = bagToList (dictsToBag solved_dicts)
emptyInertCans :: InertCans
emptyInertCans
= IC { inert_eqs = emptyTyEqs
, inert_funeqs = emptyFunEqs
, inert_given_eq_lvl = topTcLevel
, inert_given_eqs = False
, inert_dicts = emptyDictMap
, inert_safehask = emptyDictMap
, inert_insts = []
, inert_irreds = emptyBag }
emptyInert :: InertSet
emptyInert
= IS { inert_cans = emptyInertCans
, inert_cycle_breakers = emptyBag :| []
, inert_famapp_cache = emptyFunEqs
, inert_solved_dicts = emptyDictMap }
{- Note [Solved dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When we apply a top-level instance declaration, we add the "solved"
dictionary to the inert_solved_dicts. In general, we use it to avoid
creating a new EvVar when we have a new goal that we have solved in
the past.
But in particular, we can use it to create *recursive* dictionaries.
The simplest, degenerate case is
instance C [a] => C [a] where ...
If we have
[W] d1 :: C [x]
then we can apply the instance to get
d1 = $dfCList d
[W] d2 :: C [x]
Now 'd1' goes in inert_solved_dicts, and we can solve d2 directly from d1.
d1 = $dfCList d
d2 = d1
See Note [Example of recursive dictionaries]
VERY IMPORTANT INVARIANT:
(Solved Dictionary Invariant)
Every member of the inert_solved_dicts is the result
of applying an instance declaration that "takes a step"
An instance "takes a step" if it has the form
dfunDList d1 d2 = MkD (...) (...) (...)
That is, the dfun is lazy in its arguments, and guarantees to
immediately return a dictionary constructor. NB: all dictionary
data constructors are lazy in their arguments.
This property is crucial to ensure that all dictionaries are
non-bottom, which in turn ensures that the whole "recursive
dictionary" idea works at all, even if we get something like
rec { d = dfunDList d dx }
See Note [Recursive superclasses] in GHC.Tc.TyCl.Instance.
Reason:
- All instances, except two exceptions listed below, "take a step"
in the above sense
- Exception 1: local quantified constraints have no such guarantee;
indeed, adding a "solved dictionary" when applying a quantified
constraint led to the ability to define unsafeCoerce
in #17267.
- Exception 2: the magic built-in instance for (~) has no
such guarantee. It behaves as if we had
class (a ~# b) => (a ~ b) where {}
instance (a ~# b) => (a ~ b) where {}
The "dfun" for the instance is strict in the coercion.
Anyway there's no point in recording a "solved dict" for
(t1 ~ t2); it's not going to allow a recursive dictionary
to be constructed. Ditto (~~) and Coercible.
THEREFORE we only add a "solved dictionary"
- when applying an instance declaration
- subject to Exceptions 1 and 2 above
In implementation terms
- GHC.Tc.Solver.Monad.addSolvedDict adds a new solved dictionary,
conditional on the kind of instance
- It is only called when applying an instance decl,
in GHC.Tc.Solver.Dict.tryInstances
- ClsInst.InstanceWhat says what kind of instance was
used to solve the constraint. In particular
* LocalInstance identifies quantified constraints
* BuiltinEqInstance identifies the strange built-in
instances for equality.
- ClsInst.instanceReturnsDictCon says which kind of
instance guarantees to return a dictionary constructor
Other notes about solved dictionaries
* See also Note [Do not add superclasses of solved dictionaries]
* The inert_solved_dicts field is not rewritten by equalities,
so it may get out of date.
* The inert_solved_dicts are all Wanteds, never givens
* We only cache dictionaries from top-level instances, not from
local quantified constraints. Reason: if we cached the latter
we'd need to purge the cache when bringing new quantified
constraints into scope, because quantified constraints "shadow"
top-level instances.
Note [Do not add superclasses of solved dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Every member of inert_solved_dicts is the result of applying a
dictionary function, NOT of applying superclass selection to anything.
Consider
class Ord a => C a where
instance Ord [a] => C [a] where ...
Suppose we are trying to solve
[G] d1 : Ord a
[W] d2 : C [a]
Then we'll use the instance decl to give
[G] d1 : Ord a Solved: d2 : C [a] = $dfCList d3
[W] d3 : Ord [a]
We must not add d4 : Ord [a] to the 'solved' set (by taking the
superclass of d2), otherwise we'll use it to solve d3, without ever
using d1, which would be a catastrophe.
Solution: when extending the solved dictionaries, do not add superclasses.
That's why each element of the inert_solved_dicts is the result of applying
a dictionary function.
Note [Example of recursive dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
--- Example 1
data D r = ZeroD | SuccD (r (D r));
instance (Eq (r (D r))) => Eq (D r) where
ZeroD == ZeroD = True
(SuccD a) == (SuccD b) = a == b
_ == _ = False;
equalDC :: D [] -> D [] -> Bool;
equalDC = (==);
We need to prove (Eq (D [])). Here's how we go:
[W] d1 : Eq (D [])
By instance decl of Eq (D r):
[W] d2 : Eq [D []] where d1 = dfEqD d2
By instance decl of Eq [a]:
[W] d3 : Eq (D []) where d2 = dfEqList d3
d1 = dfEqD d2
Now this wanted can interact with our "solved" d1 to get:
d3 = d1
-- Example 2:
This code arises in the context of "Scrap Your Boilerplate with Class"
class Sat a
class Data ctx a
instance Sat (ctx Char) => Data ctx Char -- dfunData1
instance (Sat (ctx [a]), Data ctx a) => Data ctx [a] -- dfunData2
class Data Maybe a => Foo a
instance Foo t => Sat (Maybe t) -- dfunSat
instance Data Maybe a => Foo a -- dfunFoo1
instance Foo a => Foo [a] -- dfunFoo2
instance Foo [Char] -- dfunFoo3
Consider generating the superclasses of the instance declaration
instance Foo a => Foo [a]
So our problem is this
[G] d0 : Foo t
[W] d1 : Data Maybe [t] -- Desired superclass
We may add the given in the inert set, along with its superclasses
Inert:
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
WorkList
[W] d1 : Data Maybe [t]
Solve d1 using instance dfunData2; d1 := dfunData2 d2 d3
Inert:
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
WorkList:
[W] d2 : Sat (Maybe [t])
[W] d3 : Data Maybe t
Now, we may simplify d2 using dfunSat; d2 := dfunSat d4
Inert:
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
d2 : Sat (Maybe [t])
WorkList:
[W] d3 : Data Maybe t
[W] d4 : Foo [t]
Now, we can just solve d3 from d01; d3 := d01
Inert
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
d2 : Sat (Maybe [t])
WorkList
[W] d4 : Foo [t]
Now, solve d4 using dfunFoo2; d4 := dfunFoo2 d5
Inert
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
d2 : Sat (Maybe [t])
d4 : Foo [t]
WorkList:
[W] d5 : Foo t
Now, d5 can be solved! d5 := d0
Result
d1 := dfunData2 d2 d3
d2 := dfunSat d4
d3 := d01
d4 := dfunFoo2 d5
d5 := d0
-}
{- *********************************************************************
* *
InertCans: the canonical inerts
* *
* *
********************************************************************* -}
{- Note [Tracking Given equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
For reasons described in (UNTOUCHABLE) in GHC.Tc.Utils.Unify
Note [Unification preconditions], we can't unify
alpha[2] ~ Int
under a level-4 implication if there are any Given equalities
bound by the implications at level 3 of 4. To that end, the
InertCans tracks
inert_given_eq_lvl :: TcLevel
-- The TcLevel of the innermost implication that has a Given
-- equality of the sort that make a unification variable untouchable
-- (see Note [Unification preconditions] in GHC.Tc.Utils.Unify).
We update inert_given_eq_lvl whenever we add a Given to the
inert set, in updGivenEqs.
Then a unification variable alpha[n] is untouchable iff
n < inert_given_eq_lvl
that is, if the unification variable was born outside an
enclosing Given equality.
Exactly which constraints should trigger (UNTOUCHABLE), and hence
should update inert_given_eq_lvl?
(TGE1) We do /not/ need to worry about let-bound skolems, such as
forall[2] a. a ~ [b] => blah
See Note [Let-bound skolems] and the isOuterTyVar tests in `updGivenEqs`
(TGE2) However, solely to support better error messages (see Note [HasGivenEqs] in
GHC.Tc.Types.Constraint) we also track these "local" equalities in the
boolean inert_given_eqs field. This field is used only subsequntly (see
`getHasGivenEqs`), to set the ic_given_eqs field to LocalGivenEqs.
(TGE3) Consider an implication
forall[2]. beta[1] => alpha[1] ~ Int
where beta is a unification variable that has already been unified
to () in an outer scope. Then alpha[1] is perfectly touchable and
we can unify alpha := Int. So when deciding whether the givens contain
an equality, we should canonicalise first, rather than just looking at
the /original/ givens (#8644).
(TGE4) However, we must take account of *potential* equalities. Consider the
same example again, but this time we have /not/ yet unified beta:
forall[2] beta[1] => ...blah...
Because beta might turn into an equality, updGivenEqs conservatively
treats it as a potential equality, and updates inert_give_eq_lvl
(TGE5) We should not look at the equality relation involved (nominal vs
representational), because representational equalities can still
imply nominal ones. For example, if (G a ~R G b) and G's argument's
role is nominal, then we can deduce a ~N b.
Historical note: prior to #24938 we also ignored Given equalities that
did not mention an "outer" type variable. But that is wrong, as #24938
showed. Another example is immortalised in test LocalGivenEqs2
data T where
MkT :: F a ~ G b => a -> b -> T
f (MkT _ _) = True
We should not infer the type for `f`; let-bound-skolems does not apply.
Note [Let-bound skolems]
~~~~~~~~~~~~~~~~~~~~~~~~
If * the inert set contains a canonical Given CEqCan (a ~ ty)
and * 'a' is a skolem bound in this very implication,
then:
a) The Given is pretty much a let-binding, like
f :: (a ~ b->c) => a -> a
Here the equality constraint is like saying
let a = b->c in ...
It is not adding any new, local equality information,
and hence can be ignored by has_given_eqs
b) 'a' will have been completely substituted out in the inert set,
so we can safely discard it.
For an example, see #9211.
The actual test is in `isLetBoundSkolemCt`
Wrinkles:
(LBS1) See GHC.Tc.Utils.Unify Note [Deeper level on the left] for how we ensure
that the correct variable is on the left of the equality when both are
tyvars.
(LBS2) We also want this to work for
forall a. [G] F b ~ a (CEqCt with TyFamLHS)
Here the Given will have a TyFamLHS, with the skolem-bound tyvar on the RHS.
See tests T24938a, and LocalGivenEqs.
(LBS3) Happily (LBS2) also makes cycle-breakers work. Suppose we have
forall a. [G] (F a) Int ~ a
where F has arity 1, and `a` is the locally-bound skolem. Then, as
Note [Type equality cycles] explains, we split into
[G] F a ~ cbv, [G] cbv Int ~ a
where `cbv` is the cycle breaker variable. But cbv has the same level
as `a`, so `isOuterTyVar` (called in `isLetBoundSkolemCt`) will return False.
This actually matters occasionally: see test LocalGivenEqs.
You might wonder whether the skolem really needs to be bound "in the
very same implication" as the equality constraint.
Consider this (c.f. #15009):
data S a where
MkS :: (a ~ Int) => S a
g :: forall a. S a -> a -> blah
g x y = let h = \z. ( z :: Int
, case x of
MkS -> [y,z])
in ...
From the type signature for `g`, we get `y::a` . Then when we
encounter the `\z`, we'll assign `z :: alpha[1]`, say. Next, from the
body of the lambda we'll get
[W] alpha[1] ~ Int -- From z::Int
[W] forall[2]. (a ~ Int) => [W] alpha[1] ~ a -- From [y,z]
Now, unify alpha := a. Now we are stuck with an unsolved alpha~Int!
So we must treat alpha as untouchable under the forall[2] implication.
Possible future improvements. The current test just looks to see whether one
side of an equality is a locally-bound skolem. But actually we could, in
theory, do better: if one side (or both sides, actually) of an equality
ineluctably mentions a local skolem, then the equality cannot possibly impact
types outside of the implication (because doing to would cause those types to be
ill-scoped). The problem is the "ineluctably": this means that no expansion,
other solving, etc., could possibly get rid of the variable. This is hard,
perhaps impossible, to know for sure, especially when we think about type family
interactions. (And it's a user-visible property so we don't want it to be hard
to predict.) So we keep the existing check, looking for one lone variable,
because we're sure that variable isn't going anywhere.
Note [Detailed InertCans Invariants]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The InertCans represents a collection of constraints with the following properties:
* All canonical
* No two dictionaries with the same head
* No two CIrreds with the same type
* Family equations inert wrt top-level family axioms
* Dictionaries have no matching top-level instance
* Given family or dictionary constraints don't mention touchable
unification variables
* Non-CEqCan constraints are fully rewritten with respect
to the CEqCan equalities (modulo eqCanRewrite of course;
eg a wanted cannot rewrite a given)
* CEqCan equalities: see Note [inert_eqs: the inert equalities]
Also see documentation in Constraint.Ct for a list of invariants
Note [inert_eqs: the inert equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Our main invariant:
the EqCts in inert_eqs should be a
terminating generalised substitution
-------------- Definition [Can-rewrite relation] --------------
A "can-rewrite" relation between flavours, written f1 >= f2, is a
binary relation with the following properties
(R1) >= is transitive
(R2) If f1 >= f, and f2 >= f,
then either f1 >= f2 or f2 >= f1
(See Note [Why R2?].)
Lemma (L0). If f1 >= f then f1 >= f1
Proof. By property (R2), with f1=f2
--------- Definition [Generalised substitution] ---------------
A "generalised substitution" S is a set of triples (lhs -f-> t), where
- lhs is a type variable or an exactly-saturated type family application
(that is, lhs is a CanEqLHS)
- t is a type
- f is a flavour
such that
(WF1) if (lhs1 -f1-> t1) in S
(lhs2 -f2-> t2) in S
then (f1 >= f2) implies that lhs1 does not appear within lhs2
(WF2) if (lhs -f-> t) is in S, then t /= lhs
(WF3) No LHS in S is rewritable in an RHS in S,
in the argument of a type family application (F ty1..tyn)
where F heads a LHS in S
--------- Definition [Applying a generalised substitution] ----------
If S is a generalised substitution
S(f,lhs) = rhs, if (lhs -fs-> rhs) in S, and fs >= f
S(f,T t1..tn) = T S(f1,t1)..S(fn,tn)
S(f,t1 t2) = S(f,t1) S(f_N,t2)
S(f,t) = t
Here f1..fn are obtained from f and T using the roles of T, and f_N is
the nominal version of f. See Note [Flavours with roles].
Notation: repeated application.
S^0(f,t) = t
S^(n+1)(f,t) = S(f, S^n(t))
S*(f,t) is the result of applying S until you reach a fixpoint
--------- Definition [Terminating generalised substitution] ---------
A generalised substitution S is *terminating* iff
(IG1) for every f,t, there is an n such that
S^n(f,t) = S^(n+1)(f,t)
By (IG1) we define S*(f,t) to be the result of exahaustively
applying S(f,_) to t.
--------- End of definitions ------------------------------------
Rationale for (WF1)-(WF3)
-------------------------
* (WF1) guarantees that S is well-defined /as a function/;
see Theorem (S is a function)
Theorem (S is a function): S(f,t0) is well defined as a function.
Proof: Suppose (lhs -f1-> t1) and (lhs -f2-> t2) are both in S,
and f1 >= f and f2 >= f
Then by (R2) f1 >= f2 or f2 >= f1, which contradicts (WF1)
Note: this argument isn't quite right. WF1 ensures that lhs1 does
not appear inside lhs2, and that guarantees confluence. But I can't quite
see how to make that argument precise.
* (WF2) is a bit trivial. It means that if S is terminating, so that
S^(n+1)(f,t) = S^n(f,t), then there is no LHS of S in S^n(f,t). We
never get a silly infinite sequence a -> a -> a -> a .... which is
technically a fixed point but would still go on for ever.
* (WF3) is need for the termination proof.
Note that termination is not the same as idempotence. To apply S to a
type, you may have to apply it recursively. But termination does
guarantee that this recursive use will terminate.
Note [Why R2?]
~~~~~~~~~~~~~~
R2 states that, if we have f1 >= f and f2 >= f, then either f1 >= f2 or f2 >=
f1. If we do not have R2, we will easily fall into a loop.
To see why, imagine we have f1 >= f, f2 >= f, and that's it. Then, let our
inert set S = {a -f1-> b, b -f2-> a}. Computing S(f,a) does not terminate. And
yet, we have a hard time noticing an occurs-check problem when building S, as
the two equalities cannot rewrite one another.
R2 actually restricts our ability to accept user-written programs. See
Note [Avoiding rewriting cycles] in GHC.Tc.Types.Constraint for an example.
Note [Rewritable]
~~~~~~~~~~~~~~~~~
Definition. A CanEqLHS lhs is *rewritable* in a type t if the
lhs tree appears as a subtree within t without traversing any of the following
components of t:
* coercions (whether they appear in casts CastTy or as arguments CoercionTy)
* kinds of variable occurrences
The check for rewritability *does* look in kinds of the bound variable of a
ForAllTy.
The reason for this definition is that the rewriter does not rewrite in coercions
or variables' kinds. In turn, the rewriter does not need to rewrite there because
those places are never used for controlling the behaviour of the solver: these
places are not used in matching instances or in decomposing equalities.
This definition is used by the anyRewritableXXX family of functions and is meant
to model the actual behaviour in GHC.Tc.Solver.Rewrite.
Goal: If lhs is not rewritable in t, then t is a fixpoint of the generalised
substitution containing only {lhs -f*-> t'}, where f* is a flavour such that f* >= f
for all f.
Wrinkles
* Taking roles into account: we must consider a rewrite at a given role. That is,
a rewrite arises from some equality, and that equality has a role associated
with it. As we traverse a type, we track what role we are allowed to rewrite with.
For example, suppose we have an inert [G] b ~R# Int. Then b is rewritable in
Maybe b but not in F b, where F is a type function. This role-aware logic is
present in both the anyRewritableXXX functions and in the rewriter.
See also Note [anyRewritableTyVar must be role-aware] in GHC.Tc.Utils.TcType.
* There is one exception to the claim that non-rewritable parts of the tree do
not affect the solver: we sometimes do an occurs-check to decide e.g. how to
orient an equality. (See the comments on GHC.Tc.Solver.Equality.canEqTyVarFunEq.)
Accordingly, the presence of a variable in a kind or coercion just might
influence the solver. Here is an example:
type family Const x y where
Const x y = x
AxConst :: forall x y. Const x y ~# x
alpha :: Const Type Nat
[W] alpha ~ Int |> (sym (AxConst Type alpha) ;;
AxConst Type alpha ;;
sym (AxConst Type Nat))
The cast is clearly ludicrous (it ties together a cast and its symmetric
version), but we can't quite rule it out. (See (EQ1) from Note [Respecting
definitional equality] in GHC.Core.TyCo.Rep to see why we need the Const Type
Nat bit.) And yet this cast will (quite rightly) prevent alpha from unifying
with the RHS. I (Richard E) don't have an example of where this problem can
arise from a Haskell program, but we don't have an air-tight argument for why
the definition of *rewritable* given here is correct.
Note [Extending the inert equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Main Theorem [Stability under extension]
GIVEN a "work item" [lhs_w -fw-> rhs_w]
and a terminating generalised substitution S,
SUCH THAT
(T1) S(fw,lhs_w) = lhs_w -- LHS of work-item is a fixpoint of S(fw,_)
(T2) S(fw,rhs_w) = rhs_w -- RHS of work-item is a fixpoint of S(fw,_)
(T3) lhs_w not in rhs_w -- No occurs check in the work item
-- If lhs is a type family application, we require only that
-- lhs is not *rewritable* in rhs_w. See Note [Rewritable] and
-- Note [EqCt occurs check] in GHC.Tc.Types.Constraint.
(T4) no [lhs_s -fs-> rhs_s] in S meets [The KickOut Criteria]
(i.e. we already kicked any such items out!)
THEN the extended substitution T = S+(lhs_w -fw-> rhs_w)
is a terminating generalised substitution
How do we establish these conditions?
* (T1) and (T2) are guaranteed by exhaustively rewriting the work-item
with S(fw,_).
* (T3) is guaranteed by an occurs-check on the work item.
This is done during canonicalisation, in checkTypeEq; invariant
(TyEq:OC) of CEqCan. See also Note [EqCt occurs check] in GHC.Tc.Types.Constraint.
* (T4) is established by GHC.Tc.Solver.Monad.kickOutRewritable. If the inert
set contains a triple that meets the KickOut Criteria, we kick it out and
add it to the work list for later re-examination. See
Note [The KickOut Criteria]
Theorem: T (defined in "THEN" above) is a generalised substitution;
that is, it satisfies (WF1)-(WF3)
Proof:
(WF1) Suppose we are adding [lhs_w -fw-> rhs_w], and [lhs_s -fs-> rhs_s] is in S.
Then:
- by (T1) if fs>=fw, lhs_s does not occur within lhs_w.
- by (KK1) if fw>=fs, lhs_w is not rewritable in lhs_s, or we'd have
kicked out the stable constraint.
(WF2) is directly guaranteed by (T3)
(WF3) No lhs_s in S is rewritable in rhs_w at all, because of (T2)
And (KK2) guarantees that lhs_w is not rewritable under a type
family in rhs_s
Note [The KickOut Criteria]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
Kicking out is a Bad Thing:
* It means we have to re-process a constraint. The less we kick out, the better.
* In the limit, kicking can lead to non-termination: imagine that we /always/
kick out the entire inert set!
* Because (mid 2024) we don't support sharing in constraints, excessive kicking out
can lead to exponentially big constraints (#24984).
So we seek to do as little kicking out as possible. For example, consider this,
which happens a lot:
Inert: g1: a ~ Maybe b
Work: g2: b ~ Int
We do /not/ kick out g1 when adding g2. The new substitution S' = {g1,g2} is still
/terminating/ but it is not /idmempotent/. To apply S' to, say, (Tree a), we may
need to apply it twice: Tree a --> Tree (Maybe b) --> Tree (Maybe Int)
Here are the KickOut Criteria:
When adding [lhs_w -fw-> rhs_w] to a well-formed terminating substitution S,
element [lhs_s -fs-> rhs_s] in S meets the KickOut Criteria if:
(KK0) fw >= fs AND any of (KK1), (KK2) or (KK3) hold
* (KK1: satisfy WF1) `lhs_w` is rewritable in `lhs_s`.
* (KK2: termination) `lhs_w` is rewritable in `rhs_s` in these positions:
If not(fs>=fw)
then (KK2a) anywhere
else (KK2b) look only in the argument of type family applications,
whose type family heads some LHS in `S`
* (KK3: completeness)
If not(fs >= fw) -- If fs can rewrite fw, kick-out is redundant/harmful
* (KK3a) If the role of `fs` is Nominal:
kick out if `rhs_s = lhs_w`
* (KK3b) If the role of `fs` is Representational:
kick out if `rhs_s` is of form `(lhs_w t1 .. tn)`
Rationale
* (KK0) kick out only if `fw` can rewrite `fs`.
Reason: suppose we kick out (lhs1 -fs-> s), and add (lhs -fw-> t) to the ineart
set. The latter can't rewrite the former, so the kick-out achieved nothing
* (KK1) `lhs_w` is rewritable in `lhs_s`.
Reason: needed to guarantee (WF1). See Theorem: T is well formed
* (KK2) see Note [KK2: termination of the extended substitution]
* (KK3) see Note [KK3: completeness of solving]
The above story is a bit vague wrt roles, but the code is not.
See Note [Flavours with roles]
Note [KK2: termination of the extended substitution]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Proving termination of the extended substitution T is surprisingly tricky.
* Reason for (KK2a). Consider
Work: [G] b ~ a
Inert: [W] a ~ b
If we don't kick out the inert, then we get a loop on e.g. [W] a ~ Int.
But if both were Wanted we really should not kick out (the substitution does not
have to be idempotent). So we only look everywhere for the `lhs_w` if
not (fs>=fw), that is the inert item cannot rewrite the work item. So in the
above example we will kick out; but if both were Wanted we won't.
* Reason for (KK2b). Consider the case where (fs >= fw)
Work: [G] a ~ Int
Inert: [G] F Int ~ F a
If we just added the work item, the substitution would loop on type (F Int).
So we must kick out the inert item, even though (fs>=fw). (KK2b) does this
by looking for lhs_w under type family applications in rhs_s.
(KK2b) makes kick-out less aggressive by looking only under type-family applications,
in the case where (fs >= fw), and that made a /huge/ difference to #24944.
Tricky examples in: #19042, #17672, #24984. The last (#24984) is particular subtle:
Inert: [W] g1: F a0 ~ F a1
[W] g2: F a2 ~ F a1
[W] g3: F a3 ~ F a1
Now we add [W] g4: F a1 ~ F a7. Should we kick out g1,g2,g3? No! The
substitution doesn't need to be idempotent, merely terminating. And in #24984
it turned out that we kept adding one new constraint and kicking out all the
previous inert ones (and that rewriting led to exponentially big constraints due
to lack of contraint sharing.) So we only want to look /under/ type family applications.
The proof is hard. We start by ignoring flavours. Suppose that:
* We are adding [lhs_w -fw-> rhs_w] to a well-formed, terminating substitution S.
* None of the constraints in S meet the KickOut Criteria.
* Define T = S+[lhs_w -fw-> rhs_w]
* `f` is an arbitrary flavour
Lemma 1: for any lhs_s in S, T*(f,lhs_s) terminates.
Proof.
* We know that r1 = S*(f,lhs_s) terminates.
* Moreover, we know there are no occurrences of lhs_w under a type family (which
is the head of a LHS) in r1 (KK2)+(WF3). We need (WF3) because you might wonder
what if rhs_s is (F a), and [a --> lhs_w] was in S. But (WF3) prevents that.
* Define r2 = r1{rhs_w/lhs_w}. We know that rhs_w has no occurrences of any lhs in S,
nor of lhs_w.
* Since any occurrence of lhs_w does not occur under a type family, the substitution
won't make any F t1..tn ~ s in S match.
* So r2 is a fixed point of T.
Lemma 2: T*(f,lhs_w) teminates.
Proof: no occurrences of any LHS in rhs_w.
Theorem. For any type r, T*(r) terminates.
Proof:
1. Consider any sub-term of r, which is a LHS of T.
- Rewrite it with T*; this terminates (Lemma 1).
- Do this simultaneously to all sub-terms that match a LHS of T, yielding r1.
2. Could this new r1 have a sub-term that is an LHS of T? Yes, but only if r has a
sub-term F w, and w rewrote in Step 1 to w' and F w' matches a LHS in T.
3. Very well: apply step 1 again, but note that /doing so consumes one of the family
applications in the original r/.
4. After Step 1 either we have reached a fixed point, or we repeat Step 1 consuming at
least one family application of r.
5. There are only a finite number of family applications in r, so this process terminates.
Example:
Inert set: gs : F Int ~ b
Work item: gw : b ~ Int
F (F (F b)) --[gw]--> F (F (F Int)) --[gs]--> F (F b)
--[gw]--> F (F Int) --[gs]--> F b
--[gw]--> F Int --[gs]--> b
--[gw]--> Int
Notice that each iteration of Step 1 strips off one of the layers of F, all
of which were in the original r.
The argument is even more tricky when flavours are involved, and we have not
fleshed it out in detail.
Note [KK3: completeness of solving]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
(KK3) is not necessary for the extended substitution
to be terminating. In fact (KK0) could be made stronger by saying
... then (not (fw >= fs) or not (fs >= fs))
But it's not enough for S to be /terminating/; we also want /completeness/.
That is, we want to be able to solve all soluble wanted equalities.
Suppose we have
work-item b -G-> a
inert-item a -W-> b
Assuming (G >= W) but not (W >= W), this fulfills all the conditions,
so we could extend the inerts, thus:
inert-items b -G-> a
a -W-> b
But if we kicked-out the inert item, we'd get
work-item a -W-> b
inert-item b -G-> a
Then rewrite the work-item gives us (a -W-> a), which is soluble via Refl.
So we add one more clause (KK3) to the kick-out criteria:
* (KK3: completeness)
If not(fs >= fw) (KK3a)
* (KK3b) If the role of `fs` is Nominal:
kick out if `rhs_s = lhs_w`
* (KK3c) If the role of `fs` is Representational:
kick out if `rhs_s` is of form `(lhs_w t1 .. tn)`
Wrinkles:
* (KK3a) All this can only happen if the work-item can rewrite the inert
one, /but not vice versa/; that is not(fs >= fw). It is useless to kick
out if (fs >= fw) becuase then the work-item is already fully rewritten
by the inert item. And too much kick-out is positively harmful.
(Historical example #14363.)
* (KK3b) addresses teh main example above for KK3. Another way to understand
(KK3b) is that we treat an inert item
a -f-> b
in the same way as
b -f-> a
So if we kick out one, we should kick out the other. The orientation
is somewhat accidental.
* (KK3c) When considering roles, we also need the second clause (KK3b). Consider
work-item c -G/N-> a
inert-item a -W/R-> b c
The work-item doesn't get rewritten by the inert, because (>=) doesn't hold.
But we don't kick out the inert item because not (W/R >= W/R). So we just
add the work item. But then, consider if we hit the following:
work-item b -G/N-> Id
inert-items a -W/R-> b c
c -G/N-> a
where
newtype Id x = Id x
For similar reasons, if we only had (KK3a), we wouldn't kick the
representational inert out. And then, we'd miss solving the inert, which now
reduced to reflexivity.
The solution here is to kick out representational inerts whenever the lhs of a
work item is "exposed", where exposed means being at the head of the top-level
application chain (lhs t1 .. tn). See head_is_new_lhs. This is encoded in
(KK3c)).
Note [Flavours with roles]
~~~~~~~~~~~~~~~~~~~~~~~~~~
The system described in Note [inert_eqs: the inert equalities]
discusses an abstract
set of flavours. In GHC, flavours have two components: the flavour proper,
taken from {Wanted, Given} and the equality relation (often called
role), taken from {NomEq, ReprEq}.
When substituting w.r.t. the inert set,
as described in Note [inert_eqs: the inert equalities],
we must be careful to respect all components of a flavour.
For example, if we have
inert set: a -G/R-> Int
b -G/R-> Bool
type role T nominal representational
and we wish to compute S(W/R, T a b), the correct answer is T a Bool, NOT
T Int Bool. The reason is that T's first parameter has a nominal role, and
thus rewriting a to Int in T a b is wrong. Indeed, this non-congruence of
substitution means that the proof in Note [inert_eqs: the inert equalities] may
need to be revisited, but we don't think that the end conclusion is wrong.
-}
data InertCans -- See Note [Detailed InertCans Invariants] for more
= IC { inert_eqs :: InertEqs
-- See Note [inert_eqs: the inert equalities]
-- All EqCt with a TyVarLHS; index is the LHS tyvar
-- Domain = skolems and untouchables; a touchable would be unified
, inert_funeqs :: InertFunEqs
-- All EqCt with a TyFamLHS; index is the whole family head type.
-- LHS is fully rewritten (modulo eqCanRewrite constraints)
-- wrt inert_eqs
-- Can include both [G] and [W]
, inert_dicts :: DictMap DictCt
-- Dictionaries only
-- All fully rewritten (modulo flavour constraints)
-- wrt inert_eqs
, inert_insts :: [QCInst]
, inert_safehask :: DictMap DictCt
-- Failed dictionary resolution due to Safe Haskell overlapping
-- instances restriction. We keep this separate from inert_dicts
-- as it doesn't cause compilation failure, just safe inference
-- failure.
--
-- ^ See Note [Safe Haskell Overlapping Instances Implementation]
-- in GHC.Tc.Solver
, inert_irreds :: InertIrreds
-- Irreducible predicates that cannot be made canonical,
-- and which don't interact with others (e.g. (c a))
-- and insoluble predicates (e.g. Int ~ Bool, or a ~ [a])
, inert_given_eq_lvl :: TcLevel
-- The TcLevel of the innermost implication that has a Given
-- equality of the sort that make a unification variable untouchable
-- (see Note [Unification preconditions] in GHC.Tc.Utils.Unify).
-- See Note [Tracking Given equalities]
, inert_given_eqs :: Bool
-- True <=> The inert Givens *at this level* (tcl_tclvl)
-- could includes at least one equality /other than/ a
-- let-bound skolem equality.
-- Reason: report these givens when reporting a failed equality
-- See Note [Tracking Given equalities]
}
type InertEqs = DTyVarEnv EqualCtList
type InertFunEqs = FunEqMap EqualCtList
type InertIrreds = Bag IrredCt
instance Outputable InertCans where
ppr (IC { inert_eqs = eqs
, inert_funeqs = funeqs
, inert_dicts = dicts
, inert_safehask = safehask
, inert_irreds = irreds
, inert_given_eq_lvl = ge_lvl
, inert_given_eqs = given_eqs
, inert_insts = insts })
= braces $ vcat
[ ppUnless (isEmptyDVarEnv eqs) $
text "Equalities ="
<+> pprBag (foldTyEqs consBag eqs emptyBag)
, ppUnless (isEmptyTcAppMap funeqs) $
text "Type-function equalities ="
<+> pprBag (foldFunEqs consBag funeqs emptyBag)
, ppUnless (isEmptyTcAppMap dicts) $
text "Dictionaries =" <+> pprBag (dictsToBag dicts)
, ppUnless (isEmptyTcAppMap safehask) $
text "Safe Haskell unsafe overlap =" <+> pprBag (dictsToBag safehask)
, ppUnless (isEmptyBag irreds) $
text "Irreds =" <+> pprBag irreds
, ppUnless (null insts) $
text "Given instances =" <+> vcat (map ppr insts)
, text "Innermost given equalities =" <+> ppr ge_lvl
, text "Given eqs at this level =" <+> ppr given_eqs
]
{- *********************************************************************
* *
Inert equalities
* *
********************************************************************* -}
emptyTyEqs :: InertEqs
emptyTyEqs = emptyDVarEnv
addEqToCans :: TcLevel -> EqCt -> InertCans -> InertCans
addEqToCans tc_lvl eq_ct@(EqCt { eq_lhs = lhs })
ics@(IC { inert_funeqs = funeqs, inert_eqs = eqs })
= updGivenEqs tc_lvl (CEqCan eq_ct) $
case lhs of
TyFamLHS tc tys -> ics { inert_funeqs = addCanFunEq funeqs tc tys eq_ct }
TyVarLHS tv -> ics { inert_eqs = addTyEq eqs tv eq_ct }
addTyEq :: InertEqs -> TcTyVar -> EqCt -> InertEqs
addTyEq old_eqs tv ct
= extendDVarEnv_C add_eq old_eqs tv [ct]
where
add_eq old_eqs _ = addToEqualCtList ct old_eqs
foldTyEqs :: (EqCt -> b -> b) -> InertEqs -> b -> b
foldTyEqs k eqs z
= foldDVarEnv (\cts z -> foldr k z cts) z eqs
findTyEqs :: InertCans -> TyVar -> [EqCt]
findTyEqs icans tv = concat @Maybe (lookupDVarEnv (inert_eqs icans) tv)
delEq :: EqCt -> InertCans -> InertCans
delEq (EqCt { eq_lhs = lhs, eq_rhs = rhs }) ic = case lhs of
TyVarLHS tv
-> ic { inert_eqs = alterDVarEnv upd (inert_eqs ic) tv }
TyFamLHS tf args
-> ic { inert_funeqs = alterTcApp (inert_funeqs ic) tf args upd }
where
isThisOne :: EqCt -> Bool
isThisOne (EqCt { eq_rhs = t1 }) = tcEqTypeNoKindCheck rhs t1
upd :: Maybe EqualCtList -> Maybe EqualCtList
upd (Just eq_ct_list) = filterEqualCtList (not . isThisOne) eq_ct_list
upd Nothing = Nothing
findEq :: InertCans -> CanEqLHS -> [EqCt]
findEq icans (TyVarLHS tv) = findTyEqs icans tv
findEq icans (TyFamLHS fun_tc fun_args)
= concat @Maybe (findFunEq (inert_funeqs icans) fun_tc fun_args)
{-# INLINE partition_eqs_container #-}
partition_eqs_container
:: forall container
. container -- empty container
-> (forall b. (EqCt -> b -> b) -> container -> b -> b) -- folder
-> (EqCt -> container -> container) -- extender
-> (EqCt -> Bool)
-> container
-> ([EqCt], container)
partition_eqs_container empty_container fold_container extend_container pred orig_inerts
= fold_container folder orig_inerts ([], empty_container)
where
folder :: EqCt -> ([EqCt], container) -> ([EqCt], container)
folder eq_ct (acc_true, acc_false)
| pred eq_ct = (eq_ct : acc_true, acc_false)
| otherwise = (acc_true, extend_container eq_ct acc_false)
partitionInertEqs :: (EqCt -> Bool) -- EqCt will always have a TyVarLHS
-> InertEqs
-> ([EqCt], InertEqs)
partitionInertEqs = partition_eqs_container emptyTyEqs foldTyEqs addInertEqs
addInertEqs :: EqCt -> InertEqs -> InertEqs
-- Precondition: CanEqLHS is a TyVarLHS
addInertEqs eq_ct@(EqCt { eq_lhs = TyVarLHS tv }) eqs = addTyEq eqs tv eq_ct
addInertEqs other _ = pprPanic "extendInertEqs" (ppr other)
------------------------
addCanFunEq :: InertFunEqs -> TyCon -> [TcType] -> EqCt -> InertFunEqs
addCanFunEq old_eqs fun_tc fun_args ct
= alterTcApp old_eqs fun_tc fun_args upd
where
upd (Just old_equal_ct_list) = Just $ addToEqualCtList ct old_equal_ct_list
upd Nothing = Just [ct]
foldFunEqs :: (EqCt -> b -> b) -> FunEqMap EqualCtList -> b -> b
foldFunEqs k fun_eqs z = foldTcAppMap (\eqs z -> foldr k z eqs) fun_eqs z
partitionFunEqs :: (EqCt -> Bool) -- EqCt will have a TyFamLHS
-> InertFunEqs
-> ([EqCt], InertFunEqs)
partitionFunEqs = partition_eqs_container emptyFunEqs foldFunEqs addFunEqs
addFunEqs :: EqCt -> InertFunEqs -> InertFunEqs
-- Precondition: EqCt is a TyFamLHS
addFunEqs eq_ct@(EqCt { eq_lhs = TyFamLHS tc args }) fun_eqs
= addCanFunEq fun_eqs tc args eq_ct
addFunEqs other _ = pprPanic "extendFunEqs" (ppr other)
{- *********************************************************************
* *
Inert Dicts
* *
********************************************************************* -}
updDicts :: (DictMap DictCt -> DictMap DictCt) -> InertCans -> InertCans
updDicts upd ics = ics { inert_dicts = upd (inert_dicts ics) }
delDict :: DictCt -> DictMap a -> DictMap a
delDict (DictCt { di_cls = cls, di_tys = tys }) m
= delTcApp m (classTyCon cls) tys
addDict :: DictCt -> DictMap DictCt -> DictMap DictCt
addDict item@(DictCt { di_cls = cls, di_tys = tys }) dm
= insertTcApp dm (classTyCon cls) tys item
addSolvedDict :: DictCt -> DictMap DictCt -> DictMap DictCt
addSolvedDict item@(DictCt { di_cls = cls, di_tys = tys }) dm
= insertTcApp dm (classTyCon cls) tys item
filterDicts :: (DictCt -> Bool) -> DictMap DictCt -> DictMap DictCt
filterDicts f m = filterTcAppMap f m
partitionDicts :: (DictCt -> Bool) -> DictMap DictCt -> (Bag DictCt, DictMap DictCt)
partitionDicts f m = foldTcAppMap k m (emptyBag, emptyDictMap)
where
k ct (yeses, noes) | f ct = (ct `consBag` yeses, noes)
| otherwise = (yeses, addDict ct noes)
{- *********************************************************************
* *
Inert Irreds
* *
********************************************************************* -}
addIrredToCans :: TcLevel -> IrredCt -> InertCans -> InertCans
addIrredToCans tc_lvl irred ics
= updGivenEqs tc_lvl (CIrredCan irred) $
updIrreds (addIrred irred) ics
addIrreds :: [IrredCt] -> InertIrreds -> InertIrreds
addIrreds extras irreds
| null extras = irreds
| otherwise = irreds `unionBags` listToBag extras
addIrred :: IrredCt -> InertIrreds -> InertIrreds
addIrred extra irreds = irreds `snocBag` extra
updIrreds :: (InertIrreds -> InertIrreds) -> InertCans -> InertCans
updIrreds upd ics = ics { inert_irreds = upd (inert_irreds ics) }
delIrred :: IrredCt -> InertCans -> InertCans
-- Remove a particular (Given) Irred, on the instructions of a plugin
-- For some reason this is done vis the evidence Id, not the type
-- Compare delEq. I have not idea why
delIrred (IrredCt { ir_ev = ev }) ics
= updIrreds (filterBag keep) ics
where
ev_id = ctEvEvId ev
keep (IrredCt { ir_ev = ev' }) = ev_id /= ctEvEvId ev'
foldIrreds :: (IrredCt -> b -> b) -> InertIrreds -> b -> b
foldIrreds k irreds z = foldr k z irreds
findMatchingIrreds :: InertIrreds -> CtEvidence
-> (Bag (IrredCt, SwapFlag), InertIrreds)
findMatchingIrreds irreds ev
| EqPred eq_rel1 lty1 rty1 <- classifyPredType pred
-- See Note [Solving irreducible equalities]
= partitionBagWith (match_eq eq_rel1 lty1 rty1) irreds
| otherwise
= partitionBagWith match_non_eq irreds
where
pred = ctEvPred ev
match_non_eq irred
| irredCtPred irred `tcEqType` pred = Left (irred, NotSwapped)
| otherwise = Right irred
match_eq eq_rel1 lty1 rty1 irred
| EqPred eq_rel2 lty2 rty2 <- classifyPredType (irredCtPred irred)
, eq_rel1 == eq_rel2
, Just swap <- match_eq_help lty1 rty1 lty2 rty2
= Left (irred, swap)
| otherwise
= Right irred
match_eq_help lty1 rty1 lty2 rty2
| lty1 `tcEqType` lty2, rty1 `tcEqType` rty2
= Just NotSwapped
| lty1 `tcEqType` rty2, rty1 `tcEqType` lty2
= Just IsSwapped
| otherwise
= Nothing
{- Note [Solving irreducible equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider (#14333)
[G] a b ~R# c d
[W] c d ~R# a b
Clearly we should be able to solve this! Even though the constraints are
not decomposable. We solve this when looking up the work-item in the
irreducible constraints to look for an identical one. When doing this
lookup, findMatchingIrreds spots the equality case, and matches either
way around. It has to return a swap-flag so we can generate evidence
that is the right way round too.
-}
{- *********************************************************************
* *
Adding to and removing from the inert set
* *
* *
********************************************************************* -}
updGivenEqs :: TcLevel -> Ct -> InertCans -> InertCans
-- Set the inert_given_eq_level to the current level (tclvl)
-- if the constraint is a given equality that should prevent
-- filling in an outer unification variable.
-- See Note [Tracking Given equalities]
--
-- Precondition: Ct is either CEqCan or CIrredCan
updGivenEqs tclvl ct inerts
| not (isGivenCt ct) = inerts
-- See Note [Let-bound skolems]
| isLetBoundSkolemCt tclvl ct = inerts { inert_given_eqs = True }
-- At this point we are left with a constraint that either
-- is an equality (F a ~ ty), or /might/ be, like (c a)
| otherwise = inerts { inert_given_eq_lvl = tclvl
, inert_given_eqs = True }
isLetBoundSkolemCt :: TcLevel -> Ct -> Bool
-- See Note [Let-bound skolems]
isLetBoundSkolemCt tclvl (CEqCan (EqCt { eq_lhs = lhs, eq_rhs = rhs }))
= case lhs of
TyVarLHS tv -> not (isOuterTyVar tclvl tv)
TyFamLHS {} -> case getTyVar_maybe rhs of
Just tv -> not (isOuterTyVar tclvl tv)
Nothing -> False
isLetBoundSkolemCt _ _ = False
data KickOutSpec -- See Note [KickOutSpec]
= KOAfterUnify TcTyVarSet -- We have unified these tyvars
| KOAfterAdding CanEqLHS -- We are adding to the inert set a canonical equality
-- constraint with this LHS
instance Outputable KickOutSpec where
ppr (KOAfterUnify tvs) = text "KOAfterUnify" <> ppr tvs
ppr (KOAfterAdding lhs) = text "KOAfterAdding" <> parens (ppr lhs)
{- Note [KickOutSpec]
~~~~~~~~~~~~~~~~~~~~~~
KickOutSpec explains why we are kicking out.
Important property:
KOAfterAdding (TyVarLHS tv) should behave exactly like
KOAfterUnifying (unitVarSet tv)
The main reasons for treating the two separately are
* More efficient in the single-tyvar case
* The code is far more perspicuous
-}
data WhereToLook = LookEverywhere | LookOnlyUnderFamApps
deriving( Eq )
kickOutRewritableLHS :: KickOutSpec -> CtFlavourRole -> InertCans -> (Cts, InertCans)
-- See Note [kickOutRewritable]
kickOutRewritableLHS ko_spec new_fr@(_, new_role)
ics@(IC { inert_eqs = tv_eqs
, inert_dicts = dictmap
, inert_funeqs = funeqmap
, inert_irreds = irreds
, inert_insts = old_insts })
= (kicked_out, inert_cans_in)
where
-- inert_safehask stays unchanged; is that right?
inert_cans_in = ics { inert_eqs = tv_eqs_in
, inert_dicts = dicts_in
, inert_funeqs = feqs_in
, inert_irreds = irs_in
, inert_insts = insts_in }
kicked_out :: Cts
kicked_out = (fmap CDictCan dicts_out `andCts` fmap CIrredCan irs_out)
`extendCtsList` insts_out
`extendCtsList` map CEqCan tv_eqs_out
`extendCtsList` map CEqCan feqs_out
(tv_eqs_out, tv_eqs_in) = partitionInertEqs kick_out_eq tv_eqs
(feqs_out, feqs_in) = partitionFunEqs kick_out_eq funeqmap
(dicts_out, dicts_in) = partitionDicts kick_out_dict dictmap
(irs_out, irs_in) = partitionBag kick_out_irred irreds
-- Kick out even insolubles: See Note [Rewrite insolubles]
-- Of course we must kick out irreducibles like (c a), in case
-- we can rewrite 'c' to something more useful
-- Kick-out for inert instances
-- See Note [Quantified constraints] in GHC.Tc.Solver.Solve
insts_out :: [Ct]
insts_in :: [QCInst]
(insts_out, insts_in)
| fr_may_rewrite (Given, NomEq) -- All the insts are Givens
= partitionWith kick_out_qci old_insts
| otherwise
= ([], old_insts)
kick_out_qci qci
| let ev = qci_ev qci
, fr_can_rewrite_ty LookEverywhere NomEq (ctEvPred (qci_ev qci))
= Left (mkNonCanonical ev)
| otherwise
= Right qci
fr_tv_can_rewrite_ty :: WhereToLook -> (TyVar -> Bool) -> EqRel -> Type -> Bool
fr_tv_can_rewrite_ty where_to_look check_tv role ty
= anyRewritableTyVar role can_rewrite ty
where
can_rewrite :: UnderFam -> EqRel -> TyVar -> Bool
can_rewrite is_under_famapp old_role tv
= (where_to_look == LookEverywhere || is_under_famapp) &&
new_role `eqCanRewrite` old_role && check_tv tv
fr_tf_can_rewrite_ty :: WhereToLook -> TyCon -> [TcType] -> EqRel -> Type -> Bool
fr_tf_can_rewrite_ty where_to_look new_tf new_tf_args role ty
= anyRewritableTyFamApp role can_rewrite ty
where
can_rewrite :: UnderFam -> EqRel -> TyCon -> [TcType] -> Bool
can_rewrite is_under_famapp old_role old_tf old_tf_args
= (where_to_look == LookEverywhere || is_under_famapp) &&
new_role `eqCanRewrite` old_role &&
tcEqTyConApps new_tf new_tf_args old_tf old_tf_args
-- it's possible for old_tf_args to have too many. This is fine;
-- we'll only check what we need to.
fr_can_rewrite_ty :: WhereToLook -> EqRel -> Type -> Bool
-- UnderFam = True <=> look only under type-family applications
fr_can_rewrite_ty uf = case ko_spec of -- See Note [KickOutSpec]
KOAfterUnify tvs -> fr_tv_can_rewrite_ty uf (`elemVarSet` tvs)
KOAfterAdding (TyVarLHS tv) -> fr_tv_can_rewrite_ty uf (== tv)
KOAfterAdding (TyFamLHS tf tf_args) -> fr_tf_can_rewrite_ty uf tf tf_args
fr_may_rewrite :: CtFlavourRole -> Bool
fr_may_rewrite fs = new_fr `eqCanRewriteFR` fs
-- Can the new item rewrite the inert item?
kick_out_dict :: DictCt -> Bool
-- Kick it out if the new CEqCan can rewrite the inert one
-- See Note [kickOutRewritable]
kick_out_dict (DictCt { di_tys = tys, di_ev = ev })
= fr_may_rewrite (ctEvFlavour ev, NomEq)
&& any (fr_can_rewrite_ty LookEverywhere NomEq) tys
kick_out_irred :: IrredCt -> Bool
kick_out_irred (IrredCt { ir_ev = ev })
= fr_may_rewrite (ctEvFlavour ev, eq_rel)
&& fr_can_rewrite_ty LookEverywhere eq_rel pred
where
pred = ctEvPred ev
eq_rel = predTypeEqRel pred
-- Implements criteria K1-K3 in Note [Extending the inert equalities]
kick_out_eq :: EqCt -> Bool
kick_out_eq (EqCt { eq_lhs = lhs, eq_rhs = rhs_ty
, eq_ev = ev, eq_eq_rel = eq_rel })
-- (KK0) Keep it in the inert set if the new thing can't rewrite it
| not (fr_may_rewrite fs)
= False
-- Below here (fr_may_rewrite fs) is True
-- (KK1)
| fr_can_rewrite_ty LookEverywhere eq_rel (canEqLHSType lhs)
= True -- (KK1)
-- The above check redundantly checks the role & flavour,
-- but it's very convenient
-- (KK2)
| let where_to_look | fs_can_rewrite_fr = LookOnlyUnderFamApps
| otherwise = LookEverywhere
, fr_can_rewrite_ty where_to_look eq_rel rhs_ty
= True
-- (KK3)
| not fs_can_rewrite_fr -- (KK3a)
, case eq_rel of
NomEq -> is_new_lhs rhs_ty -- (KK3b)
ReprEq -> head_is_new_lhs rhs_ty -- (KK3c)
= True
| otherwise = False
where
fs_can_rewrite_fr = fs `eqCanRewriteFR` new_fr
fs = (ctEvFlavour ev, eq_rel)
is_new_lhs :: Type -> Bool
is_new_lhs = case ko_spec of -- See Note [KickOutSpec]
KOAfterUnify tvs -> is_tyvar_ty_for tvs
KOAfterAdding lhs -> (`eqType` canEqLHSType lhs)
is_tyvar_ty_for :: TcTyVarSet -> Type -> Bool
-- True if the type is equal to one of the tyvars
is_tyvar_ty_for tvs ty
= case getTyVar_maybe ty of
Nothing -> False
Just tv -> tv `elemVarSet` tvs
head_is_new_lhs :: Type -> Bool
head_is_new_lhs = case ko_spec of -- See Note [KickOutSpec]
KOAfterUnify tvs -> tv_at_head (`elemVarSet` tvs)
KOAfterAdding (TyVarLHS tv) -> tv_at_head (== tv)
KOAfterAdding (TyFamLHS tf tf_args) -> fam_at_head tf tf_args
tv_at_head :: (TyVar -> Bool) -> Type -> Bool
tv_at_head is_tv = go
where
go (Rep.TyVarTy tv) = is_tv tv
go (Rep.AppTy fun _) = go fun
go (Rep.CastTy ty _) = go ty
go (Rep.TyConApp {}) = False
go (Rep.LitTy {}) = False
go (Rep.ForAllTy {}) = False
go (Rep.FunTy {}) = False
go (Rep.CoercionTy {}) = False
fam_at_head :: TyCon -> [Type] -> Type -> Bool
fam_at_head fun_tc fun_args = go
where
go (Rep.TyVarTy {}) = False
go (Rep.AppTy {}) = False -- no TyConApp to the left of an AppTy
go (Rep.CastTy ty _) = go ty
go (Rep.TyConApp tc args) = tcEqTyConApps fun_tc fun_args tc args
go (Rep.LitTy {}) = False
go (Rep.ForAllTy {}) = False
go (Rep.FunTy {}) = False
go (Rep.CoercionTy {}) = False
{- Note [kickOutRewritable]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
See also Note [inert_eqs: the inert equalities].
When we add a new inert equality (lhs ~N ty) to the inert set,
we must kick out any inert items that could be rewritten by the
new equality, to maintain the inert-set invariants.
- We want to kick out an existing inert constraint if
a) the new constraint can rewrite the inert one
b) 'lhs' is free in the inert constraint (so that it *will*)
rewrite it if we kick it out.
For (b) we use anyRewritableCanLHS, which examines the types /and
kinds/ that are directly visible in the type. Hence
we will have exposed all the rewriting we care about to make the
most precise kinds visible for matching classes etc. No need to
kick out constraints that mention type variables whose kinds
contain this LHS!
- We don't kick out constraints from inert_solved_dicts, and
inert_solved_funeqs optimistically. But when we lookup we have to
take the substitution into account
NB: we could in principle avoid kick-out:
a) When unifying a meta-tyvar from an outer level, because
then the entire implication will be iterated; see
Note [The Unification Level Flag] in GHC.Tc.Solver.Monad.
b) For Givens, after a unification. By (GivenInv) in GHC.Tc.Utils.TcType
Note [TcLevel invariants], a Given can't include a meta-tyvar from
its own level, so it falls under (a). Of course, we must still
kick out Givens when adding a new non-unification Given.
But kicking out more vigorously may lead to earlier unification and fewer
iterations, so we don't take advantage of these possibilities.
Note [Rewrite insolubles]
~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have an insoluble alpha ~ [alpha], which is insoluble
because an occurs check. And then we unify alpha := [Int]. Then we
really want to rewrite the insoluble to [Int] ~ [[Int]]. Now it can
be decomposed. Otherwise we end up with a "Can't match [Int] ~
[[Int]]" which is true, but a bit confusing because the outer type
constructors match.
Hence:
* In the main simplifier loops in GHC.Tc.Solver (solveWanteds,
simpl_loop), we feed the insolubles in solveSimpleWanteds,
so that they get rewritten (albeit not solved).
* We kick insolubles out of the inert set, if they can be
rewritten (see GHC.Tc.Solver.Monad.kick_out_rewritable)
* We rewrite those insolubles in GHC.Tc.Solver.Equality
See Note [Make sure that insolubles are fully rewritten]
in GHC.Tc.Solver.Equality
-}
{- *********************************************************************
* *
Queries
* *
* *
********************************************************************* -}
isOuterTyVar :: TcLevel -> TyCoVar -> Bool
-- True of a type variable that comes from a
-- shallower level than the ambient level (tclvl)
isOuterTyVar tclvl tv
| isTyVar tv = assertPpr (not (isTouchableMetaTyVar tclvl tv)) (ppr tv <+> ppr tclvl) $
tclvl `strictlyDeeperThan` tcTyVarLevel tv
-- ASSERT: we are dealing with Givens here, and invariant (GivenInv) from
-- Note Note [TcLevel invariants] in GHC.Tc.Utils.TcType ensures that there can't
-- be a touchable meta tyvar. If this wasn't true, you might worry that,
-- at level 3, a meta-tv alpha[3] gets unified with skolem b[2], and thereby
-- becomes "outer" even though its level numbers says it isn't.
| otherwise = False -- Coercion variables; doesn't much matter
noGivenNewtypeReprEqs :: TyCon -> InertSet -> Bool
-- True <=> there is no Irred looking like (N tys1 ~ N tys2)
-- See Note [Decomposing newtype equalities] (EX2) in GHC.Tc.Solver.Equality
-- This is the only call site.
noGivenNewtypeReprEqs tc inerts
= not (anyBag might_help (inert_irreds (inert_cans inerts)))
where
might_help irred
= case classifyPredType (ctEvPred (irredCtEvidence irred)) of
EqPred ReprEq t1 t2
| Just (tc1,_) <- tcSplitTyConApp_maybe t1
, tc == tc1
, Just (tc2,_) <- tcSplitTyConApp_maybe t2
, tc == tc2
-> True
_ -> False
-- | Returns True iff there are no Given constraints that might,
-- potentially, match the given class constraint. This is used when checking to see if a
-- Given might overlap with an instance. See Note [Instance and Given overlap]
-- in GHC.Tc.Solver.Dict
noMatchableGivenDicts :: InertSet -> CtLoc -> Class -> [TcType] -> Bool
noMatchableGivenDicts inerts@(IS { inert_cans = inert_cans }) loc_w clas tys
= not $ anyBag matchable_given $
findDictsByClass (inert_dicts inert_cans) clas
where
pred_w = mkClassPred clas tys
matchable_given :: DictCt -> Bool
matchable_given (DictCt { di_ev = ev })
| CtGiven { ctev_loc = loc_g, ctev_pred = pred_g } <- ev
= isJust $ mightEqualLater inerts pred_g loc_g pred_w loc_w
| otherwise
= False
mightEqualLater :: InertSet -> TcPredType -> CtLoc -> TcPredType -> CtLoc -> Maybe Subst
-- See Note [What might equal later?]
-- Used to implement logic in Note [Instance and Given overlap] in GHC.Tc.Solver.Dict
mightEqualLater inert_set given_pred given_loc wanted_pred wanted_loc
| prohibitedSuperClassSolve given_loc wanted_loc
= Nothing
| otherwise
= case tcUnifyTysFG bind_fun [flattened_given] [flattened_wanted] of
Unifiable subst
-> Just subst
MaybeApart reason subst
| MARInfinite <- reason -- see Example 7 in the Note.
-> Nothing
| otherwise
-> Just subst
SurelyApart -> Nothing
where
in_scope = mkInScopeSet $ tyCoVarsOfTypes [given_pred, wanted_pred]
-- NB: flatten both at the same time, so that we can share mappings
-- from type family applications to variables, and also to guarantee
-- that the fresh variables are really fresh between the given and
-- the wanted. Flattening both at the same time is needed to get
-- Example 10 from the Note.
([flattened_given, flattened_wanted], var_mapping)
= flattenTysX in_scope [given_pred, wanted_pred]
bind_fun :: BindFun
bind_fun tv rhs_ty
| isMetaTyVar tv
, can_unify tv (metaTyVarInfo tv) rhs_ty
-- this checks for CycleBreakerTvs and TyVarTvs; forgetting
-- the latter was #19106.
= BindMe
-- See Examples 4, 5, and 6 from the Note
| Just (_fam_tc, fam_args) <- lookupVarEnv var_mapping tv
, anyFreeVarsOfTypes mentions_meta_ty_var fam_args
= BindMe
| otherwise
= Apart
-- True for TauTv and TyVarTv (and RuntimeUnkTv) meta-tyvars
-- (as they can be unified)
-- and also for CycleBreakerTvs that mentions meta-tyvars
mentions_meta_ty_var :: TyVar -> Bool
mentions_meta_ty_var tv
| isMetaTyVar tv
= case metaTyVarInfo tv of
-- See Examples 8 and 9 in the Note
CycleBreakerTv
-> anyFreeVarsOfType mentions_meta_ty_var
(lookupCycleBreakerVar tv inert_set)
_ -> True
| otherwise
= False
-- Like checkTopShape, but allows cbv variables to unify
can_unify :: TcTyVar -> MetaInfo -> Type -> Bool
can_unify _lhs_tv TyVarTv rhs_ty -- see Example 3 from the Note
| Just rhs_tv <- getTyVar_maybe rhs_ty
= case tcTyVarDetails rhs_tv of
MetaTv { mtv_info = TyVarTv } -> True
MetaTv {} -> False -- Could unify with anything
SkolemTv {} -> True
RuntimeUnk -> True
| otherwise -- not a var on the RHS
= False
can_unify lhs_tv _other _rhs_ty = mentions_meta_ty_var lhs_tv
-- | Is it (potentially) loopy to use the first @ct1@ to solve @ct2@?
--
-- Necessary (but not sufficient) conditions for this function to return @True@:
--
-- - @ct1@ and @ct2@ both arise from superclass expansion,
-- - @ct1@ is a Given and @ct2@ is a Wanted.
--
-- See Note [Solving superclass constraints] in GHC.Tc.TyCl.Instance, (sc2).
prohibitedSuperClassSolve :: CtLoc -- ^ is it loopy to use this one ...
-> CtLoc -- ^ ... to solve this one?
-> Bool -- ^ True ==> don't solve it
prohibitedSuperClassSolve given_loc wanted_loc
| GivenSCOrigin _ _ blocked <- ctLocOrigin given_loc
, blocked
, ScOrigin _ NakedSc <- ctLocOrigin wanted_loc
= True -- Prohibited if the Wanted is a superclass
-- and the Given has come via a superclass selection from
-- a predicate bigger than the head
| otherwise
= False
{- Note [What might equal later?]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We must determine whether a Given might later equal a Wanted. We
definitely need to account for the possibility that any metavariable
might be arbitrarily instantiated. Yet we do *not* want
to allow skolems in to be instantiated, as we've already rewritten
with respect to any Givens. (We're solving a Wanted here, and so
all Givens have already been processed.)
This is best understood by example.
1. C alpha ~? C Int
That given certainly might match later.
2. C a ~? C Int
No. No new givens are going to arise that will get the `a` to rewrite
to Int.
3. C alpha[tv] ~? C Int
That alpha[tv] is a TyVarTv, unifiable only with other type variables.
It cannot equal later.
4. C (F alpha) ~? C Int
Sure -- that can equal later, if we learn something useful about alpha.
5. C (F alpha[tv]) ~? C Int
This, too, might equal later. Perhaps we have [G] F b ~ Int elsewhere.
Or maybe we have C (F alpha[tv] beta[tv]), these unify with each other,
and F x x = Int. Remember: returning True doesn't commit ourselves to
anything.
6. C (F a) ~? C a
No, this won't match later. If we could rewrite (F a) or a, we would
have by now. But see also Red Herring below.
7. C (Maybe alpha) ~? C alpha
We say this cannot equal later, because it would require
alpha := Maybe (Maybe (Maybe ...)). While such a type can be contrived,
we choose not to worry about it. See Note [Infinitary substitution in lookup]
in GHC.Core.InstEnv. Getting this wrong let to #19107, tested in
typecheck/should_compile/T19107.
8. C cbv ~? C Int
where cbv = F a
The cbv is a cycle-breaker var which stands for F a. See
Note [Type equality cycles] in GHC.Tc.Solver.Equality
This is just like case 6, and we say "no". Saying "no" here is
essential in getting the parser to type-check, with its use of DisambECP.
9. C cbv ~? C Int
where cbv = F alpha
Here, we might indeed equal later. Distinguishing between
this case and Example 8 is why we need the InertSet in mightEqualLater.
10. C (F alpha, Int) ~? C (Bool, F alpha)
This cannot equal later, because F a would have to equal both Bool and
Int.
To deal with type family applications, we use the Core flattener. See
Note [Flattening type-family applications when matching instances] in GHC.Core.Unify.
The Core flattener replaces all type family applications with
fresh variables. The next question: should we allow these fresh
variables in the domain of a unifying substitution?
A type family application that mentions only skolems (example 6) is settled:
any skolems would have been rewritten w.r.t. Givens by now. These type family
applications match only themselves. A type family application that mentions
metavariables, on the other hand, can match anything. So, if the original type
family application contains a metavariable, we use BindMe to tell the unifier
to allow it in the substitution. On the other hand, a type family application
with only skolems is considered rigid.
This treatment fixes #18910 and is tested in
typecheck/should_compile/InstanceGivenOverlap{,2}
Red Herring
~~~~~~~~~~~
In #21208, we have this scenario:
instance forall b. C b
[G] C a[sk]
[W] C (F a[sk])
What should we do with that wanted? According to the logic above, the Given
cannot match later (this is example 6), and so we use the global instance.
But wait, you say: What if we learn later (say by a future type instance F a = a)
that F a unifies with a? That looks like the Given might really match later!
This mechanism described in this Note is *not* about this kind of situation, however.
It is all asking whether a Given might match the Wanted *in this run of the solver*.
It is *not* about whether a variable might be instantiated so that the Given matches,
or whether a type instance introduced in a downstream module might make the Given match.
The reason we care about what might match later is only about avoiding order-dependence.
That is, we don't want to commit to a course of action that depends on seeing constraints
in a certain order. But an instantiation of a variable and a later type instance
don't introduce order dependency in this way, and so mightMatchLater is right to ignore
these possibilities.
Here is an example, with no type families, that is perhaps clearer:
instance forall b. C (Maybe b)
[G] C (Maybe Int)
[W] C (Maybe a)
What to do? We *might* say that the Given could match later and should thus block
us from using the global instance. But we don't do this. Instead, we rely on class
coherence to say that choosing the global instance is just fine, even if later we
call a function with (a := Int). After all, in this run of the solver, [G] C (Maybe Int)
will definitely never match [W] C (Maybe a). (Recall that we process Givens before
Wanteds, so there is no [G] a ~ Int hanging about unseen.)
Interestingly, in the first case (from #21208), the behavior changed between
GHC 8.10.7 and GHC 9.2, with the latter behaving correctly and the former
reporting overlapping instances.
Test case: typecheck/should_compile/T21208.
-}
{- *********************************************************************
* *
Cycle breakers
* *
********************************************************************* -}
-- | Return the type family application a CycleBreakerTv maps to.
lookupCycleBreakerVar :: TcTyVar -- ^ cbv, must be a CycleBreakerTv
-> InertSet
-> TcType -- ^ type family application the cbv maps to
lookupCycleBreakerVar cbv (IS { inert_cycle_breakers = cbvs_stack })
-- This function looks at every environment in the stack. This is necessary
-- to avoid #20231. This function (and its one usage site) is the only reason
-- that we store a stack instead of just the top environment.
| Just tyfam_app <- assert (isCycleBreakerTyVar cbv) $
firstJusts (NE.map (lookupBag cbv) cbvs_stack)
= tyfam_app
| otherwise
= pprPanic "lookupCycleBreakerVar found an unbound cycle breaker" (ppr cbv $$ ppr cbvs_stack)
-- | Push a fresh environment onto the cycle-breaker var stack. Useful
-- when entering a nested implication.
pushCycleBreakerVarStack :: CycleBreakerVarStack -> CycleBreakerVarStack
pushCycleBreakerVarStack = (emptyBag <|)
-- | Add a new cycle-breaker binding to the top environment on the stack.
addCycleBreakerBindings :: Bag (TcTyVar, Type) -- ^ (cbv,expansion) pairs
-> InertSet -> InertSet
addCycleBreakerBindings prs ics
= assertPpr (all (isCycleBreakerTyVar . fst) prs) (ppr prs) $
ics { inert_cycle_breakers = add_to (inert_cycle_breakers ics) }
where
add_to (top_env :| rest_envs) = (prs `unionBags` top_env) :| rest_envs
-- | Perform a monadic operation on all pairs in the top environment
-- in the stack.
forAllCycleBreakerBindings_ :: Monad m
=> CycleBreakerVarStack
-> (TcTyVar -> TcType -> m ()) -> m ()
forAllCycleBreakerBindings_ (top_env :| _rest_envs) action
= forM_ top_env (uncurry action)
{-# INLINABLE forAllCycleBreakerBindings_ #-} -- to allow SPECIALISE later
{- *********************************************************************
* *
Solving one from another
* *
********************************************************************* -}
data InteractResult
= KeepInert -- Keep the inert item, and solve the work item from it
-- (if the latter is Wanted; just discard it if not)
| KeepWork -- Keep the work item, and solve the inert item from it
instance Outputable InteractResult where
ppr KeepInert = text "keep inert"
ppr KeepWork = text "keep work-item"
solveOneFromTheOther :: Ct -- Inert (Dict or Irred)
-> Ct -- WorkItem (same predicate as inert)
-> InteractResult
-- Precondition:
-- * inert and work item represent evidence for the /same/ predicate
-- * Both are CDictCan or CIrredCan
--
-- We can always solve one from the other: even if both are wanted,
-- although we don't rewrite wanteds with wanteds, we can combine
-- two wanteds into one by solving one from the other
solveOneFromTheOther ct_i ct_w
| CtWanted { ctev_loc = loc_w } <- ev_w
, prohibitedSuperClassSolve loc_i loc_w
-- See Note [Solving superclass constraints] in GHC.Tc.TyCl.Instance
= -- Inert must be Given
KeepWork
| CtWanted {} <- ev_w
= -- Inert is Given or Wanted
case ev_i of
CtGiven {} -> KeepInert
-- work is Wanted; inert is Given: easy choice.
CtWanted {} -- Both are Wanted
-- If only one has no pending superclasses, use it
-- Otherwise we can get infinite superclass expansion (#22516)
-- in silly cases like class C T b => C a b where ...
| not is_psc_i, is_psc_w -> KeepInert
| is_psc_i, not is_psc_w -> KeepWork
-- If only one is a WantedSuperclassOrigin (arising from expanding
-- a Wanted class constraint), keep the other: wanted superclasses
-- may be unexpected by users
| not is_wsc_orig_i, is_wsc_orig_w -> KeepInert
| is_wsc_orig_i, not is_wsc_orig_w -> KeepWork
-- otherwise, just choose the lower span
-- reason: if we have something like (abs 1) (where the
-- Num constraint cannot be satisfied), it's better to
-- get an error about abs than about 1.
-- This test might become more elaborate if we see an
-- opportunity to improve the error messages
| ((<) `on` ctLocSpan) loc_i loc_w -> KeepInert
| otherwise -> KeepWork
-- From here on the work-item is Given
| CtWanted { ctev_loc = loc_i } <- ev_i
, prohibitedSuperClassSolve loc_w loc_i
= KeepInert -- Just discard the un-usable Given
-- This never actually happens because
-- Givens get processed first
| CtWanted {} <- ev_i
= KeepWork
-- From here on both are Given
-- See Note [Replacement vs keeping]
| lvl_i `sameDepthAs` lvl_w
= same_level_strategy
| otherwise -- Both are Given, levels differ
= different_level_strategy
where
ev_i = ctEvidence ct_i
ev_w = ctEvidence ct_w
pred = ctEvPred ev_i
loc_i = ctEvLoc ev_i
loc_w = ctEvLoc ev_w
orig_i = ctLocOrigin loc_i
orig_w = ctLocOrigin loc_w
lvl_i = ctLocLevel loc_i
lvl_w = ctLocLevel loc_w
is_psc_w = isPendingScDict ct_w
is_psc_i = isPendingScDict ct_i
is_wsc_orig_i = isWantedSuperclassOrigin orig_i
is_wsc_orig_w = isWantedSuperclassOrigin orig_w
different_level_strategy -- Both Given
| isIPLikePred pred = if lvl_w `strictlyDeeperThan` lvl_i then KeepWork else KeepInert
| otherwise = if lvl_w `strictlyDeeperThan` lvl_i then KeepInert else KeepWork
-- See Note [Replacement vs keeping] part (1)
-- For the isIPLikePred case see Note [Shadowing of implicit parameters]
-- in GHC.Tc.Solver.Dict
same_level_strategy -- Both Given
= case (orig_i, orig_w) of
(GivenSCOrigin _ depth_i blocked_i, GivenSCOrigin _ depth_w blocked_w)
| blocked_i, not blocked_w -> KeepWork -- Case 2(a) from
| not blocked_i, blocked_w -> KeepInert -- Note [Replacement vs keeping]
-- Both blocked or both not blocked
| depth_w < depth_i -> KeepWork -- Case 2(c) from
| otherwise -> KeepInert -- Note [Replacement vs keeping]
(GivenSCOrigin {}, _) -> KeepWork -- Case 2(b) from Note [Replacement vs keeping]
_ -> KeepInert -- Case 2(d) from Note [Replacement vs keeping]
{-
Note [Replacement vs keeping]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When we have two Given constraints both of type (C tys), say, which should
we keep? More subtle than you might think! This is all implemented in
solveOneFromTheOther.
1) Constraints come from different levels (different_level_strategy)
- For implicit parameters we want to keep the innermost (deepest)
one, so that it overrides the outer one.
See Note [Shadowing of implicit parameters] in GHC.Tc.Solver.Dict
- For everything else, we want to keep the outermost one. Reason: that
makes it more likely that the inner one will turn out to be unused,
and can be reported as redundant. See Note [Tracking redundant constraints]
in GHC.Tc.Solver.
It transpires that using the outermost one is responsible for an
8% performance improvement in nofib cryptarithm2, compared to
just rolling the dice. I didn't investigate why.
2) Constraints coming from the same level (i.e. same implication)
(a) If both are GivenSCOrigin, choose the one that is unblocked if possible
according to Note [Solving superclass constraints] in GHC.Tc.TyCl.Instance.
(b) Prefer constraints that are not superclass selections. Example:
f :: (Eq a, Ord a) => a -> Bool
f x = x == x
Eager superclass expansion gives us two [G] Eq a constraints. We
want to keep the one from the user-written Eq a, not the superclass
selection. This means we report the Ord a as redundant with
-Wredundant-constraints, not the Eq a.
Getting this wrong was #20602. See also
Note [Tracking redundant constraints] in GHC.Tc.Solver.
(c) If both are GivenSCOrigin, chooose the one with the shallower
superclass-selection depth, in the hope of identifying more correct
redundant constraints. This is really a generalization of point (b),
because the superclass depth of a non-superclass constraint is 0.
(If the levels differ, we definitely won't have both with GivenSCOrigin.)
(d) Finally, when there is still a choice, use KeepInert rather than
KeepWork, for two reasons:
- to avoid unnecessary munging of the inert set.
- to cut off superclass loops; see Note [Superclass loops] in GHC.Tc.Solver.Dict
Doing the level-check for implicit parameters, rather than making the work item
always override, is important. Consider
data T a where { T1 :: (?x::Int) => T Int; T2 :: T a }
f :: (?x::a) => T a -> Int
f T1 = ?x
f T2 = 3
We have a [G] (?x::a) in the inert set, and at the pattern match on T1 we add
two new givens in the work-list: [G] (?x::Int)
[G] (a ~ Int)
Now consider these steps
- process a~Int, kicking out (?x::a)
- process (?x::Int), the inner given, adding to inert set
- process (?x::a), the outer given, overriding the inner given
Wrong! The level-check ensures that the inner implicit parameter wins.
(Actually I think that the order in which the work-list is processed means
that this chain of events won't happen, but that's very fragile.)
-}